diff options
Diffstat (limited to 'Documentation')
113 files changed, 10051 insertions, 1115 deletions
diff --git a/Documentation/ABI/testing/sysfs-class-net-cdc_ncm b/Documentation/ABI/testing/sysfs-class-net-cdc_ncm index 5cedf72df358..f7be0e88b139 100644 --- a/Documentation/ABI/testing/sysfs-class-net-cdc_ncm +++ b/Documentation/ABI/testing/sysfs-class-net-cdc_ncm @@ -19,6 +19,25 @@ Description: Set to 0 to pad all frames. Set greater than tx_max to disable all padding. +What: /sys/class/net/<iface>/cdc_ncm/ndp_to_end +Date: Dec 2015 +KernelVersion: 4.5 +Contact: Bjørn Mork <bjorn@mork.no> +Description: + Boolean attribute showing the status of the "NDP to + end" quirk. Defaults to 'N', except for devices + already known to need it enabled. + + The "NDP to end" quirk makes the driver place the NDP + (the packet index table) after the payload. The NCM + specification does not mandate this, but some devices + are known to be more restrictive. Write 'Y' to this + attribute for temporary testing of a suspect device + failing to work with the default driver settings. + + A device entry should be added to the driver if this + quirk is found to be required. + What: /sys/class/net/<iface>/cdc_ncm/rx_max Date: May 2014 KernelVersion: 3.16 diff --git a/Documentation/ABI/testing/sysfs-class-net-mesh b/Documentation/ABI/testing/sysfs-class-net-mesh index c46406296631..c2b956d44a95 100644 --- a/Documentation/ABI/testing/sysfs-class-net-mesh +++ b/Documentation/ABI/testing/sysfs-class-net-mesh @@ -8,7 +8,7 @@ Description: What: /sys/class/net/<mesh_iface>/mesh/<vlan_subdir>/ap_isolation Date: May 2011 -Contact: Antonio Quartulli <antonio@meshcoding.com> +Contact: Antonio Quartulli <a@unstable.cc> Description: Indicates whether the data traffic going from a wireless client to another wireless client will be @@ -70,7 +70,7 @@ Description: What: /sys/class/net/<mesh_iface>/mesh/isolation_mark Date: Nov 2013 -Contact: Antonio Quartulli <antonio@meshcoding.com> +Contact: Antonio Quartulli <a@unstable.cc> Description: Defines the isolation mark (and its bitmask) which is used to classify clients as "isolated" by the diff --git a/Documentation/ABI/testing/sysfs-class-net-qmi b/Documentation/ABI/testing/sysfs-class-net-qmi new file mode 100644 index 000000000000..fa5a00bb1143 --- /dev/null +++ b/Documentation/ABI/testing/sysfs-class-net-qmi @@ -0,0 +1,23 @@ +What: /sys/class/net/<iface>/qmi/raw_ip +Date: Dec 2015 +KernelVersion: 4.4 +Contact: Bjørn Mork <bjorn@mork.no> +Description: + Boolean. Default: 'N' + + Set this to 'Y' to change the network device link + framing from '802.3' to 'raw-ip'. + + The netdev will change to reflect the link framing + mode. The netdev is an ordinary ethernet device in + '802.3' mode, and the driver expects to exchange + frames with an ethernet header over the USB link. The + netdev is a headerless p-t-p device in 'raw-ip' mode, + and the driver expects to echange IPv4 or IPv6 packets + without any L2 header over the USB link. + + Userspace is in full control of firmware configuration + through the delegation of the QMI protocol. Userspace + is responsible for coordination of driver and firmware + link framing mode, changing this setting to 'Y' if the + firmware is configured for 'raw-ip' mode. diff --git a/Documentation/DocBook/device-drivers.tmpl b/Documentation/DocBook/device-drivers.tmpl index 42a2d8593e39..7b3fcc5effcd 100644 --- a/Documentation/DocBook/device-drivers.tmpl +++ b/Documentation/DocBook/device-drivers.tmpl @@ -238,78 +238,26 @@ X!Isound/sound_firmware.c !Iinclude/media/videobuf2-memops.h </sect1> <sect1><title>Digital TV (DVB) devices</title> -!Idrivers/media/dvb-core/dvb_ca_en50221.h -!Idrivers/media/dvb-core/dvb_frontend.h + <sect1><title>Digital TV Common functions</title> !Idrivers/media/dvb-core/dvb_math.h !Idrivers/media/dvb-core/dvb_ringbuffer.h !Idrivers/media/dvb-core/dvbdev.h - <sect1><title>Digital TV Demux API</title> - <para>The kernel demux API defines a driver-internal interface for - registering low-level, hardware specific driver to a hardware - independent demux layer. It is only of interest for Digital TV - device driver writers. The header file for this API is named - <constant>demux.h</constant> and located in - <constant>drivers/media/dvb-core</constant>.</para> - - <para>The demux API should be implemented for each demux in the - system. It is used to select the TS source of a demux and to manage - the demux resources. When the demux client allocates a resource via - the demux API, it receives a pointer to the API of that - resource.</para> - <para>Each demux receives its TS input from a DVB front-end or from - memory, as set via this demux API. In a system with more than one - front-end, the API can be used to select one of the DVB front-ends - as a TS source for a demux, unless this is fixed in the HW platform. - The demux API only controls front-ends regarding to their connections - with demuxes; the APIs used to set the other front-end parameters, - such as tuning, are not defined in this document.</para> - <para>The functions that implement the abstract interface demux should - be defined static or module private and registered to the Demux - core for external access. It is not necessary to implement every - function in the struct <constant>dmx_demux</constant>. For example, - a demux interface might support Section filtering, but not PES - filtering. The API client is expected to check the value of any - function pointer before calling the function: the value of NULL means - that the “function is not available”.</para> - <para>Whenever the functions of the demux API modify shared data, - the possibilities of lost update and race condition problems should - be addressed, e.g. by protecting parts of code with mutexes.</para> - <para>Note that functions called from a bottom half context must not - sleep. Even a simple memory allocation without using GFP_ATOMIC can - result in a kernel thread being put to sleep if swapping is needed. - For example, the Linux kernel calls the functions of a network device - interface from a bottom half context. Thus, if a demux API function - is called from network device code, the function must not sleep. - </para> - </sect1> - - <section id="demux_callback_api"> - <title>Demux Callback API</title> - <para>This kernel-space API comprises the callback functions that - deliver filtered data to the demux client. Unlike the other DVB - kABIs, these functions are provided by the client and called from - the demux code.</para> - <para>The function pointers of this abstract interface are not - packed into a structure as in the other demux APIs, because the - callback functions are registered and used independent of each - other. As an example, it is possible for the API client to provide - several callback functions for receiving TS packets and no - callbacks for PES packets or sections.</para> - <para>The functions that implement the callback API need not be - re-entrant: when a demux driver calls one of these functions, - the driver is not allowed to call the function again before - the original call returns. If a callback is triggered by a - hardware interrupt, it is recommended to use the Linux - “bottom half” mechanism or start a tasklet instead of - making the callback function call directly from a hardware - interrupt.</para> - <para>This mechanism is implemented by - <link linkend='API-dmx-ts-cb'>dmx_ts_cb()</link> and - <link linkend='API-dmx-section-cb'>dmx_section_cb()</link>.</para> - </section> - + </sect1> + <sect1><title>Digital TV Frontend kABI</title> +!Pdrivers/media/dvb-core/dvb_frontend.h Digital TV Frontend +!Idrivers/media/dvb-core/dvb_frontend.h + </sect1> + <sect1><title>Digital TV Demux kABI</title> +!Pdrivers/media/dvb-core/demux.h Digital TV Demux + <sect1><title>Demux Callback API</title> +!Pdrivers/media/dvb-core/demux.h Demux Callback + </sect1> !Idrivers/media/dvb-core/demux.h - </sect1> + </sect1> + <sect1><title>Digital TV Conditional Access kABI</title> +!Idrivers/media/dvb-core/dvb_ca_en50221.h + </sect1> + </sect1> <sect1><title>Remote Controller devices</title> !Iinclude/media/rc-core.h !Iinclude/media/lirc_dev.h diff --git a/Documentation/DocBook/media/Makefile b/Documentation/DocBook/media/Makefile index 08527e7ea4d0..2840ff483d5a 100644 --- a/Documentation/DocBook/media/Makefile +++ b/Documentation/DocBook/media/Makefile @@ -199,8 +199,10 @@ DVB_DOCUMENTED = \ # install_media_images = \ - $(Q)-mkdir $(MEDIA_OBJ_DIR)/media_api; \ - cp $(OBJIMGFILES) $(MEDIA_SRC_DIR)/*.svg $(MEDIA_SRC_DIR)/v4l/*.svg $(MEDIA_OBJ_DIR)/media_api + $(Q)if [ "x$(findstring media_api.xml,$(DOCBOOKS))" != "x" ]; then \ + mkdir -p $(MEDIA_OBJ_DIR)/media_api; \ + cp $(OBJIMGFILES) $(MEDIA_SRC_DIR)/*.svg $(MEDIA_SRC_DIR)/v4l/*.svg $(MEDIA_OBJ_DIR)/media_api; \ + fi $(MEDIA_OBJ_DIR)/%: $(MEDIA_SRC_DIR)/%.b64 $(Q)base64 -d $< >$@ diff --git a/Documentation/DocBook/media/dvb/dvbproperty.xml b/Documentation/DocBook/media/dvb/dvbproperty.xml index 08227d4e9150..e579ae5088ae 100644 --- a/Documentation/DocBook/media/dvb/dvbproperty.xml +++ b/Documentation/DocBook/media/dvb/dvbproperty.xml @@ -76,7 +76,7 @@ int main(void) <para>NOTE: While it is possible to directly call the Kernel code like the above example, it is strongly recommended to use - <ulink url="http://linuxtv.org/docs/libdvbv5/index.html">libdvbv5</ulink>, + <ulink url="https://linuxtv.org/docs/libdvbv5/index.html">libdvbv5</ulink>, as it provides abstraction to work with the supported digital TV standards and provides methods for usual operations like program scanning and to read/write channel descriptor files.</para> diff --git a/Documentation/DocBook/media/dvb/examples.xml b/Documentation/DocBook/media/dvb/examples.xml index c9f68c7183cc..837fb3b64b72 100644 --- a/Documentation/DocBook/media/dvb/examples.xml +++ b/Documentation/DocBook/media/dvb/examples.xml @@ -3,7 +3,7 @@ </para> <para>NOTE: This section is out of date, and the code below won't even compile. Please refer to the - <ulink url="http://linuxtv.org/docs/libdvbv5/index.html">libdvbv5</ulink> + <ulink url="https://linuxtv.org/docs/libdvbv5/index.html">libdvbv5</ulink> for updated/recommended examples. </para> diff --git a/Documentation/DocBook/media/dvb/intro.xml b/Documentation/DocBook/media/dvb/intro.xml index 51db15648099..b5b701f5d8c2 100644 --- a/Documentation/DocBook/media/dvb/intro.xml +++ b/Documentation/DocBook/media/dvb/intro.xml @@ -32,7 +32,7 @@ and filtering several section and PES data streams at the same time. new standard Linux DVB API. As a commitment to the development of terminals based on open standards, Nokia and Convergence made it available to all Linux developers and published it on -<ulink url="http://www.linuxtv.org/" /> in September 2000. +<ulink url="https://linuxtv.org" /> in September 2000. Convergence is the maintainer of the Linux DVB API. Together with the LinuxTV community (i.e. you, the reader of this document), the Linux DVB API will be constantly reviewed and improved. With the Linux driver for diff --git a/Documentation/DocBook/media/v4l/capture.c.xml b/Documentation/DocBook/media/v4l/capture.c.xml index 1c5c49a2de59..22126a991b34 100644 --- a/Documentation/DocBook/media/v4l/capture.c.xml +++ b/Documentation/DocBook/media/v4l/capture.c.xml @@ -5,7 +5,7 @@ * This program can be used and distributed without restrictions. * * This program is provided with the V4L2 API - * see http://linuxtv.org/docs.php for more information + * see https://linuxtv.org/docs.php for more information */ #include <stdio.h> diff --git a/Documentation/DocBook/media/v4l/compat.xml b/Documentation/DocBook/media/v4l/compat.xml index 5701a08ed792..5399e8904715 100644 --- a/Documentation/DocBook/media/v4l/compat.xml +++ b/Documentation/DocBook/media/v4l/compat.xml @@ -2666,7 +2666,7 @@ is useful to display images captured with V4L2 devices.</para> <para>V4L2 does not support digital terrestrial, cable or satellite broadcast. A separate project aiming at digital receivers exists. You can find its homepage at <ulink -url="http://linuxtv.org">http://linuxtv.org</ulink>. The Linux DVB API +url="https://linuxtv.org">https://linuxtv.org</ulink>. The Linux DVB API has no connection to the V4L2 API except that drivers for hybrid hardware may support both.</para> </section> diff --git a/Documentation/DocBook/media/v4l/io.xml b/Documentation/DocBook/media/v4l/io.xml index da654031ef3f..144158b3a5ac 100644 --- a/Documentation/DocBook/media/v4l/io.xml +++ b/Documentation/DocBook/media/v4l/io.xml @@ -699,7 +699,7 @@ linkend="v4l2-buf-type" /></entry> buffer. It depends on the negotiated data format and may change with each buffer for compressed variable size data like JPEG images. Drivers must set this field when <structfield>type</structfield> -refers to an input stream, applications when it refers to an output stream. +refers to a capture stream, applications when it refers to an output stream. If the application sets this to 0 for an output stream, then <structfield>bytesused</structfield> will be set to the size of the buffer (see the <structfield>length</structfield> field of this struct) by @@ -720,14 +720,14 @@ linkend="buffer-flags" />.</entry> <entry>Indicates the field order of the image in the buffer, see <xref linkend="v4l2-field" />. This field is not used when the buffer contains VBI data. Drivers must set it when -<structfield>type</structfield> refers to an input stream, +<structfield>type</structfield> refers to a capture stream, applications when it refers to an output stream.</entry> </row> <row> <entry>struct timeval</entry> <entry><structfield>timestamp</structfield></entry> <entry></entry> - <entry><para>For input streams this is time when the first data + <entry><para>For capture streams this is time when the first data byte was captured, as returned by the <function>clock_gettime()</function> function for the relevant clock id; see <constant>V4L2_BUF_FLAG_TIMESTAMP_*</constant> in @@ -866,7 +866,7 @@ must set this to 0.</entry> <entry></entry> <entry>The number of bytes occupied by data in the plane (its payload). Drivers must set this field when <structfield>type</structfield> - refers to an input stream, applications when it refers to an output stream. + refers to a capture stream, applications when it refers to an output stream. If the application sets this to 0 for an output stream, then <structfield>bytesused</structfield> will be set to the size of the plane (see the <structfield>length</structfield> field of this struct) @@ -919,7 +919,7 @@ must set this to 0.</entry> <entry></entry> <entry>Offset in bytes to video data in the plane. Drivers must set this field when <structfield>type</structfield> - refers to an input stream, applications when it refers to an output stream. + refers to a capture stream, applications when it refers to an output stream. Note that data_offset is included in <structfield>bytesused</structfield>. So the size of the image in the plane is <structfield>bytesused</structfield>-<structfield>data_offset</structfield> at diff --git a/Documentation/DocBook/media/v4l/v4l2.xml b/Documentation/DocBook/media/v4l/v4l2.xml index 7e61643358de..42e626d6c936 100644 --- a/Documentation/DocBook/media/v4l/v4l2.xml +++ b/Documentation/DocBook/media/v4l/v4l2.xml @@ -152,6 +152,16 @@ structs, ioctls) must be noted in more detail in the history chapter (compat.xml), along with the possible impact on existing drivers and applications. --> <revision> + <revnumber>4.5</revnumber> + <date>2015-10-29</date> + <authorinitials>rr</authorinitials> + <revremark>Extend vidioc-g-ext-ctrls;. Replace ctrl_class with a new +union with ctrl_class and which. Which is used to select the current value of +the control or the default value. + </revremark> + </revision> + + <revision> <revnumber>4.4</revnumber> <date>2015-05-26</date> <authorinitials>ap</authorinitials> diff --git a/Documentation/DocBook/media/v4l/vidioc-create-bufs.xml b/Documentation/DocBook/media/v4l/vidioc-create-bufs.xml index 8ffe74f84af1..d81fa0d4016b 100644 --- a/Documentation/DocBook/media/v4l/vidioc-create-bufs.xml +++ b/Documentation/DocBook/media/v4l/vidioc-create-bufs.xml @@ -58,7 +58,7 @@ <para>This ioctl is used to create buffers for <link linkend="mmap">memory mapped</link> or <link linkend="userp">user pointer</link> or <link linkend="dmabuf">DMA buffer</link> I/O. It can be used as an alternative or in -addition to the <constant>VIDIOC_REQBUFS</constant> ioctl, when a tighter +addition to the &VIDIOC-REQBUFS; ioctl, when a tighter control over buffers is required. This ioctl can be called multiple times to create buffers of different sizes.</para> @@ -71,30 +71,28 @@ zeroed.</para> <para>The <structfield>format</structfield> field specifies the image format that the buffers must be able to handle. The application has to fill in this -&v4l2-format;. Usually this will be done using the -<constant>VIDIOC_TRY_FMT</constant> or <constant>VIDIOC_G_FMT</constant> ioctl() -to ensure that the requested format is supported by the driver. Unsupported -formats will result in an error.</para> +&v4l2-format;. Usually this will be done using the &VIDIOC-TRY-FMT; or &VIDIOC-G-FMT; ioctls +to ensure that the requested format is supported by the driver. +Based on the format's <structfield>type</structfield> field the requested buffer +size (for single-planar) or plane sizes (for multi-planar formats) will be +used for the allocated buffers. The driver may return an error if the size(s) +are not supported by the hardware (usually because they are too small).</para> <para>The buffers created by this ioctl will have as minimum size the size -defined by the <structfield>format.pix.sizeimage</structfield> field. If the +defined by the <structfield>format.pix.sizeimage</structfield> field (or the +corresponding fields for other format types). Usually if the <structfield>format.pix.sizeimage</structfield> field is less than the minimum -required for the given format, then <structfield>sizeimage</structfield> will be -increased by the driver to that minimum to allocate the buffers. If it is -larger, then the value will be used as-is. The same applies to the -<structfield>sizeimage</structfield> field of the -<structname>v4l2_plane_pix_format</structname> structure in the case of -multiplanar formats.</para> +required for the given format, then an error will be returned since drivers will +typically not allow this. If it is larger, then the value will be used as-is. +In other words, the driver may reject the requested size, but if it is accepted +the driver will use it unchanged.</para> <para>When the ioctl is called with a pointer to this structure the driver will attempt to allocate up to the requested number of buffers and store the actual number allocated and the starting index in the <structfield>count</structfield> and the <structfield>index</structfield> fields respectively. On return <structfield>count</structfield> can be smaller than -the number requested. The driver may also increase buffer sizes if required, -however, it will not update <structfield>sizeimage</structfield> field values. -The user has to use <constant>VIDIOC_QUERYBUF</constant> to retrieve that -information.</para> +the number requested.</para> <table pgwide="1" frame="none" id="v4l2-create-buffers"> <title>struct <structname>v4l2_create_buffers</structname></title> diff --git a/Documentation/DocBook/media/v4l/vidioc-dbg-g-chip-info.xml b/Documentation/DocBook/media/v4l/vidioc-dbg-g-chip-info.xml index 4c4603c135fe..f14a3bb1afaa 100644 --- a/Documentation/DocBook/media/v4l/vidioc-dbg-g-chip-info.xml +++ b/Documentation/DocBook/media/v4l/vidioc-dbg-g-chip-info.xml @@ -99,7 +99,7 @@ if the driver supports writing registers to the device.</para> <para>We recommended the <application>v4l2-dbg</application> utility over calling this ioctl directly. It is available from the LinuxTV v4l-dvb repository; see <ulink -url="http://linuxtv.org/repo/">http://linuxtv.org/repo/</ulink> for +url="https://linuxtv.org/repo/">https://linuxtv.org/repo/</ulink> for access instructions.</para> <!-- Note for convenience vidioc-dbg-g-register.sgml diff --git a/Documentation/DocBook/media/v4l/vidioc-dbg-g-register.xml b/Documentation/DocBook/media/v4l/vidioc-dbg-g-register.xml index 3d038e75d12b..5877f68a5820 100644 --- a/Documentation/DocBook/media/v4l/vidioc-dbg-g-register.xml +++ b/Documentation/DocBook/media/v4l/vidioc-dbg-g-register.xml @@ -117,7 +117,7 @@ However when a driver supports these ioctls it must also support <para>We recommended the <application>v4l2-dbg</application> utility over calling these ioctls directly. It is available from the LinuxTV v4l-dvb repository; see <ulink -url="http://linuxtv.org/repo/">http://linuxtv.org/repo/</ulink> for +url="https://linuxtv.org/repo/">https://linuxtv.org/repo/</ulink> for access instructions.</para> <!-- Note for convenience vidioc-dbg-g-chip-info.sgml diff --git a/Documentation/DocBook/media/v4l/vidioc-enumstd.xml b/Documentation/DocBook/media/v4l/vidioc-enumstd.xml index 8065099401d1..f18454e91752 100644 --- a/Documentation/DocBook/media/v4l/vidioc-enumstd.xml +++ b/Documentation/DocBook/media/v4l/vidioc-enumstd.xml @@ -198,7 +198,7 @@ video4linux-list@redhat.com on 17 Oct 2002 <constant>V4L2_STD_ATSC_16_VSB</constant> are U.S. terrestrial digital TV standards. Presently the V4L2 API does not support digital TV. See also the Linux DVB API at <ulink -url="http://linuxtv.org">http://linuxtv.org</ulink>.</para> +url="https://linuxtv.org">https://linuxtv.org</ulink>.</para> <para><programlisting> #define V4L2_STD_PAL_BG (V4L2_STD_PAL_B |\ V4L2_STD_PAL_B1 |\ diff --git a/Documentation/DocBook/media/v4l/vidioc-g-ext-ctrls.xml b/Documentation/DocBook/media/v4l/vidioc-g-ext-ctrls.xml index 842536aae8b4..eb82f7e7d06b 100644 --- a/Documentation/DocBook/media/v4l/vidioc-g-ext-ctrls.xml +++ b/Documentation/DocBook/media/v4l/vidioc-g-ext-ctrls.xml @@ -61,7 +61,7 @@ must belong to the same control class.</para> <para>Applications must always fill in the <structfield>count</structfield>, -<structfield>ctrl_class</structfield>, +<structfield>which</structfield>, <structfield>controls</structfield> and <structfield>reserved</structfield> fields of &v4l2-ext-controls;, and initialize the &v4l2-ext-control; array pointed to by the @@ -109,7 +109,7 @@ the driver whether wrong values are automatically adjusted to a valid value or if an error is returned.</para> <para>When the <structfield>id</structfield> or -<structfield>ctrl_class</structfield> is invalid drivers return an +<structfield>which</structfield> is invalid drivers return an &EINVAL;. When the value is out of bounds drivers can choose to take the closest valid value or return an &ERANGE;, whatever seems more appropriate. In the first case the new value is set in @@ -223,7 +223,12 @@ Valid if <constant>V4L2_CTRL_FLAG_HAS_PAYLOAD</constant> is set for this control <tgroup cols="3"> &cs-str; <tbody valign="top"> + <row> + <entry>union</entry> + <entry>(anonymous)</entry> + </row> <row> + <entry></entry> <entry>__u32</entry> <entry><structfield>ctrl_class</structfield></entry> <entry>The control class to which all controls belong, see @@ -235,6 +240,23 @@ with a <structfield>count</structfield> of 0. If that succeeds, then the driver supports this feature.</entry> </row> <row> + <entry></entry> + <entry>__u32</entry> + <entry><structfield>which</structfield></entry> + <entry><para>Which value of the control to get/set/try. <constant>V4L2_CTRL_WHICH_CUR_VAL</constant> +will return the current value of the control and <constant>V4L2_CTRL_WHICH_DEF_VAL</constant> will +return the default value of the control. Please note that you can only get the default value of the +control, you cannot set or try it.</para> +<para>For backwards compatibility you can also use a control class here (see +<xref linkend="ctrl-class" />). In that case all controls have to belong to that +control class. This usage is deprecated, instead just use <constant>V4L2_CTRL_WHICH_CUR_VAL</constant>. +There are some very old drivers that do not yet support <constant>V4L2_CTRL_WHICH_CUR_VAL</constant> +and that require a control class here. You can test for such drivers by setting ctrl_class to +<constant>V4L2_CTRL_WHICH_CUR_VAL</constant> and calling VIDIOC_TRY_EXT_CTRLS with a count of 0. +If that fails, then the driver does not support <constant>V4L2_CTRL_WHICH_CUR_VAL</constant>.</para> +</entry> + </row> + <row> <entry>__u32</entry> <entry><structfield>count</structfield></entry> <entry>The number of controls in the controls array. May @@ -390,7 +412,7 @@ These controls are described in <xref linkend="rf-tuner-controls" />.</entry> <listitem> <para>The &v4l2-ext-control; <structfield>id</structfield> is invalid, the &v4l2-ext-controls; -<structfield>ctrl_class</structfield> is invalid, or the &v4l2-ext-control; +<structfield>which</structfield> is invalid, or the &v4l2-ext-control; <structfield>value</structfield> was inappropriate (e.g. the given menu index is not supported by the driver). This error code is also returned by the <constant>VIDIOC_S_EXT_CTRLS</constant> and diff --git a/Documentation/DocBook/media_api.tmpl b/Documentation/DocBook/media_api.tmpl index 92037033f5eb..7b77e0f7b87d 100644 --- a/Documentation/DocBook/media_api.tmpl +++ b/Documentation/DocBook/media_api.tmpl @@ -19,10 +19,10 @@ <!ENTITY cs-def "<colspec colname='c1' colwidth='3*' /><colspec colname='c2' colwidth='1*' /><colspec colname='c3' colwidth='4*' /><spanspec spanname='hspan' namest='c1' nameend='c3' />"> <!-- Video for Linux mailing list address. --> -<!ENTITY v4l-ml "<ulink url='http://www.linuxtv.org/lists.php'>http://www.linuxtv.org/lists.php</ulink>"> +<!ENTITY v4l-ml "<ulink url='https://linuxtv.org/lists.php'>https://linuxtv.org/lists.php</ulink>"> <!-- LinuxTV v4l-dvb repository. --> -<!ENTITY v4l-dvb "<ulink url='http://linuxtv.org/repo/'>http://linuxtv.org/repo/</ulink>"> +<!ENTITY v4l-dvb "<ulink url='https://linuxtv.org/repo/'>https://linuxtv.org/repo/</ulink>"> <!ENTITY dash-ent-8 "<entry>-</entry><entry>-</entry><entry>-</entry><entry>-</entry><entry>-</entry><entry>-</entry><entry>-</entry><entry>-</entry>"> <!ENTITY dash-ent-10 "<entry>-</entry><entry>-</entry><entry>-</entry><entry>-</entry><entry>-</entry><entry>-</entry><entry>-</entry><entry>-</entry><entry>-</entry><entry>-</entry>"> <!ENTITY dash-ent-12 "<entry>-</entry><entry>-</entry><entry>-</entry><entry>-</entry><entry>-</entry><entry>-</entry><entry>-</entry><entry>-</entry><entry>-</entry><entry>-</entry><entry>-</entry><entry>-</entry>"> @@ -91,7 +91,7 @@ components, like mixers, PCM capture, PCM playback, etc, which are controlled via ALSA API.</para> <para>For additional information and for the latest development code, - see: <ulink url="http://linuxtv.org">http://linuxtv.org</ulink>.</para> + see: <ulink url="https://linuxtv.org">https://linuxtv.org</ulink>.</para> <para>For discussing improvements, reporting troubles, sending new drivers, etc, please mail to: <ulink 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L;-inkscape-font-specification:Nimbus Sans L" + id="tspan3017">Read-Mostly, Stale &</tspan></text> + <!-- Text --> + <text + xml:space="preserve" + x="7200" + y="1350" + font-style="normal" + font-weight="normal" + font-size="324" + id="text16" + sodipodi:linespacing="125%" + style="font-size:427.63009644px;font-style:normal;font-variant:normal;font-weight:normal;font-stretch:normal;text-align:center;line-height:125%;writing-mode:lr-tb;text-anchor:middle;fill:#000000;font-family:Nimbus Sans L;-inkscape-font-specification:Nimbus Sans L"><tspan + style="font-size:427.63009644px;font-style:normal;font-variant:normal;font-weight:normal;font-stretch:normal;text-align:center;line-height:125%;writing-mode:lr-tb;text-anchor:middle;font-family:Nimbus Sans L;-inkscape-font-specification:Nimbus Sans L" + id="tspan3019">Inconsistent Data OK</tspan></text> + <!-- Text --> + <text + xml:space="preserve" + x="7200" + y="1800" + font-style="normal" + font-weight="normal" + font-size="324" + id="text18" + sodipodi:linespacing="125%" + style="font-size:427.63009644px;font-style:normal;font-variant:normal;font-weight:normal;font-stretch:normal;text-align:center;line-height:125%;writing-mode:lr-tb;text-anchor:middle;fill:#000000;font-family:Nimbus Sans L;-inkscape-font-specification:Nimbus Sans L"><tspan + style="font-size:427.63009644px;font-style:normal;font-variant:normal;font-weight:normal;font-stretch:normal;text-align:center;line-height:125%;writing-mode:lr-tb;text-anchor:middle;font-family:Nimbus Sans L;-inkscape-font-specification:Nimbus Sans L" + id="tspan3021">(RCU Works Great!!!)</tspan></text> + <!-- Text --> + <text + xml:space="preserve" + x="7200" + y="3825" + font-style="normal" + font-weight="normal" + font-size="324" + id="text20" + sodipodi:linespacing="125%" + style="font-size:427.63009644px;font-style:normal;font-variant:normal;font-weight:normal;font-stretch:normal;text-align:center;line-height:125%;writing-mode:lr-tb;text-anchor:middle;fill:#000000;font-family:Nimbus Sans L;-inkscape-font-specification:Nimbus Sans L"><tspan + style="font-size:427.63009644px;font-style:normal;font-variant:normal;font-weight:normal;font-stretch:normal;text-align:center;line-height:125%;writing-mode:lr-tb;text-anchor:middle;font-family:Nimbus Sans L;-inkscape-font-specification:Nimbus Sans L" + id="tspan3023">(RCU Works Well)</tspan></text> + <!-- Text --> + <text + xml:space="preserve" + x="7200" + y="3375" + font-style="normal" + font-weight="normal" + font-size="324" + id="text22" + sodipodi:linespacing="125%" + style="font-size:427.63009644px;font-style:normal;font-variant:normal;font-weight:normal;font-stretch:normal;text-align:center;line-height:125%;writing-mode:lr-tb;text-anchor:middle;fill:#000000;font-family:Nimbus Sans L;-inkscape-font-specification:Nimbus Sans L"><tspan + style="font-size:427.63009644px;font-style:normal;font-variant:normal;font-weight:normal;font-stretch:normal;text-align:center;line-height:125%;writing-mode:lr-tb;text-anchor:middle;font-family:Nimbus Sans L;-inkscape-font-specification:Nimbus Sans L" + id="tspan3025">Read-Mostly, Need Consistent Data</tspan></text> + <!-- Text --> + <text + xml:space="preserve" + x="7200" + y="5175" + font-style="normal" + font-weight="normal" + font-size="324" + id="text24" + sodipodi:linespacing="125%" + style="font-size:427.63009644px;font-style:normal;font-variant:normal;font-weight:normal;font-stretch:normal;text-align:center;line-height:125%;writing-mode:lr-tb;text-anchor:middle;fill:#000000;font-family:Nimbus Sans L;-inkscape-font-specification:Nimbus Sans L"><tspan + style="font-size:427.63009644px;font-style:normal;font-variant:normal;font-weight:normal;font-stretch:normal;text-align:center;line-height:125%;writing-mode:lr-tb;text-anchor:middle;font-family:Nimbus Sans L;-inkscape-font-specification:Nimbus Sans L" + id="tspan3027">Read-Write, Need Consistent Data</tspan></text> + <!-- Text --> + <text + xml:space="preserve" + x="7200" + y="6975" + font-style="normal" + font-weight="normal" + font-size="324" + id="text26" + style="font-size:427.63009644px;font-style:normal;font-variant:normal;font-weight:normal;font-stretch:normal;text-align:center;line-height:125%;writing-mode:lr-tb;text-anchor:middle;fill:#000000;font-family:Nimbus Sans L;-inkscape-font-specification:Nimbus Sans L" + sodipodi:linespacing="125%">Update-Mostly, Need Consistent Data</text> + <!-- Text --> + <text + xml:space="preserve" + x="7200" + y="5625" + font-style="normal" + font-weight="normal" + font-size="324" + id="text28" + sodipodi:linespacing="125%" + style="font-size:427.63009644px;font-style:normal;font-variant:normal;font-weight:normal;font-stretch:normal;text-align:center;line-height:125%;writing-mode:lr-tb;text-anchor:middle;fill:#000000;font-family:Nimbus Sans L;-inkscape-font-specification:Nimbus Sans L"><tspan + style="font-size:427.63009644px;font-style:normal;font-variant:normal;font-weight:normal;font-stretch:normal;text-align:center;line-height:125%;writing-mode:lr-tb;text-anchor:middle;font-family:Nimbus Sans L;-inkscape-font-specification:Nimbus Sans L" + id="tspan3029">(RCU Might Be OK...)</tspan></text> + <!-- Text --> + <text + xml:space="preserve" + x="7200" + y="7875" + font-style="normal" + font-weight="normal" + font-size="324" + id="text30" + style="font-size:427.63009644px;font-style:normal;font-variant:normal;font-weight:normal;font-stretch:normal;text-align:center;line-height:125%;writing-mode:lr-tb;text-anchor:middle;fill:#000000;font-family:Nimbus Sans L;-inkscape-font-specification:Nimbus Sans L" + sodipodi:linespacing="125%">(1) Provide Existence Guarantees For Update-Friendly Mechanisms</text> + <!-- Text --> + <text + xml:space="preserve" + x="7200" + y="8325" + font-style="normal" + font-weight="normal" + font-size="324" + id="text32" + style="font-size:427.63009644px;font-style:normal;font-variant:normal;font-weight:normal;font-stretch:normal;text-align:center;line-height:125%;writing-mode:lr-tb;text-anchor:middle;fill:#000000;font-family:Nimbus Sans L;-inkscape-font-specification:Nimbus Sans L" + sodipodi:linespacing="125%">(2) Provide Wait-Free Read-Side Primitives for Real-Time Use)</text> + <!-- Text --> + <text + xml:space="preserve" + x="7200" + y="7425" + font-style="normal" + font-weight="normal" + font-size="324" + id="text34" + style="font-size:427.63009644px;font-style:normal;font-variant:normal;font-weight:normal;font-stretch:normal;text-align:center;line-height:125%;writing-mode:lr-tb;text-anchor:middle;fill:#000000;font-family:Nimbus Sans 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y="-261.66608" + x="436.12299" + id="tspan2995" + sodipodi:role="line">synchronize_rcu()</tspan></text> + <g + id="g4417" + transform="matrix(0,1,-1,0,730.90257,222.4928)"> + <path + style="fill:none;stroke:#000000;stroke-width:1px;stroke-linecap:butt;stroke-linejoin:miter;stroke-opacity:1;marker-start:url(#Arrow2Lstart);marker-end:url(#Arrow2Lend)" + d="M 97.580736,477.4048 327.57913,476.09759" + id="path2997" + inkscape:connector-curvature="0" + sodipodi:nodetypes="cc" /> + <path + style="fill:none;stroke:#000000;stroke-width:1px;stroke-linecap:butt;stroke-linejoin:miter;stroke-opacity:1" + d="m 96.752718,465.38398 0,22.62742" + id="path4397" + inkscape:connector-curvature="0" + sodipodi:nodetypes="cc" /> + <path + style="fill:none;stroke:#000000;stroke-width:1px;stroke-linecap:butt;stroke-linejoin:miter;stroke-opacity:1" + d="m 328.40703,465.38397 0,22.62742" + id="path4397-5" + inkscape:connector-curvature="0" + sodipodi:nodetypes="cc" /> + </g> + </g> + <text + 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x="684.00067" + y="213.91006" + id="text4461-2-1" + sodipodi:linespacing="125%"><tspan + sodipodi:role="line" + id="tspan4463-2-0" + x="684.00067" + y="213.91006">thread4()</tspan></text> + </g> +</svg> diff --git a/Documentation/RCU/Design/Requirements/Requirements.html b/Documentation/RCU/Design/Requirements/Requirements.html new file mode 100644 index 000000000000..a725f9900ec8 --- /dev/null +++ b/Documentation/RCU/Design/Requirements/Requirements.html @@ -0,0 +1,2897 @@ +<!-- DO NOT HAND EDIT. --> +<!-- Instead, edit Documentation/RCU/Design/Requirements/Requirements.htmlx and run 'sh htmlqqz.sh Documentation/RCU/Design/Requirements/Requirements' --> +<!DOCTYPE HTML PUBLIC "-//W3C//DTD HTML 4.01 Transitional//EN" + "http://www.w3.org/TR/html4/loose.dtd"> + <html> + <head><title>A Tour Through RCU's Requirements [LWN.net]</title> + <meta HTTP-EQUIV="Content-Type" CONTENT="text/html; charset=utf-8"> + +<h1>A Tour Through RCU's Requirements</h1> + +<p>Copyright IBM Corporation, 2015</p> +<p>Author: Paul E. McKenney</p> +<p><i>The initial version of this document appeared in the +<a href="https://lwn.net/">LWN</a> articles +<a href="https://lwn.net/Articles/652156/">here</a>, +<a href="https://lwn.net/Articles/652677/">here</a>, and +<a href="https://lwn.net/Articles/653326/">here</a>.</i></p> + +<h2>Introduction</h2> + +<p> +Read-copy update (RCU) is a synchronization mechanism that is often +used as a replacement for reader-writer locking. +RCU is unusual in that updaters do not block readers, +which means that RCU's read-side primitives can be exceedingly fast +and scalable. +In addition, updaters can make useful forward progress concurrently +with readers. +However, all this concurrency between RCU readers and updaters does raise +the question of exactly what RCU readers are doing, which in turn +raises the question of exactly what RCU's requirements are. + +<p> +This document therefore summarizes RCU's requirements, and can be thought +of as an informal, high-level specification for RCU. +It is important to understand that RCU's specification is primarily +empirical in nature; +in fact, I learned about many of these requirements the hard way. +This situation might cause some consternation, however, not only +has this learning process been a lot of fun, but it has also been +a great privilege to work with so many people willing to apply +technologies in interesting new ways. + +<p> +All that aside, here are the categories of currently known RCU requirements: +</p> + +<ol> +<li> <a href="#Fundamental Requirements"> + Fundamental Requirements</a> +<li> <a href="#Fundamental Non-Requirements">Fundamental Non-Requirements</a> +<li> <a href="#Parallelism Facts of Life"> + Parallelism Facts of Life</a> +<li> <a href="#Quality-of-Implementation Requirements"> + Quality-of-Implementation Requirements</a> +<li> <a href="#Linux Kernel Complications"> + Linux Kernel Complications</a> +<li> <a href="#Software-Engineering Requirements"> + Software-Engineering Requirements</a> +<li> <a href="#Other RCU Flavors"> + Other RCU Flavors</a> +<li> <a href="#Possible Future Changes"> + Possible Future Changes</a> +</ol> + +<p> +This is followed by a <a href="#Summary">summary</a>, +which is in turn followed by the inevitable +<a href="#Answers to Quick Quizzes">answers to the quick quizzes</a>. + +<h2><a name="Fundamental Requirements">Fundamental Requirements</a></h2> + +<p> +RCU's fundamental requirements are the closest thing RCU has to hard +mathematical requirements. +These are: + +<ol> +<li> <a href="#Grace-Period Guarantee"> + Grace-Period Guarantee</a> +<li> <a href="#Publish-Subscribe Guarantee"> + Publish-Subscribe Guarantee</a> +<li> <a href="#Memory-Barrier Guarantees"> + Memory-Barrier Guarantees</a> +<li> <a href="#RCU Primitives Guaranteed to Execute Unconditionally"> + RCU Primitives Guaranteed to Execute Unconditionally</a> +<li> <a href="#Guaranteed Read-to-Write Upgrade"> + Guaranteed Read-to-Write Upgrade</a> +</ol> + +<h3><a name="Grace-Period Guarantee">Grace-Period Guarantee</a></h3> + +<p> +RCU's grace-period guarantee is unusual in being premeditated: +Jack Slingwine and I had this guarantee firmly in mind when we started +work on RCU (then called “rclock”) in the early 1990s. +That said, the past two decades of experience with RCU have produced +a much more detailed understanding of this guarantee. + +<p> +RCU's grace-period guarantee allows updaters to wait for the completion +of all pre-existing RCU read-side critical sections. +An RCU read-side critical section +begins with the marker <tt>rcu_read_lock()</tt> and ends with +the marker <tt>rcu_read_unlock()</tt>. +These markers may be nested, and RCU treats a nested set as one +big RCU read-side critical section. +Production-quality implementations of <tt>rcu_read_lock()</tt> and +<tt>rcu_read_unlock()</tt> are extremely lightweight, and in +fact have exactly zero overhead in Linux kernels built for production +use with <tt>CONFIG_PREEMPT=n</tt>. + +<p> +This guarantee allows ordering to be enforced with extremely low +overhead to readers, for example: + +<blockquote> +<pre> + 1 int x, y; + 2 + 3 void thread0(void) + 4 { + 5 rcu_read_lock(); + 6 r1 = READ_ONCE(x); + 7 r2 = READ_ONCE(y); + 8 rcu_read_unlock(); + 9 } +10 +11 void thread1(void) +12 { +13 WRITE_ONCE(x, 1); +14 synchronize_rcu(); +15 WRITE_ONCE(y, 1); +16 } +</pre> +</blockquote> + +<p> +Because the <tt>synchronize_rcu()</tt> on line 14 waits for +all pre-existing readers, any instance of <tt>thread0()</tt> that +loads a value of zero from <tt>x</tt> must complete before +<tt>thread1()</tt> stores to <tt>y</tt>, so that instance must +also load a value of zero from <tt>y</tt>. +Similarly, any instance of <tt>thread0()</tt> that loads a value of +one from <tt>y</tt> must have started after the +<tt>synchronize_rcu()</tt> started, and must therefore also load +a value of one from <tt>x</tt>. +Therefore, the outcome: +<blockquote> +<pre> +(r1 == 0 && r2 == 1) +</pre> +</blockquote> +cannot happen. + +<p><a name="Quick Quiz 1"><b>Quick Quiz 1</b>:</a> +Wait a minute! +You said that updaters can make useful forward progress concurrently +with readers, but pre-existing readers will block +<tt>synchronize_rcu()</tt>!!! +Just who are you trying to fool??? +<br><a href="#qq1answer">Answer</a> + +<p> +This scenario resembles one of the first uses of RCU in +<a href="https://en.wikipedia.org/wiki/DYNIX">DYNIX/ptx</a>, +which managed a distributed lock manager's transition into +a state suitable for handling recovery from node failure, +more or less as follows: + +<blockquote> +<pre> + 1 #define STATE_NORMAL 0 + 2 #define STATE_WANT_RECOVERY 1 + 3 #define STATE_RECOVERING 2 + 4 #define STATE_WANT_NORMAL 3 + 5 + 6 int state = STATE_NORMAL; + 7 + 8 void do_something_dlm(void) + 9 { +10 int state_snap; +11 +12 rcu_read_lock(); +13 state_snap = READ_ONCE(state); +14 if (state_snap == STATE_NORMAL) +15 do_something(); +16 else +17 do_something_carefully(); +18 rcu_read_unlock(); +19 } +20 +21 void start_recovery(void) +22 { +23 WRITE_ONCE(state, STATE_WANT_RECOVERY); +24 synchronize_rcu(); +25 WRITE_ONCE(state, STATE_RECOVERING); +26 recovery(); +27 WRITE_ONCE(state, STATE_WANT_NORMAL); +28 synchronize_rcu(); +29 WRITE_ONCE(state, STATE_NORMAL); +30 } +</pre> +</blockquote> + +<p> +The RCU read-side critical section in <tt>do_something_dlm()</tt> +works with the <tt>synchronize_rcu()</tt> in <tt>start_recovery()</tt> +to guarantee that <tt>do_something()</tt> never runs concurrently +with <tt>recovery()</tt>, but with little or no synchronization +overhead in <tt>do_something_dlm()</tt>. + +<p><a name="Quick Quiz 2"><b>Quick Quiz 2</b>:</a> +Why is the <tt>synchronize_rcu()</tt> on line 28 needed? +<br><a href="#qq2answer">Answer</a> + +<p> +In order to avoid fatal problems such as deadlocks, +an RCU read-side critical section must not contain calls to +<tt>synchronize_rcu()</tt>. +Similarly, an RCU read-side critical section must not +contain anything that waits, directly or indirectly, on completion of +an invocation of <tt>synchronize_rcu()</tt>. + +<p> +Although RCU's grace-period guarantee is useful in and of itself, with +<a href="https://lwn.net/Articles/573497/">quite a few use cases</a>, +it would be good to be able to use RCU to coordinate read-side +access to linked data structures. +For this, the grace-period guarantee is not sufficient, as can +be seen in function <tt>add_gp_buggy()</tt> below. +We will look at the reader's code later, but in the meantime, just think of +the reader as locklessly picking up the <tt>gp</tt> pointer, +and, if the value loaded is non-<tt>NULL</tt>, locklessly accessing the +<tt>->a</tt> and <tt>->b</tt> fields. + +<blockquote> +<pre> + 1 bool add_gp_buggy(int a, int b) + 2 { + 3 p = kmalloc(sizeof(*p), GFP_KERNEL); + 4 if (!p) + 5 return -ENOMEM; + 6 spin_lock(&gp_lock); + 7 if (rcu_access_pointer(gp)) { + 8 spin_unlock(&gp_lock); + 9 return false; +10 } +11 p->a = a; +12 p->b = a; +13 gp = p; /* ORDERING BUG */ +14 spin_unlock(&gp_lock); +15 return true; +16 } +</pre> +</blockquote> + +<p> +The problem is that both the compiler and weakly ordered CPUs are within +their rights to reorder this code as follows: + +<blockquote> +<pre> + 1 bool add_gp_buggy_optimized(int a, int b) + 2 { + 3 p = kmalloc(sizeof(*p), GFP_KERNEL); + 4 if (!p) + 5 return -ENOMEM; + 6 spin_lock(&gp_lock); + 7 if (rcu_access_pointer(gp)) { + 8 spin_unlock(&gp_lock); + 9 return false; +10 } +<b>11 gp = p; /* ORDERING BUG */ +12 p->a = a; +13 p->b = a;</b> +14 spin_unlock(&gp_lock); +15 return true; +16 } +</pre> +</blockquote> + +<p> +If an RCU reader fetches <tt>gp</tt> just after +<tt>add_gp_buggy_optimized</tt> executes line 11, +it will see garbage in the <tt>->a</tt> and <tt>->b</tt> +fields. +And this is but one of many ways in which compiler and hardware optimizations +could cause trouble. +Therefore, we clearly need some way to prevent the compiler and the CPU from +reordering in this manner, which brings us to the publish-subscribe +guarantee discussed in the next section. + +<h3><a name="Publish-Subscribe Guarantee">Publish/Subscribe Guarantee</a></h3> + +<p> +RCU's publish-subscribe guarantee allows data to be inserted +into a linked data structure without disrupting RCU readers. +The updater uses <tt>rcu_assign_pointer()</tt> to insert the +new data, and readers use <tt>rcu_dereference()</tt> to +access data, whether new or old. +The following shows an example of insertion: + +<blockquote> +<pre> + 1 bool add_gp(int a, int b) + 2 { + 3 p = kmalloc(sizeof(*p), GFP_KERNEL); + 4 if (!p) + 5 return -ENOMEM; + 6 spin_lock(&gp_lock); + 7 if (rcu_access_pointer(gp)) { + 8 spin_unlock(&gp_lock); + 9 return false; +10 } +11 p->a = a; +12 p->b = a; +13 rcu_assign_pointer(gp, p); +14 spin_unlock(&gp_lock); +15 return true; +16 } +</pre> +</blockquote> + +<p> +The <tt>rcu_assign_pointer()</tt> on line 13 is conceptually +equivalent to a simple assignment statement, but also guarantees +that its assignment will +happen after the two assignments in lines 11 and 12, +similar to the C11 <tt>memory_order_release</tt> store operation. +It also prevents any number of “interesting” compiler +optimizations, for example, the use of <tt>gp</tt> as a scratch +location immediately preceding the assignment. + +<p><a name="Quick Quiz 3"><b>Quick Quiz 3</b>:</a> +But <tt>rcu_assign_pointer()</tt> does nothing to prevent the +two assignments to <tt>p->a</tt> and <tt>p->b</tt> +from being reordered. +Can't that also cause problems? +<br><a href="#qq3answer">Answer</a> + +<p> +It is tempting to assume that the reader need not do anything special +to control its accesses to the RCU-protected data, +as shown in <tt>do_something_gp_buggy()</tt> below: + +<blockquote> +<pre> + 1 bool do_something_gp_buggy(void) + 2 { + 3 rcu_read_lock(); + 4 p = gp; /* OPTIMIZATIONS GALORE!!! */ + 5 if (p) { + 6 do_something(p->a, p->b); + 7 rcu_read_unlock(); + 8 return true; + 9 } +10 rcu_read_unlock(); +11 return false; +12 } +</pre> +</blockquote> + +<p> +However, this temptation must be resisted because there are a +surprisingly large number of ways that the compiler +(to say nothing of +<a href="https://h71000.www7.hp.com/wizard/wiz_2637.html">DEC Alpha CPUs</a>) +can trip this code up. +For but one example, if the compiler were short of registers, it +might choose to refetch from <tt>gp</tt> rather than keeping +a separate copy in <tt>p</tt> as follows: + +<blockquote> +<pre> + 1 bool do_something_gp_buggy_optimized(void) + 2 { + 3 rcu_read_lock(); + 4 if (gp) { /* OPTIMIZATIONS GALORE!!! */ +<b> 5 do_something(gp->a, gp->b);</b> + 6 rcu_read_unlock(); + 7 return true; + 8 } + 9 rcu_read_unlock(); +10 return false; +11 } +</pre> +</blockquote> + +<p> +If this function ran concurrently with a series of updates that +replaced the current structure with a new one, +the fetches of <tt>gp->a</tt> +and <tt>gp->b</tt> might well come from two different structures, +which could cause serious confusion. +To prevent this (and much else besides), <tt>do_something_gp()</tt> uses +<tt>rcu_dereference()</tt> to fetch from <tt>gp</tt>: + +<blockquote> +<pre> + 1 bool do_something_gp(void) + 2 { + 3 rcu_read_lock(); + 4 p = rcu_dereference(gp); + 5 if (p) { + 6 do_something(p->a, p->b); + 7 rcu_read_unlock(); + 8 return true; + 9 } +10 rcu_read_unlock(); +11 return false; +12 } +</pre> +</blockquote> + +<p> +The <tt>rcu_dereference()</tt> uses volatile casts and (for DEC Alpha) +memory barriers in the Linux kernel. +Should a +<a href="http://www.rdrop.com/users/paulmck/RCU/consume.2015.07.13a.pdf">high-quality implementation of C11 <tt>memory_order_consume</tt> [PDF]</a> +ever appear, then <tt>rcu_dereference()</tt> could be implemented +as a <tt>memory_order_consume</tt> load. +Regardless of the exact implementation, a pointer fetched by +<tt>rcu_dereference()</tt> may not be used outside of the +outermost RCU read-side critical section containing that +<tt>rcu_dereference()</tt>, unless protection of +the corresponding data element has been passed from RCU to some +other synchronization mechanism, most commonly locking or +<a href="https://www.kernel.org/doc/Documentation/RCU/rcuref.txt">reference counting</a>. + +<p> +In short, updaters use <tt>rcu_assign_pointer()</tt> and readers +use <tt>rcu_dereference()</tt>, and these two RCU API elements +work together to ensure that readers have a consistent view of +newly added data elements. + +<p> +Of course, it is also necessary to remove elements from RCU-protected +data structures, for example, using the following process: + +<ol> +<li> Remove the data element from the enclosing structure. +<li> Wait for all pre-existing RCU read-side critical sections + to complete (because only pre-existing readers can possibly have + a reference to the newly removed data element). +<li> At this point, only the updater has a reference to the + newly removed data element, so it can safely reclaim + the data element, for example, by passing it to <tt>kfree()</tt>. +</ol> + +This process is implemented by <tt>remove_gp_synchronous()</tt>: + +<blockquote> +<pre> + 1 bool remove_gp_synchronous(void) + 2 { + 3 struct foo *p; + 4 + 5 spin_lock(&gp_lock); + 6 p = rcu_access_pointer(gp); + 7 if (!p) { + 8 spin_unlock(&gp_lock); + 9 return false; +10 } +11 rcu_assign_pointer(gp, NULL); +12 spin_unlock(&gp_lock); +13 synchronize_rcu(); +14 kfree(p); +15 return true; +16 } +</pre> +</blockquote> + +<p> +This function is straightforward, with line 13 waiting for a grace +period before line 14 frees the old data element. +This waiting ensures that readers will reach line 7 of +<tt>do_something_gp()</tt> before the data element referenced by +<tt>p</tt> is freed. +The <tt>rcu_access_pointer()</tt> on line 6 is similar to +<tt>rcu_dereference()</tt>, except that: + +<ol> +<li> The value returned by <tt>rcu_access_pointer()</tt> + cannot be dereferenced. + If you want to access the value pointed to as well as + the pointer itself, use <tt>rcu_dereference()</tt> + instead of <tt>rcu_access_pointer()</tt>. +<li> The call to <tt>rcu_access_pointer()</tt> need not be + protected. + In contrast, <tt>rcu_dereference()</tt> must either be + within an RCU read-side critical section or in a code + segment where the pointer cannot change, for example, in + code protected by the corresponding update-side lock. +</ol> + +<p><a name="Quick Quiz 4"><b>Quick Quiz 4</b>:</a> +Without the <tt>rcu_dereference()</tt> or the +<tt>rcu_access_pointer()</tt>, what destructive optimizations +might the compiler make use of? +<br><a href="#qq4answer">Answer</a> + +<p> +In short, RCU's publish-subscribe guarantee is provided by the combination +of <tt>rcu_assign_pointer()</tt> and <tt>rcu_dereference()</tt>. +This guarantee allows data elements to be safely added to RCU-protected +linked data structures without disrupting RCU readers. +This guarantee can be used in combination with the grace-period +guarantee to also allow data elements to be removed from RCU-protected +linked data structures, again without disrupting RCU readers. + +<p> +This guarantee was only partially premeditated. +DYNIX/ptx used an explicit memory barrier for publication, but had nothing +resembling <tt>rcu_dereference()</tt> for subscription, nor did it +have anything resembling the <tt>smp_read_barrier_depends()</tt> +that was later subsumed into <tt>rcu_dereference()</tt>. +The need for these operations made itself known quite suddenly at a +late-1990s meeting with the DEC Alpha architects, back in the days when +DEC was still a free-standing company. +It took the Alpha architects a good hour to convince me that any sort +of barrier would ever be needed, and it then took me a good <i>two</i> hours +to convince them that their documentation did not make this point clear. +More recent work with the C and C++ standards committees have provided +much education on tricks and traps from the compiler. +In short, compilers were much less tricky in the early 1990s, but in +2015, don't even think about omitting <tt>rcu_dereference()</tt>! + +<h3><a name="Memory-Barrier Guarantees">Memory-Barrier Guarantees</a></h3> + +<p> +The previous section's simple linked-data-structure scenario clearly +demonstrates the need for RCU's stringent memory-ordering guarantees on +systems with more than one CPU: + +<ol> +<li> Each CPU that has an RCU read-side critical section that + begins before <tt>synchronize_rcu()</tt> starts is + guaranteed to execute a full memory barrier between the time + that the RCU read-side critical section ends and the time that + <tt>synchronize_rcu()</tt> returns. + Without this guarantee, a pre-existing RCU read-side critical section + might hold a reference to the newly removed <tt>struct foo</tt> + after the <tt>kfree()</tt> on line 14 of + <tt>remove_gp_synchronous()</tt>. +<li> Each CPU that has an RCU read-side critical section that ends + after <tt>synchronize_rcu()</tt> returns is guaranteed + to execute a full memory barrier between the time that + <tt>synchronize_rcu()</tt> begins and the time that the RCU + read-side critical section begins. + Without this guarantee, a later RCU read-side critical section + running after the <tt>kfree()</tt> on line 14 of + <tt>remove_gp_synchronous()</tt> might + later run <tt>do_something_gp()</tt> and find the + newly deleted <tt>struct foo</tt>. +<li> If the task invoking <tt>synchronize_rcu()</tt> remains + on a given CPU, then that CPU is guaranteed to execute a full + memory barrier sometime during the execution of + <tt>synchronize_rcu()</tt>. + This guarantee ensures that the <tt>kfree()</tt> on + line 14 of <tt>remove_gp_synchronous()</tt> really does + execute after the removal on line 11. +<li> If the task invoking <tt>synchronize_rcu()</tt> migrates + among a group of CPUs during that invocation, then each of the + CPUs in that group is guaranteed to execute a full memory barrier + sometime during the execution of <tt>synchronize_rcu()</tt>. + This guarantee also ensures that the <tt>kfree()</tt> on + line 14 of <tt>remove_gp_synchronous()</tt> really does + execute after the removal on + line 11, but also in the case where the thread executing the + <tt>synchronize_rcu()</tt> migrates in the meantime. +</ol> + +<p><a name="Quick Quiz 5"><b>Quick Quiz 5</b>:</a> +Given that multiple CPUs can start RCU read-side critical sections +at any time without any ordering whatsoever, how can RCU possibly tell whether +or not a given RCU read-side critical section starts before a +given instance of <tt>synchronize_rcu()</tt>? +<br><a href="#qq5answer">Answer</a> + +<p><a name="Quick Quiz 6"><b>Quick Quiz 6</b>:</a> +The first and second guarantees require unbelievably strict ordering! +Are all these memory barriers <i> really</i> required? +<br><a href="#qq6answer">Answer</a> + +<p> +Note that these memory-barrier requirements do not replace the fundamental +RCU requirement that a grace period wait for all pre-existing readers. +On the contrary, the memory barriers called out in this section must operate in +such a way as to <i>enforce</i> this fundamental requirement. +Of course, different implementations enforce this requirement in different +ways, but enforce it they must. + +<h3><a name="RCU Primitives Guaranteed to Execute Unconditionally">RCU Primitives Guaranteed to Execute Unconditionally</a></h3> + +<p> +The common-case RCU primitives are unconditional. +They are invoked, they do their job, and they return, with no possibility +of error, and no need to retry. +This is a key RCU design philosophy. + +<p> +However, this philosophy is pragmatic rather than pigheaded. +If someone comes up with a good justification for a particular conditional +RCU primitive, it might well be implemented and added. +After all, this guarantee was reverse-engineered, not premeditated. +The unconditional nature of the RCU primitives was initially an +accident of implementation, and later experience with synchronization +primitives with conditional primitives caused me to elevate this +accident to a guarantee. +Therefore, the justification for adding a conditional primitive to +RCU would need to be based on detailed and compelling use cases. + +<h3><a name="Guaranteed Read-to-Write Upgrade">Guaranteed Read-to-Write Upgrade</a></h3> + +<p> +As far as RCU is concerned, it is always possible to carry out an +update within an RCU read-side critical section. +For example, that RCU read-side critical section might search for +a given data element, and then might acquire the update-side +spinlock in order to update that element, all while remaining +in that RCU read-side critical section. +Of course, it is necessary to exit the RCU read-side critical section +before invoking <tt>synchronize_rcu()</tt>, however, this +inconvenience can be avoided through use of the +<tt>call_rcu()</tt> and <tt>kfree_rcu()</tt> API members +described later in this document. + +<p><a name="Quick Quiz 7"><b>Quick Quiz 7</b>:</a> +But how does the upgrade-to-write operation exclude other readers? +<br><a href="#qq7answer">Answer</a> + +<p> +This guarantee allows lookup code to be shared between read-side +and update-side code, and was premeditated, appearing in the earliest +DYNIX/ptx RCU documentation. + +<h2><a name="Fundamental Non-Requirements">Fundamental Non-Requirements</a></h2> + +<p> +RCU provides extremely lightweight readers, and its read-side guarantees, +though quite useful, are correspondingly lightweight. +It is therefore all too easy to assume that RCU is guaranteeing more +than it really is. +Of course, the list of things that RCU does not guarantee is infinitely +long, however, the following sections list a few non-guarantees that +have caused confusion. +Except where otherwise noted, these non-guarantees were premeditated. + +<ol> +<li> <a href="#Readers Impose Minimal Ordering"> + Readers Impose Minimal Ordering</a> +<li> <a href="#Readers Do Not Exclude Updaters"> + Readers Do Not Exclude Updaters</a> +<li> <a href="#Updaters Only Wait For Old Readers"> + Updaters Only Wait For Old Readers</a> +<li> <a href="#Grace Periods Don't Partition Read-Side Critical Sections"> + Grace Periods Don't Partition Read-Side Critical Sections</a> +<li> <a href="#Read-Side Critical Sections Don't Partition Grace Periods"> + Read-Side Critical Sections Don't Partition Grace Periods</a> +<li> <a href="#Disabling Preemption Does Not Block Grace Periods"> + Disabling Preemption Does Not Block Grace Periods</a> +</ol> + +<h3><a name="Readers Impose Minimal Ordering">Readers Impose Minimal Ordering</a></h3> + +<p> +Reader-side markers such as <tt>rcu_read_lock()</tt> and +<tt>rcu_read_unlock()</tt> provide absolutely no ordering guarantees +except through their interaction with the grace-period APIs such as +<tt>synchronize_rcu()</tt>. +To see this, consider the following pair of threads: + +<blockquote> +<pre> + 1 void thread0(void) + 2 { + 3 rcu_read_lock(); + 4 WRITE_ONCE(x, 1); + 5 rcu_read_unlock(); + 6 rcu_read_lock(); + 7 WRITE_ONCE(y, 1); + 8 rcu_read_unlock(); + 9 } +10 +11 void thread1(void) +12 { +13 rcu_read_lock(); +14 r1 = READ_ONCE(y); +15 rcu_read_unlock(); +16 rcu_read_lock(); +17 r2 = READ_ONCE(x); +18 rcu_read_unlock(); +19 } +</pre> +</blockquote> + +<p> +After <tt>thread0()</tt> and <tt>thread1()</tt> execute +concurrently, it is quite possible to have + +<blockquote> +<pre> +(r1 == 1 && r2 == 0) +</pre> +</blockquote> + +(that is, <tt>y</tt> appears to have been assigned before <tt>x</tt>), +which would not be possible if <tt>rcu_read_lock()</tt> and +<tt>rcu_read_unlock()</tt> had much in the way of ordering +properties. +But they do not, so the CPU is within its rights +to do significant reordering. +This is by design: Any significant ordering constraints would slow down +these fast-path APIs. + +<p><a name="Quick Quiz 8"><b>Quick Quiz 8</b>:</a> +Can't the compiler also reorder this code? +<br><a href="#qq8answer">Answer</a> + +<h3><a name="Readers Do Not Exclude Updaters">Readers Do Not Exclude Updaters</a></h3> + +<p> +Neither <tt>rcu_read_lock()</tt> nor <tt>rcu_read_unlock()</tt> +exclude updates. +All they do is to prevent grace periods from ending. +The following example illustrates this: + +<blockquote> +<pre> + 1 void thread0(void) + 2 { + 3 rcu_read_lock(); + 4 r1 = READ_ONCE(y); + 5 if (r1) { + 6 do_something_with_nonzero_x(); + 7 r2 = READ_ONCE(x); + 8 WARN_ON(!r2); /* BUG!!! */ + 9 } +10 rcu_read_unlock(); +11 } +12 +13 void thread1(void) +14 { +15 spin_lock(&my_lock); +16 WRITE_ONCE(x, 1); +17 WRITE_ONCE(y, 1); +18 spin_unlock(&my_lock); +19 } +</pre> +</blockquote> + +<p> +If the <tt>thread0()</tt> function's <tt>rcu_read_lock()</tt> +excluded the <tt>thread1()</tt> function's update, +the <tt>WARN_ON()</tt> could never fire. +But the fact is that <tt>rcu_read_lock()</tt> does not exclude +much of anything aside from subsequent grace periods, of which +<tt>thread1()</tt> has none, so the +<tt>WARN_ON()</tt> can and does fire. + +<h3><a name="Updaters Only Wait For Old Readers">Updaters Only Wait For Old Readers</a></h3> + +<p> +It might be tempting to assume that after <tt>synchronize_rcu()</tt> +completes, there are no readers executing. +This temptation must be avoided because +new readers can start immediately after <tt>synchronize_rcu()</tt> +starts, and <tt>synchronize_rcu()</tt> is under no +obligation to wait for these new readers. + +<p><a name="Quick Quiz 9"><b>Quick Quiz 9</b>:</a> +Suppose that synchronize_rcu() did wait until all readers had completed. +Would the updater be able to rely on this? +<br><a href="#qq9answer">Answer</a> + +<h3><a name="Grace Periods Don't Partition Read-Side Critical Sections"> +Grace Periods Don't Partition Read-Side Critical Sections</a></h3> + +<p> +It is tempting to assume that if any part of one RCU read-side critical +section precedes a given grace period, and if any part of another RCU +read-side critical section follows that same grace period, then all of +the first RCU read-side critical section must precede all of the second. +However, this just isn't the case: A single grace period does not +partition the set of RCU read-side critical sections. +An example of this situation can be illustrated as follows, where +<tt>x</tt>, <tt>y</tt>, and <tt>z</tt> are initially all zero: + +<blockquote> +<pre> + 1 void thread0(void) + 2 { + 3 rcu_read_lock(); + 4 WRITE_ONCE(a, 1); + 5 WRITE_ONCE(b, 1); + 6 rcu_read_unlock(); + 7 } + 8 + 9 void thread1(void) +10 { +11 r1 = READ_ONCE(a); +12 synchronize_rcu(); +13 WRITE_ONCE(c, 1); +14 } +15 +16 void thread2(void) +17 { +18 rcu_read_lock(); +19 r2 = READ_ONCE(b); +20 r3 = READ_ONCE(c); +21 rcu_read_unlock(); +22 } +</pre> +</blockquote> + +<p> +It turns out that the outcome: + +<blockquote> +<pre> +(r1 == 1 && r2 == 0 && r3 == 1) +</pre> +</blockquote> + +is entirely possible. +The following figure show how this can happen, with each circled +<tt>QS</tt> indicating the point at which RCU recorded a +<i>quiescent state</i> for each thread, that is, a state in which +RCU knows that the thread cannot be in the midst of an RCU read-side +critical section that started before the current grace period: + +<p><img src="GPpartitionReaders1.svg" alt="GPpartitionReaders1.svg" width="60%"></p> + +<p> +If it is necessary to partition RCU read-side critical sections in this +manner, it is necessary to use two grace periods, where the first +grace period is known to end before the second grace period starts: + +<blockquote> +<pre> + 1 void thread0(void) + 2 { + 3 rcu_read_lock(); + 4 WRITE_ONCE(a, 1); + 5 WRITE_ONCE(b, 1); + 6 rcu_read_unlock(); + 7 } + 8 + 9 void thread1(void) +10 { +11 r1 = READ_ONCE(a); +12 synchronize_rcu(); +13 WRITE_ONCE(c, 1); +14 } +15 +16 void thread2(void) +17 { +18 r2 = READ_ONCE(c); +19 synchronize_rcu(); +20 WRITE_ONCE(d, 1); +21 } +22 +23 void thread3(void) +24 { +25 rcu_read_lock(); +26 r3 = READ_ONCE(b); +27 r4 = READ_ONCE(d); +28 rcu_read_unlock(); +29 } +</pre> +</blockquote> + +<p> +Here, if <tt>(r1 == 1)</tt>, then +<tt>thread0()</tt>'s write to <tt>b</tt> must happen +before the end of <tt>thread1()</tt>'s grace period. +If in addition <tt>(r4 == 1)</tt>, then +<tt>thread3()</tt>'s read from <tt>b</tt> must happen +after the beginning of <tt>thread2()</tt>'s grace period. +If it is also the case that <tt>(r2 == 1)</tt>, then the +end of <tt>thread1()</tt>'s grace period must precede the +beginning of <tt>thread2()</tt>'s grace period. +This mean that the two RCU read-side critical sections cannot overlap, +guaranteeing that <tt>(r3 == 1)</tt>. +As a result, the outcome: + +<blockquote> +<pre> +(r1 == 1 && r2 == 1 && r3 == 0 && r4 == 1) +</pre> +</blockquote> + +cannot happen. + +<p> +This non-requirement was also non-premeditated, but became apparent +when studying RCU's interaction with memory ordering. + +<h3><a name="Read-Side Critical Sections Don't Partition Grace Periods"> +Read-Side Critical Sections Don't Partition Grace Periods</a></h3> + +<p> +It is also tempting to assume that if an RCU read-side critical section +happens between a pair of grace periods, then those grace periods cannot +overlap. +However, this temptation leads nowhere good, as can be illustrated by +the following, with all variables initially zero: + +<blockquote> +<pre> + 1 void thread0(void) + 2 { + 3 rcu_read_lock(); + 4 WRITE_ONCE(a, 1); + 5 WRITE_ONCE(b, 1); + 6 rcu_read_unlock(); + 7 } + 8 + 9 void thread1(void) +10 { +11 r1 = READ_ONCE(a); +12 synchronize_rcu(); +13 WRITE_ONCE(c, 1); +14 } +15 +16 void thread2(void) +17 { +18 rcu_read_lock(); +19 WRITE_ONCE(d, 1); +20 r2 = READ_ONCE(c); +21 rcu_read_unlock(); +22 } +23 +24 void thread3(void) +25 { +26 r3 = READ_ONCE(d); +27 synchronize_rcu(); +28 WRITE_ONCE(e, 1); +29 } +30 +31 void thread4(void) +32 { +33 rcu_read_lock(); +34 r4 = READ_ONCE(b); +35 r5 = READ_ONCE(e); +36 rcu_read_unlock(); +37 } +</pre> +</blockquote> + +<p> +In this case, the outcome: + +<blockquote> +<pre> +(r1 == 1 && r2 == 1 && r3 == 1 && r4 == 0 && r5 == 1) +</pre> +</blockquote> + +is entirely possible, as illustrated below: + +<p><img src="ReadersPartitionGP1.svg" alt="ReadersPartitionGP1.svg" width="100%"></p> + +<p> +Again, an RCU read-side critical section can overlap almost all of a +given grace period, just so long as it does not overlap the entire +grace period. +As a result, an RCU read-side critical section cannot partition a pair +of RCU grace periods. + +<p><a name="Quick Quiz 10"><b>Quick Quiz 10</b>:</a> +How long a sequence of grace periods, each separated by an RCU read-side +critical section, would be required to partition the RCU read-side +critical sections at the beginning and end of the chain? +<br><a href="#qq10answer">Answer</a> + +<h3><a name="Disabling Preemption Does Not Block Grace Periods"> +Disabling Preemption Does Not Block Grace Periods</a></h3> + +<p> +There was a time when disabling preemption on any given CPU would block +subsequent grace periods. +However, this was an accident of implementation and is not a requirement. +And in the current Linux-kernel implementation, disabling preemption +on a given CPU in fact does not block grace periods, as Oleg Nesterov +<a href="https://lkml.kernel.org/g/20150614193825.GA19582@redhat.com">demonstrated</a>. + +<p> +If you need a preempt-disable region to block grace periods, you need to add +<tt>rcu_read_lock()</tt> and <tt>rcu_read_unlock()</tt>, for example +as follows: + +<blockquote> +<pre> + 1 preempt_disable(); + 2 rcu_read_lock(); + 3 do_something(); + 4 rcu_read_unlock(); + 5 preempt_enable(); + 6 + 7 /* Spinlocks implicitly disable preemption. */ + 8 spin_lock(&mylock); + 9 rcu_read_lock(); +10 do_something(); +11 rcu_read_unlock(); +12 spin_unlock(&mylock); +</pre> +</blockquote> + +<p> +In theory, you could enter the RCU read-side critical section first, +but it is more efficient to keep the entire RCU read-side critical +section contained in the preempt-disable region as shown above. +Of course, RCU read-side critical sections that extend outside of +preempt-disable regions will work correctly, but such critical sections +can be preempted, which forces <tt>rcu_read_unlock()</tt> to do +more work. +And no, this is <i>not</i> an invitation to enclose all of your RCU +read-side critical sections within preempt-disable regions, because +doing so would degrade real-time response. + +<p> +This non-requirement appeared with preemptible RCU. +If you need a grace period that waits on non-preemptible code regions, use +<a href="#Sched Flavor">RCU-sched</a>. + +<h2><a name="Parallelism Facts of Life">Parallelism Facts of Life</a></h2> + +<p> +These parallelism facts of life are by no means specific to RCU, but +the RCU implementation must abide by them. +They therefore bear repeating: + +<ol> +<li> Any CPU or task may be delayed at any time, + and any attempts to avoid these delays by disabling + preemption, interrupts, or whatever are completely futile. + This is most obvious in preemptible user-level + environments and in virtualized environments (where + a given guest OS's VCPUs can be preempted at any time by + the underlying hypervisor), but can also happen in bare-metal + environments due to ECC errors, NMIs, and other hardware + events. + Although a delay of more than about 20 seconds can result + in splats, the RCU implementation is obligated to use + algorithms that can tolerate extremely long delays, but where + “extremely long” is not long enough to allow + wrap-around when incrementing a 64-bit counter. +<li> Both the compiler and the CPU can reorder memory accesses. + Where it matters, RCU must use compiler directives and + memory-barrier instructions to preserve ordering. +<li> Conflicting writes to memory locations in any given cache line + will result in expensive cache misses. + Greater numbers of concurrent writes and more-frequent + concurrent writes will result in more dramatic slowdowns. + RCU is therefore obligated to use algorithms that have + sufficient locality to avoid significant performance and + scalability problems. +<li> As a rough rule of thumb, only one CPU's worth of processing + may be carried out under the protection of any given exclusive + lock. + RCU must therefore use scalable locking designs. +<li> Counters are finite, especially on 32-bit systems. + RCU's use of counters must therefore tolerate counter wrap, + or be designed such that counter wrap would take way more + time than a single system is likely to run. + An uptime of ten years is quite possible, a runtime + of a century much less so. + As an example of the latter, RCU's dyntick-idle nesting counter + allows 54 bits for interrupt nesting level (this counter + is 64 bits even on a 32-bit system). + Overflowing this counter requires 2<sup>54</sup> + half-interrupts on a given CPU without that CPU ever going idle. + If a half-interrupt happened every microsecond, it would take + 570 years of runtime to overflow this counter, which is currently + believed to be an acceptably long time. +<li> Linux systems can have thousands of CPUs running a single + Linux kernel in a single shared-memory environment. + RCU must therefore pay close attention to high-end scalability. +</ol> + +<p> +This last parallelism fact of life means that RCU must pay special +attention to the preceding facts of life. +The idea that Linux might scale to systems with thousands of CPUs would +have been met with some skepticism in the 1990s, but these requirements +would have otherwise have been unsurprising, even in the early 1990s. + +<h2><a name="Quality-of-Implementation Requirements">Quality-of-Implementation Requirements</a></h2> + +<p> +These sections list quality-of-implementation requirements. +Although an RCU implementation that ignores these requirements could +still be used, it would likely be subject to limitations that would +make it inappropriate for industrial-strength production use. +Classes of quality-of-implementation requirements are as follows: + +<ol> +<li> <a href="#Specialization">Specialization</a> +<li> <a href="#Performance and Scalability">Performance and Scalability</a> +<li> <a href="#Composability">Composability</a> +<li> <a href="#Corner Cases">Corner Cases</a> +</ol> + +<p> +These classes is covered in the following sections. + +<h3><a name="Specialization">Specialization</a></h3> + +<p> +RCU is and always has been intended primarily for read-mostly situations, as +illustrated by the following figure. +This means that RCU's read-side primitives are optimized, often at the +expense of its update-side primitives. + +<p><img src="RCUApplicability.svg" alt="RCUApplicability.svg" width="70%"></p> + +<p> +This focus on read-mostly situations means that RCU must interoperate +with other synchronization primitives. +For example, the <tt>add_gp()</tt> and <tt>remove_gp_synchronous()</tt> +examples discussed earlier use RCU to protect readers and locking to +coordinate updaters. +However, the need extends much farther, requiring that a variety of +synchronization primitives be legal within RCU read-side critical sections, +including spinlocks, sequence locks, atomic operations, reference +counters, and memory barriers. + +<p><a name="Quick Quiz 11"><b>Quick Quiz 11</b>:</a> +What about sleeping locks? +<br><a href="#qq11answer">Answer</a> + +<p> +It often comes as a surprise that many algorithms do not require a +consistent view of data, but many can function in that mode, +with network routing being the poster child. +Internet routing algorithms take significant time to propagate +updates, so that by the time an update arrives at a given system, +that system has been sending network traffic the wrong way for +a considerable length of time. +Having a few threads continue to send traffic the wrong way for a +few more milliseconds is clearly not a problem: In the worst case, +TCP retransmissions will eventually get the data where it needs to go. +In general, when tracking the state of the universe outside of the +computer, some level of inconsistency must be tolerated due to +speed-of-light delays if nothing else. + +<p> +Furthermore, uncertainty about external state is inherent in many cases. +For example, a pair of veternarians might use heartbeat to determine +whether or not a given cat was alive. +But how long should they wait after the last heartbeat to decide that +the cat is in fact dead? +Waiting less than 400 milliseconds makes no sense because this would +mean that a relaxed cat would be considered to cycle between death +and life more than 100 times per minute. +Moreover, just as with human beings, a cat's heart might stop for +some period of time, so the exact wait period is a judgment call. +One of our pair of veternarians might wait 30 seconds before pronouncing +the cat dead, while the other might insist on waiting a full minute. +The two veternarians would then disagree on the state of the cat during +the final 30 seconds of the minute following the last heartbeat, as +fancifully illustrated below: + +<p><img src="2013-08-is-it-dead.png" alt="2013-08-is-it-dead.png" width="431"></p> + +<p> +Interestingly enough, this same situation applies to hardware. +When push comes to shove, how do we tell whether or not some +external server has failed? +We send messages to it periodically, and declare it failed if we +don't receive a response within a given period of time. +Policy decisions can usually tolerate short +periods of inconsistency. +The policy was decided some time ago, and is only now being put into +effect, so a few milliseconds of delay is normally inconsequential. + +<p> +However, there are algorithms that absolutely must see consistent data. +For example, the translation between a user-level SystemV semaphore +ID to the corresponding in-kernel data structure is protected by RCU, +but it is absolutely forbidden to update a semaphore that has just been +removed. +In the Linux kernel, this need for consistency is accommodated by acquiring +spinlocks located in the in-kernel data structure from within +the RCU read-side critical section, and this is indicated by the +green box in the figure above. +Many other techniques may be used, and are in fact used within the +Linux kernel. + +<p> +In short, RCU is not required to maintain consistency, and other +mechanisms may be used in concert with RCU when consistency is required. +RCU's specialization allows it to do its job extremely well, and its +ability to interoperate with other synchronization mechanisms allows +the right mix of synchronization tools to be used for a given job. + +<h3><a name="Performance and Scalability">Performance and Scalability</a></h3> + +<p> +Energy efficiency is a critical component of performance today, +and Linux-kernel RCU implementations must therefore avoid unnecessarily +awakening idle CPUs. +I cannot claim that this requirement was premeditated. +In fact, I learned of it during a telephone conversation in which I +was given “frank and open” feedback on the importance +of energy efficiency in battery-powered systems and on specific +energy-efficiency shortcomings of the Linux-kernel RCU implementation. +In my experience, the battery-powered embedded community will consider +any unnecessary wakeups to be extremely unfriendly acts. +So much so that mere Linux-kernel-mailing-list posts are +insufficient to vent their ire. + +<p> +Memory consumption is not particularly important for in most +situations, and has become decreasingly +so as memory sizes have expanded and memory +costs have plummeted. +However, as I learned from Matt Mackall's +<a href="http://elinux.org/Linux_Tiny-FAQ">bloatwatch</a> +efforts, memory footprint is critically important on single-CPU systems with +non-preemptible (<tt>CONFIG_PREEMPT=n</tt>) kernels, and thus +<a href="https://lkml.kernel.org/g/20090113221724.GA15307@linux.vnet.ibm.com">tiny RCU</a> +was born. +Josh Triplett has since taken over the small-memory banner with his +<a href="https://tiny.wiki.kernel.org/">Linux kernel tinification</a> +project, which resulted in +<a href="#Sleepable RCU">SRCU</a> +becoming optional for those kernels not needing it. + +<p> +The remaining performance requirements are, for the most part, +unsurprising. +For example, in keeping with RCU's read-side specialization, +<tt>rcu_dereference()</tt> should have negligible overhead (for +example, suppression of a few minor compiler optimizations). +Similarly, in non-preemptible environments, <tt>rcu_read_lock()</tt> and +<tt>rcu_read_unlock()</tt> should have exactly zero overhead. + +<p> +In preemptible environments, in the case where the RCU read-side +critical section was not preempted (as will be the case for the +highest-priority real-time process), <tt>rcu_read_lock()</tt> and +<tt>rcu_read_unlock()</tt> should have minimal overhead. +In particular, they should not contain atomic read-modify-write +operations, memory-barrier instructions, preemption disabling, +interrupt disabling, or backwards branches. +However, in the case where the RCU read-side critical section was preempted, +<tt>rcu_read_unlock()</tt> may acquire spinlocks and disable interrupts. +This is why it is better to nest an RCU read-side critical section +within a preempt-disable region than vice versa, at least in cases +where that critical section is short enough to avoid unduly degrading +real-time latencies. + +<p> +The <tt>synchronize_rcu()</tt> grace-period-wait primitive is +optimized for throughput. +It may therefore incur several milliseconds of latency in addition to +the duration of the longest RCU read-side critical section. +On the other hand, multiple concurrent invocations of +<tt>synchronize_rcu()</tt> are required to use batching optimizations +so that they can be satisfied by a single underlying grace-period-wait +operation. +For example, in the Linux kernel, it is not unusual for a single +grace-period-wait operation to serve more than +<a href="https://www.usenix.org/conference/2004-usenix-annual-technical-conference/making-rcu-safe-deep-sub-millisecond-response">1,000 separate invocations</a> +of <tt>synchronize_rcu()</tt>, thus amortizing the per-invocation +overhead down to nearly zero. +However, the grace-period optimization is also required to avoid +measurable degradation of real-time scheduling and interrupt latencies. + +<p> +In some cases, the multi-millisecond <tt>synchronize_rcu()</tt> +latencies are unacceptable. +In these cases, <tt>synchronize_rcu_expedited()</tt> may be used +instead, reducing the grace-period latency down to a few tens of +microseconds on small systems, at least in cases where the RCU read-side +critical sections are short. +There are currently no special latency requirements for +<tt>synchronize_rcu_expedited()</tt> on large systems, but, +consistent with the empirical nature of the RCU specification, +that is subject to change. +However, there most definitely are scalability requirements: +A storm of <tt>synchronize_rcu_expedited()</tt> invocations on 4096 +CPUs should at least make reasonable forward progress. +In return for its shorter latencies, <tt>synchronize_rcu_expedited()</tt> +is permitted to impose modest degradation of real-time latency +on non-idle online CPUs. +That said, it will likely be necessary to take further steps to reduce this +degradation, hopefully to roughly that of a scheduling-clock interrupt. + +<p> +There are a number of situations where even +<tt>synchronize_rcu_expedited()</tt>'s reduced grace-period +latency is unacceptable. +In these situations, the asynchronous <tt>call_rcu()</tt> can be +used in place of <tt>synchronize_rcu()</tt> as follows: + +<blockquote> +<pre> + 1 struct foo { + 2 int a; + 3 int b; + 4 struct rcu_head rh; + 5 }; + 6 + 7 static void remove_gp_cb(struct rcu_head *rhp) + 8 { + 9 struct foo *p = container_of(rhp, struct foo, rh); +10 +11 kfree(p); +12 } +13 +14 bool remove_gp_asynchronous(void) +15 { +16 struct foo *p; +17 +18 spin_lock(&gp_lock); +19 p = rcu_dereference(gp); +20 if (!p) { +21 spin_unlock(&gp_lock); +22 return false; +23 } +24 rcu_assign_pointer(gp, NULL); +25 call_rcu(&p->rh, remove_gp_cb); +26 spin_unlock(&gp_lock); +27 return true; +28 } +</pre> +</blockquote> + +<p> +A definition of <tt>struct foo</tt> is finally needed, and appears +on lines 1-5. +The function <tt>remove_gp_cb()</tt> is passed to <tt>call_rcu()</tt> +on line 25, and will be invoked after the end of a subsequent +grace period. +This gets the same effect as <tt>remove_gp_synchronous()</tt>, +but without forcing the updater to wait for a grace period to elapse. +The <tt>call_rcu()</tt> function may be used in a number of +situations where neither <tt>synchronize_rcu()</tt> nor +<tt>synchronize_rcu_expedited()</tt> would be legal, +including within preempt-disable code, <tt>local_bh_disable()</tt> code, +interrupt-disable code, and interrupt handlers. +However, even <tt>call_rcu()</tt> is illegal within NMI handlers. +The callback function (<tt>remove_gp_cb()</tt> in this case) will be +executed within softirq (software interrupt) environment within the +Linux kernel, +either within a real softirq handler or under the protection +of <tt>local_bh_disable()</tt>. +In both the Linux kernel and in userspace, it is bad practice to +write an RCU callback function that takes too long. +Long-running operations should be relegated to separate threads or +(in the Linux kernel) workqueues. + +<p><a name="Quick Quiz 12"><b>Quick Quiz 12</b>:</a> +Why does line 19 use <tt>rcu_access_pointer()</tt>? +After all, <tt>call_rcu()</tt> on line 25 stores into the +structure, which would interact badly with concurrent insertions. +Doesn't this mean that <tt>rcu_dereference()</tt> is required? +<br><a href="#qq12answer">Answer</a> + +<p> +However, all that <tt>remove_gp_cb()</tt> is doing is +invoking <tt>kfree()</tt> on the data element. +This is a common idiom, and is supported by <tt>kfree_rcu()</tt>, +which allows “fire and forget” operation as shown below: + +<blockquote> +<pre> + 1 struct foo { + 2 int a; + 3 int b; + 4 struct rcu_head rh; + 5 }; + 6 + 7 bool remove_gp_faf(void) + 8 { + 9 struct foo *p; +10 +11 spin_lock(&gp_lock); +12 p = rcu_dereference(gp); +13 if (!p) { +14 spin_unlock(&gp_lock); +15 return false; +16 } +17 rcu_assign_pointer(gp, NULL); +18 kfree_rcu(p, rh); +19 spin_unlock(&gp_lock); +20 return true; +21 } +</pre> +</blockquote> + +<p> +Note that <tt>remove_gp_faf()</tt> simply invokes +<tt>kfree_rcu()</tt> and proceeds, without any need to pay any +further attention to the subsequent grace period and <tt>kfree()</tt>. +It is permissible to invoke <tt>kfree_rcu()</tt> from the same +environments as for <tt>call_rcu()</tt>. +Interestingly enough, DYNIX/ptx had the equivalents of +<tt>call_rcu()</tt> and <tt>kfree_rcu()</tt>, but not +<tt>synchronize_rcu()</tt>. +This was due to the fact that RCU was not heavily used within DYNIX/ptx, +so the very few places that needed something like +<tt>synchronize_rcu()</tt> simply open-coded it. + +<p><a name="Quick Quiz 13"><b>Quick Quiz 13</b>:</a> +Earlier it was claimed that <tt>call_rcu()</tt> and +<tt>kfree_rcu()</tt> allowed updaters to avoid being blocked +by readers. +But how can that be correct, given that the invocation of the callback +and the freeing of the memory (respectively) must still wait for +a grace period to elapse? +<br><a href="#qq13answer">Answer</a> + +<p> +But what if the updater must wait for the completion of code to be +executed after the end of the grace period, but has other tasks +that can be carried out in the meantime? +The polling-style <tt>get_state_synchronize_rcu()</tt> and +<tt>cond_synchronize_rcu()</tt> functions may be used for this +purpose, as shown below: + +<blockquote> +<pre> + 1 bool remove_gp_poll(void) + 2 { + 3 struct foo *p; + 4 unsigned long s; + 5 + 6 spin_lock(&gp_lock); + 7 p = rcu_access_pointer(gp); + 8 if (!p) { + 9 spin_unlock(&gp_lock); +10 return false; +11 } +12 rcu_assign_pointer(gp, NULL); +13 spin_unlock(&gp_lock); +14 s = get_state_synchronize_rcu(); +15 do_something_while_waiting(); +16 cond_synchronize_rcu(s); +17 kfree(p); +18 return true; +19 } +</pre> +</blockquote> + +<p> +On line 14, <tt>get_state_synchronize_rcu()</tt> obtains a +“cookie” from RCU, +then line 15 carries out other tasks, +and finally, line 16 returns immediately if a grace period has +elapsed in the meantime, but otherwise waits as required. +The need for <tt>get_state_synchronize_rcu</tt> and +<tt>cond_synchronize_rcu()</tt> has appeared quite recently, +so it is too early to tell whether they will stand the test of time. + +<p> +RCU thus provides a range of tools to allow updaters to strike the +required tradeoff between latency, flexibility and CPU overhead. + +<h3><a name="Composability">Composability</a></h3> + +<p> +Composability has received much attention in recent years, perhaps in part +due to the collision of multicore hardware with object-oriented techniques +designed in single-threaded environments for single-threaded use. +And in theory, RCU read-side critical sections may be composed, and in +fact may be nested arbitrarily deeply. +In practice, as with all real-world implementations of composable +constructs, there are limitations. + +<p> +Implementations of RCU for which <tt>rcu_read_lock()</tt> +and <tt>rcu_read_unlock()</tt> generate no code, such as +Linux-kernel RCU when <tt>CONFIG_PREEMPT=n</tt>, can be +nested arbitrarily deeply. +After all, there is no overhead. +Except that if all these instances of <tt>rcu_read_lock()</tt> +and <tt>rcu_read_unlock()</tt> are visible to the compiler, +compilation will eventually fail due to exhausting memory, +mass storage, or user patience, whichever comes first. +If the nesting is not visible to the compiler, as is the case with +mutually recursive functions each in its own translation unit, +stack overflow will result. +If the nesting takes the form of loops, either the control variable +will overflow or (in the Linux kernel) you will get an RCU CPU stall warning. +Nevertheless, this class of RCU implementations is one +of the most composable constructs in existence. + +<p> +RCU implementations that explicitly track nesting depth +are limited by the nesting-depth counter. +For example, the Linux kernel's preemptible RCU limits nesting to +<tt>INT_MAX</tt>. +This should suffice for almost all practical purposes. +That said, a consecutive pair of RCU read-side critical sections +between which there is an operation that waits for a grace period +cannot be enclosed in another RCU read-side critical section. +This is because it is not legal to wait for a grace period within +an RCU read-side critical section: To do so would result either +in deadlock or +in RCU implicitly splitting the enclosing RCU read-side critical +section, neither of which is conducive to a long-lived and prosperous +kernel. + +<p> +It is worth noting that RCU is not alone in limiting composability. +For example, many transactional-memory implementations prohibit +composing a pair of transactions separated by an irrevocable +operation (for example, a network receive operation). +For another example, lock-based critical sections can be composed +surprisingly freely, but only if deadlock is avoided. + +<p> +In short, although RCU read-side critical sections are highly composable, +care is required in some situations, just as is the case for any other +composable synchronization mechanism. + +<h3><a name="Corner Cases">Corner Cases</a></h3> + +<p> +A given RCU workload might have an endless and intense stream of +RCU read-side critical sections, perhaps even so intense that there +was never a point in time during which there was not at least one +RCU read-side critical section in flight. +RCU cannot allow this situation to block grace periods: As long as +all the RCU read-side critical sections are finite, grace periods +must also be finite. + +<p> +That said, preemptible RCU implementations could potentially result +in RCU read-side critical sections being preempted for long durations, +which has the effect of creating a long-duration RCU read-side +critical section. +This situation can arise only in heavily loaded systems, but systems using +real-time priorities are of course more vulnerable. +Therefore, RCU priority boosting is provided to help deal with this +case. +That said, the exact requirements on RCU priority boosting will likely +evolve as more experience accumulates. + +<p> +Other workloads might have very high update rates. +Although one can argue that such workloads should instead use +something other than RCU, the fact remains that RCU must +handle such workloads gracefully. +This requirement is another factor driving batching of grace periods, +but it is also the driving force behind the checks for large numbers +of queued RCU callbacks in the <tt>call_rcu()</tt> code path. +Finally, high update rates should not delay RCU read-side critical +sections, although some read-side delays can occur when using +<tt>synchronize_rcu_expedited()</tt>, courtesy of this function's use +of <tt>try_stop_cpus()</tt>. +(In the future, <tt>synchronize_rcu_expedited()</tt> will be +converted to use lighter-weight inter-processor interrupts (IPIs), +but this will still disturb readers, though to a much smaller degree.) + +<p> +Although all three of these corner cases were understood in the early +1990s, a simple user-level test consisting of <tt>close(open(path))</tt> +in a tight loop +in the early 2000s suddenly provided a much deeper appreciation of the +high-update-rate corner case. +This test also motivated addition of some RCU code to react to high update +rates, for example, if a given CPU finds itself with more than 10,000 +RCU callbacks queued, it will cause RCU to take evasive action by +more aggressively starting grace periods and more aggressively forcing +completion of grace-period processing. +This evasive action causes the grace period to complete more quickly, +but at the cost of restricting RCU's batching optimizations, thus +increasing the CPU overhead incurred by that grace period. + +<h2><a name="Software-Engineering Requirements"> +Software-Engineering Requirements</a></h2> + +<p> +Between Murphy's Law and “To err is human”, it is necessary to +guard against mishaps and misuse: + +<ol> +<li> It is all too easy to forget to use <tt>rcu_read_lock()</tt> + everywhere that it is needed, so kernels built with + <tt>CONFIG_PROVE_RCU=y</tt> will spat if + <tt>rcu_dereference()</tt> is used outside of an + RCU read-side critical section. + Update-side code can use <tt>rcu_dereference_protected()</tt>, + which takes a + <a href="https://lwn.net/Articles/371986/">lockdep expression</a> + to indicate what is providing the protection. + If the indicated protection is not provided, a lockdep splat + is emitted. + + <p> + Code shared between readers and updaters can use + <tt>rcu_dereference_check()</tt>, which also takes a + lockdep expression, and emits a lockdep splat if neither + <tt>rcu_read_lock()</tt> nor the indicated protection + is in place. + In addition, <tt>rcu_dereference_raw()</tt> is used in those + (hopefully rare) cases where the required protection cannot + be easily described. + Finally, <tt>rcu_read_lock_held()</tt> is provided to + allow a function to verify that it has been invoked within + an RCU read-side critical section. + I was made aware of this set of requirements shortly after Thomas + Gleixner audited a number of RCU uses. +<li> A given function might wish to check for RCU-related preconditions + upon entry, before using any other RCU API. + The <tt>rcu_lockdep_assert()</tt> does this job, + asserting the expression in kernels having lockdep enabled + and doing nothing otherwise. +<li> It is also easy to forget to use <tt>rcu_assign_pointer()</tt> + and <tt>rcu_dereference()</tt>, perhaps (incorrectly) + substituting a simple assignment. + To catch this sort of error, a given RCU-protected pointer may be + tagged with <tt>__rcu</tt>, after which running sparse + with <tt>CONFIG_SPARSE_RCU_POINTER=y</tt> will complain + about simple-assignment accesses to that pointer. + Arnd Bergmann made me aware of this requirement, and also + supplied the needed + <a href="https://lwn.net/Articles/376011/">patch series</a>. +<li> Kernels built with <tt>CONFIG_DEBUG_OBJECTS_RCU_HEAD=y</tt> + will splat if a data element is passed to <tt>call_rcu()</tt> + twice in a row, without a grace period in between. + (This error is similar to a double free.) + The corresponding <tt>rcu_head</tt> structures that are + dynamically allocated are automatically tracked, but + <tt>rcu_head</tt> structures allocated on the stack + must be initialized with <tt>init_rcu_head_on_stack()</tt> + and cleaned up with <tt>destroy_rcu_head_on_stack()</tt>. + Similarly, statically allocated non-stack <tt>rcu_head</tt> + structures must be initialized with <tt>init_rcu_head()</tt> + and cleaned up with <tt>destroy_rcu_head()</tt>. + Mathieu Desnoyers made me aware of this requirement, and also + supplied the needed + <a href="https://lkml.kernel.org/g/20100319013024.GA28456@Krystal">patch</a>. +<li> An infinite loop in an RCU read-side critical section will + eventually trigger an RCU CPU stall warning splat, with + the duration of “eventually” being controlled by the + <tt>RCU_CPU_STALL_TIMEOUT</tt> <tt>Kconfig</tt> option, or, + alternatively, by the + <tt>rcupdate.rcu_cpu_stall_timeout</tt> boot/sysfs + parameter. + However, RCU is not obligated to produce this splat + unless there is a grace period waiting on that particular + RCU read-side critical section. + <p> + Some extreme workloads might intentionally delay + RCU grace periods, and systems running those workloads can + be booted with <tt>rcupdate.rcu_cpu_stall_suppress</tt> + to suppress the splats. + This kernel parameter may also be set via <tt>sysfs</tt>. + Furthermore, RCU CPU stall warnings are counter-productive + during sysrq dumps and during panics. + RCU therefore supplies the <tt>rcu_sysrq_start()</tt> and + <tt>rcu_sysrq_end()</tt> API members to be called before + and after long sysrq dumps. + RCU also supplies the <tt>rcu_panic()</tt> notifier that is + automatically invoked at the beginning of a panic to suppress + further RCU CPU stall warnings. + + <p> + This requirement made itself known in the early 1990s, pretty + much the first time that it was necessary to debug a CPU stall. + That said, the initial implementation in DYNIX/ptx was quite + generic in comparison with that of Linux. +<li> Although it would be very good to detect pointers leaking out + of RCU read-side critical sections, there is currently no + good way of doing this. + One complication is the need to distinguish between pointers + leaking and pointers that have been handed off from RCU to + some other synchronization mechanism, for example, reference + counting. +<li> In kernels built with <tt>CONFIG_RCU_TRACE=y</tt>, RCU-related + information is provided via both debugfs and event tracing. +<li> Open-coded use of <tt>rcu_assign_pointer()</tt> and + <tt>rcu_dereference()</tt> to create typical linked + data structures can be surprisingly error-prone. + Therefore, RCU-protected + <a href="https://lwn.net/Articles/609973/#RCU List APIs">linked lists</a> + and, more recently, RCU-protected + <a href="https://lwn.net/Articles/612100/">hash tables</a> + are available. + Many other special-purpose RCU-protected data structures are + available in the Linux kernel and the userspace RCU library. +<li> Some linked structures are created at compile time, but still + require <tt>__rcu</tt> checking. + The <tt>RCU_POINTER_INITIALIZER()</tt> macro serves this + purpose. +<li> It is not necessary to use <tt>rcu_assign_pointer()</tt> + when creating linked structures that are to be published via + a single external pointer. + The <tt>RCU_INIT_POINTER()</tt> macro is provided for + this task and also for assigning <tt>NULL</tt> pointers + at runtime. +</ol> + +<p> +This not a hard-and-fast list: RCU's diagnostic capabilities will +continue to be guided by the number and type of usage bugs found +in real-world RCU usage. + +<h2><a name="Linux Kernel Complications">Linux Kernel Complications</a></h2> + +<p> +The Linux kernel provides an interesting environment for all kinds of +software, including RCU. +Some of the relevant points of interest are as follows: + +<ol> +<li> <a href="#Configuration">Configuration</a>. +<li> <a href="#Firmware Interface">Firmware Interface</a>. +<li> <a href="#Early Boot">Early Boot</a>. +<li> <a href="#Interrupts and NMIs"> + Interrupts and non-maskable interrupts (NMIs)</a>. +<li> <a href="#Loadable Modules">Loadable Modules</a>. +<li> <a href="#Hotplug CPU">Hotplug CPU</a>. +<li> <a href="#Scheduler and RCU">Scheduler and RCU</a>. +<li> <a href="#Tracing and RCU">Tracing and RCU</a>. +<li> <a href="#Energy Efficiency">Energy Efficiency</a>. +<li> <a href="#Memory Efficiency">Memory Efficiency</a>. +<li> <a href="#Performance, Scalability, Response Time, and Reliability"> + Performance, Scalability, Response Time, and Reliability</a>. +</ol> + +<p> +This list is probably incomplete, but it does give a feel for the +most notable Linux-kernel complications. +Each of the following sections covers one of the above topics. + +<h3><a name="Configuration">Configuration</a></h3> + +<p> +RCU's goal is automatic configuration, so that almost nobody +needs to worry about RCU's <tt>Kconfig</tt> options. +And for almost all users, RCU does in fact work well +“out of the box.” + +<p> +However, there are specialized use cases that are handled by +kernel boot parameters and <tt>Kconfig</tt> options. +Unfortunately, the <tt>Kconfig</tt> system will explicitly ask users +about new <tt>Kconfig</tt> options, which requires almost all of them +be hidden behind a <tt>CONFIG_RCU_EXPERT</tt> <tt>Kconfig</tt> option. + +<p> +This all should be quite obvious, but the fact remains that +Linus Torvalds recently had to +<a href="https://lkml.kernel.org/g/CA+55aFy4wcCwaL4okTs8wXhGZ5h-ibecy_Meg9C4MNQrUnwMcg@mail.gmail.com">remind</a> +me of this requirement. + +<h3><a name="Firmware Interface">Firmware Interface</a></h3> + +<p> +In many cases, kernel obtains information about the system from the +firmware, and sometimes things are lost in translation. +Or the translation is accurate, but the original message is bogus. + +<p> +For example, some systems' firmware overreports the number of CPUs, +sometimes by a large factor. +If RCU naively believed the firmware, as it used to do, +it would create too many per-CPU kthreads. +Although the resulting system will still run correctly, the extra +kthreads needlessly consume memory and can cause confusion +when they show up in <tt>ps</tt> listings. + +<p> +RCU must therefore wait for a given CPU to actually come online before +it can allow itself to believe that the CPU actually exists. +The resulting “ghost CPUs” (which are never going to +come online) cause a number of +<a href="https://paulmck.livejournal.com/37494.html">interesting complications</a>. + +<h3><a name="Early Boot">Early Boot</a></h3> + +<p> +The Linux kernel's boot sequence is an interesting process, +and RCU is used early, even before <tt>rcu_init()</tt> +is invoked. +In fact, a number of RCU's primitives can be used as soon as the +initial task's <tt>task_struct</tt> is available and the +boot CPU's per-CPU variables are set up. +The read-side primitives (<tt>rcu_read_lock()</tt>, +<tt>rcu_read_unlock()</tt>, <tt>rcu_dereference()</tt>, +and <tt>rcu_access_pointer()</tt>) will operate normally very early on, +as will <tt>rcu_assign_pointer()</tt>. + +<p> +Although <tt>call_rcu()</tt> may be invoked at any +time during boot, callbacks are not guaranteed to be invoked until after +the scheduler is fully up and running. +This delay in callback invocation is due to the fact that RCU does not +invoke callbacks until it is fully initialized, and this full initialization +cannot occur until after the scheduler has initialized itself to the +point where RCU can spawn and run its kthreads. +In theory, it would be possible to invoke callbacks earlier, +however, this is not a panacea because there would be severe restrictions +on what operations those callbacks could invoke. + +<p> +Perhaps surprisingly, <tt>synchronize_rcu()</tt>, +<a href="#Bottom-Half Flavor"><tt>synchronize_rcu_bh()</tt></a> +(<a href="#Bottom-Half Flavor">discussed below</a>), +and +<a href="#Sched Flavor"><tt>synchronize_sched()</tt></a> +will all operate normally +during very early boot, the reason being that there is only one CPU +and preemption is disabled. +This means that the call <tt>synchronize_rcu()</tt> (or friends) +itself is a quiescent +state and thus a grace period, so the early-boot implementation can +be a no-op. + +<p> +Both <tt>synchronize_rcu_bh()</tt> and <tt>synchronize_sched()</tt> +continue to operate normally through the remainder of boot, courtesy +of the fact that preemption is disabled across their RCU read-side +critical sections and also courtesy of the fact that there is still +only one CPU. +However, once the scheduler starts initializing, preemption is enabled. +There is still only a single CPU, but the fact that preemption is enabled +means that the no-op implementation of <tt>synchronize_rcu()</tt> no +longer works in <tt>CONFIG_PREEMPT=y</tt> kernels. +Therefore, as soon as the scheduler starts initializing, the early-boot +fastpath is disabled. +This means that <tt>synchronize_rcu()</tt> switches to its runtime +mode of operation where it posts callbacks, which in turn means that +any call to <tt>synchronize_rcu()</tt> will block until the corresponding +callback is invoked. +Unfortunately, the callback cannot be invoked until RCU's runtime +grace-period machinery is up and running, which cannot happen until +the scheduler has initialized itself sufficiently to allow RCU's +kthreads to be spawned. +Therefore, invoking <tt>synchronize_rcu()</tt> during scheduler +initialization can result in deadlock. + +<p><a name="Quick Quiz 14"><b>Quick Quiz 14</b>:</a> +So what happens with <tt>synchronize_rcu()</tt> during +scheduler initialization for <tt>CONFIG_PREEMPT=n</tt> +kernels? +<br><a href="#qq14answer">Answer</a> + +<p> +I learned of these boot-time requirements as a result of a series of +system hangs. + +<h3><a name="Interrupts and NMIs">Interrupts and NMIs</a></h3> + +<p> +The Linux kernel has interrupts, and RCU read-side critical sections are +legal within interrupt handlers and within interrupt-disabled regions +of code, as are invocations of <tt>call_rcu()</tt>. + +<p> +Some Linux-kernel architectures can enter an interrupt handler from +non-idle process context, and then just never leave it, instead stealthily +transitioning back to process context. +This trick is sometimes used to invoke system calls from inside the kernel. +These “half-interrupts” mean that RCU has to be very careful +about how it counts interrupt nesting levels. +I learned of this requirement the hard way during a rewrite +of RCU's dyntick-idle code. + +<p> +The Linux kernel has non-maskable interrupts (NMIs), and +RCU read-side critical sections are legal within NMI handlers. +Thankfully, RCU update-side primitives, including +<tt>call_rcu()</tt>, are prohibited within NMI handlers. + +<p> +The name notwithstanding, some Linux-kernel architectures +can have nested NMIs, which RCU must handle correctly. +Andy Lutomirski +<a href="https://lkml.kernel.org/g/CALCETrXLq1y7e_dKFPgou-FKHB6Pu-r8+t-6Ds+8=va7anBWDA@mail.gmail.com">surprised me</a> +with this requirement; +he also kindly surprised me with +<a href="https://lkml.kernel.org/g/CALCETrXSY9JpW3uE6H8WYk81sg56qasA2aqmjMPsq5dOtzso=g@mail.gmail.com">an algorithm</a> +that meets this requirement. + +<h3><a name="Loadable Modules">Loadable Modules</a></h3> + +<p> +The Linux kernel has loadable modules, and these modules can +also be unloaded. +After a given module has been unloaded, any attempt to call +one of its functions results in a segmentation fault. +The module-unload functions must therefore cancel any +delayed calls to loadable-module functions, for example, +any outstanding <tt>mod_timer()</tt> must be dealt with +via <tt>del_timer_sync()</tt> or similar. + +<p> +Unfortunately, there is no way to cancel an RCU callback; +once you invoke <tt>call_rcu()</tt>, the callback function is +going to eventually be invoked, unless the system goes down first. +Because it is normally considered socially irresponsible to crash the system +in response to a module unload request, we need some other way +to deal with in-flight RCU callbacks. + +<p> +RCU therefore provides +<tt><a href="https://lwn.net/Articles/217484/">rcu_barrier()</a></tt>, +which waits until all in-flight RCU callbacks have been invoked. +If a module uses <tt>call_rcu()</tt>, its exit function should therefore +prevent any future invocation of <tt>call_rcu()</tt>, then invoke +<tt>rcu_barrier()</tt>. +In theory, the underlying module-unload code could invoke +<tt>rcu_barrier()</tt> unconditionally, but in practice this would +incur unacceptable latencies. + +<p> +Nikita Danilov noted this requirement for an analogous filesystem-unmount +situation, and Dipankar Sarma incorporated <tt>rcu_barrier()</tt> into RCU. +The need for <tt>rcu_barrier()</tt> for module unloading became +apparent later. + +<h3><a name="Hotplug CPU">Hotplug CPU</a></h3> + +<p> +The Linux kernel supports CPU hotplug, which means that CPUs +can come and go. +It is of course illegal to use any RCU API member from an offline CPU. +This requirement was present from day one in DYNIX/ptx, but +on the other hand, the Linux kernel's CPU-hotplug implementation +is “interesting.” + +<p> +The Linux-kernel CPU-hotplug implementation has notifiers that +are used to allow the various kernel subsystems (including RCU) +to respond appropriately to a given CPU-hotplug operation. +Most RCU operations may be invoked from CPU-hotplug notifiers, +including even normal synchronous grace-period operations +such as <tt>synchronize_rcu()</tt>. +However, expedited grace-period operations such as +<tt>synchronize_rcu_expedited()</tt> are not supported, +due to the fact that current implementations block CPU-hotplug +operations, which could result in deadlock. + +<p> +In addition, all-callback-wait operations such as +<tt>rcu_barrier()</tt> are also not supported, due to the +fact that there are phases of CPU-hotplug operations where +the outgoing CPU's callbacks will not be invoked until after +the CPU-hotplug operation ends, which could also result in deadlock. + +<h3><a name="Scheduler and RCU">Scheduler and RCU</a></h3> + +<p> +RCU depends on the scheduler, and the scheduler uses RCU to +protect some of its data structures. +This means the scheduler is forbidden from acquiring +the runqueue locks and the priority-inheritance locks +in the middle of an outermost RCU read-side critical section unless either +(1) it releases them before exiting that same +RCU read-side critical section, or +(2) interrupts are disabled across +that entire RCU read-side critical section. +This same prohibition also applies (recursively!) to any lock that is acquired +while holding any lock to which this prohibition applies. +Adhering to this rule prevents preemptible RCU from invoking +<tt>rcu_read_unlock_special()</tt> while either runqueue or +priority-inheritance locks are held, thus avoiding deadlock. + +<p> +Prior to v4.4, it was only necessary to disable preemption across +RCU read-side critical sections that acquired scheduler locks. +In v4.4, expedited grace periods started using IPIs, and these +IPIs could force a <tt>rcu_read_unlock()</tt> to take the slowpath. +Therefore, this expedited-grace-period change required disabling of +interrupts, not just preemption. + +<p> +For RCU's part, the preemptible-RCU <tt>rcu_read_unlock()</tt> +implementation must be written carefully to avoid similar deadlocks. +In particular, <tt>rcu_read_unlock()</tt> must tolerate an +interrupt where the interrupt handler invokes both +<tt>rcu_read_lock()</tt> and <tt>rcu_read_unlock()</tt>. +This possibility requires <tt>rcu_read_unlock()</tt> to use +negative nesting levels to avoid destructive recursion via +interrupt handler's use of RCU. + +<p> +This pair of mutual scheduler-RCU requirements came as a +<a href="https://lwn.net/Articles/453002/">complete surprise</a>. + +<p> +As noted above, RCU makes use of kthreads, and it is necessary to +avoid excessive CPU-time accumulation by these kthreads. +This requirement was no surprise, but RCU's violation of it +when running context-switch-heavy workloads when built with +<tt>CONFIG_NO_HZ_FULL=y</tt> +<a href="http://www.rdrop.com/users/paulmck/scalability/paper/BareMetal.2015.01.15b.pdf">did come as a surprise [PDF]</a>. +RCU has made good progress towards meeting this requirement, even +for context-switch-have <tt>CONFIG_NO_HZ_FULL=y</tt> workloads, +but there is room for further improvement. + +<h3><a name="Tracing and RCU">Tracing and RCU</a></h3> + +<p> +It is possible to use tracing on RCU code, but tracing itself +uses RCU. +For this reason, <tt>rcu_dereference_raw_notrace()</tt> +is provided for use by tracing, which avoids the destructive +recursion that could otherwise ensue. +This API is also used by virtualization in some architectures, +where RCU readers execute in environments in which tracing +cannot be used. +The tracing folks both located the requirement and provided the +needed fix, so this surprise requirement was relatively painless. + +<h3><a name="Energy Efficiency">Energy Efficiency</a></h3> + +<p> +Interrupting idle CPUs is considered socially unacceptable, +especially by people with battery-powered embedded systems. +RCU therefore conserves energy by detecting which CPUs are +idle, including tracking CPUs that have been interrupted from idle. +This is a large part of the energy-efficiency requirement, +so I learned of this via an irate phone call. + +<p> +Because RCU avoids interrupting idle CPUs, it is illegal to +execute an RCU read-side critical section on an idle CPU. +(Kernels built with <tt>CONFIG_PROVE_RCU=y</tt> will splat +if you try it.) +The <tt>RCU_NONIDLE()</tt> macro and <tt>_rcuidle</tt> +event tracing is provided to work around this restriction. +In addition, <tt>rcu_is_watching()</tt> may be used to +test whether or not it is currently legal to run RCU read-side +critical sections on this CPU. +I learned of the need for diagnostics on the one hand +and <tt>RCU_NONIDLE()</tt> on the other while inspecting +idle-loop code. +Steven Rostedt supplied <tt>_rcuidle</tt> event tracing, +which is used quite heavily in the idle loop. + +<p> +It is similarly socially unacceptable to interrupt an +<tt>nohz_full</tt> CPU running in userspace. +RCU must therefore track <tt>nohz_full</tt> userspace +execution. +And in +<a href="https://lwn.net/Articles/558284/"><tt>CONFIG_NO_HZ_FULL_SYSIDLE=y</tt></a> +kernels, RCU must separately track idle CPUs on the one hand and +CPUs that are either idle or executing in userspace on the other. +In both cases, RCU must be able to sample state at two points in +time, and be able to determine whether or not some other CPU spent +any time idle and/or executing in userspace. + +<p> +These energy-efficiency requirements have proven quite difficult to +understand and to meet, for example, there have been more than five +clean-sheet rewrites of RCU's energy-efficiency code, the last of +which was finally able to demonstrate +<a href="http://www.rdrop.com/users/paulmck/realtime/paper/AMPenergy.2013.04.19a.pdf">real energy savings running on real hardware [PDF]</a>. +As noted earlier, +I learned of many of these requirements via angry phone calls: +Flaming me on the Linux-kernel mailing list was apparently not +sufficient to fully vent their ire at RCU's energy-efficiency bugs! + +<h3><a name="Memory Efficiency">Memory Efficiency</a></h3> + +<p> +Although small-memory non-realtime systems can simply use Tiny RCU, +code size is only one aspect of memory efficiency. +Another aspect is the size of the <tt>rcu_head</tt> structure +used by <tt>call_rcu()</tt> and <tt>kfree_rcu()</tt>. +Although this structure contains nothing more than a pair of pointers, +it does appear in many RCU-protected data structures, including +some that are size critical. +The <tt>page</tt> structure is a case in point, as evidenced by +the many occurrences of the <tt>union</tt> keyword within that structure. + +<p> +This need for memory efficiency is one reason that RCU uses hand-crafted +singly linked lists to track the <tt>rcu_head</tt> structures that +are waiting for a grace period to elapse. +It is also the reason why <tt>rcu_head</tt> structures do not contain +debug information, such as fields tracking the file and line of the +<tt>call_rcu()</tt> or <tt>kfree_rcu()</tt> that posted them. +Although this information might appear in debug-only kernel builds at some +point, in the meantime, the <tt>->func</tt> field will often provide +the needed debug information. + +<p> +However, in some cases, the need for memory efficiency leads to even +more extreme measures. +Returning to the <tt>page</tt> structure, the <tt>rcu_head</tt> field +shares storage with a great many other structures that are used at +various points in the corresponding page's lifetime. +In order to correctly resolve certain +<a href="https://lkml.kernel.org/g/1439976106-137226-1-git-send-email-kirill.shutemov@linux.intel.com">race conditions</a>, +the Linux kernel's memory-management subsystem needs a particular bit +to remain zero during all phases of grace-period processing, +and that bit happens to map to the bottom bit of the +<tt>rcu_head</tt> structure's <tt>->next</tt> field. +RCU makes this guarantee as long as <tt>call_rcu()</tt> +is used to post the callback, as opposed to <tt>kfree_rcu()</tt> +or some future “lazy” +variant of <tt>call_rcu()</tt> that might one day be created for +energy-efficiency purposes. + +<h3><a name="Performance, Scalability, Response Time, and Reliability"> +Performance, Scalability, Response Time, and Reliability</a></h3> + +<p> +Expanding on the +<a href="#Performance and Scalability">earlier discussion</a>, +RCU is used heavily by hot code paths in performance-critical +portions of the Linux kernel's networking, security, virtualization, +and scheduling code paths. +RCU must therefore use efficient implementations, especially in its +read-side primitives. +To that end, it would be good if preemptible RCU's implementation +of <tt>rcu_read_lock()</tt> could be inlined, however, doing +this requires resolving <tt>#include</tt> issues with the +<tt>task_struct</tt> structure. + +<p> +The Linux kernel supports hardware configurations with up to +4096 CPUs, which means that RCU must be extremely scalable. +Algorithms that involve frequent acquisitions of global locks or +frequent atomic operations on global variables simply cannot be +tolerated within the RCU implementation. +RCU therefore makes heavy use of a combining tree based on the +<tt>rcu_node</tt> structure. +RCU is required to tolerate all CPUs continuously invoking any +combination of RCU's runtime primitives with minimal per-operation +overhead. +In fact, in many cases, increasing load must <i>decrease</i> the +per-operation overhead, witness the batching optimizations for +<tt>synchronize_rcu()</tt>, <tt>call_rcu()</tt>, +<tt>synchronize_rcu_expedited()</tt>, and <tt>rcu_barrier()</tt>. +As a general rule, RCU must cheerfully accept whatever the +rest of the Linux kernel decides to throw at it. + +<p> +The Linux kernel is used for real-time workloads, especially +in conjunction with the +<a href="https://rt.wiki.kernel.org/index.php/Main_Page">-rt patchset</a>. +The real-time-latency response requirements are such that the +traditional approach of disabling preemption across RCU +read-side critical sections is inappropriate. +Kernels built with <tt>CONFIG_PREEMPT=y</tt> therefore +use an RCU implementation that allows RCU read-side critical +sections to be preempted. +This requirement made its presence known after users made it +clear that an earlier +<a href="https://lwn.net/Articles/107930/">real-time patch</a> +did not meet their needs, in conjunction with some +<a href="https://lkml.kernel.org/g/20050318002026.GA2693@us.ibm.com">RCU issues</a> +encountered by a very early version of the -rt patchset. + +<p> +In addition, RCU must make do with a sub-100-microsecond real-time latency +budget. +In fact, on smaller systems with the -rt patchset, the Linux kernel +provides sub-20-microsecond real-time latencies for the whole kernel, +including RCU. +RCU's scalability and latency must therefore be sufficient for +these sorts of configurations. +To my surprise, the sub-100-microsecond real-time latency budget +<a href="http://www.rdrop.com/users/paulmck/realtime/paper/bigrt.2013.01.31a.LCA.pdf"> +applies to even the largest systems [PDF]</a>, +up to and including systems with 4096 CPUs. +This real-time requirement motivated the grace-period kthread, which +also simplified handling of a number of race conditions. + +<p> +Finally, RCU's status as a synchronization primitive means that +any RCU failure can result in arbitrary memory corruption that can be +extremely difficult to debug. +This means that RCU must be extremely reliable, which in +practice also means that RCU must have an aggressive stress-test +suite. +This stress-test suite is called <tt>rcutorture</tt>. + +<p> +Although the need for <tt>rcutorture</tt> was no surprise, +the current immense popularity of the Linux kernel is posing +interesting—and perhaps unprecedented—validation +challenges. +To see this, keep in mind that there are well over one billion +instances of the Linux kernel running today, given Android +smartphones, Linux-powered televisions, and servers. +This number can be expected to increase sharply with the advent of +the celebrated Internet of Things. + +<p> +Suppose that RCU contains a race condition that manifests on average +once per million years of runtime. +This bug will be occurring about three times per <i>day</i> across +the installed base. +RCU could simply hide behind hardware error rates, given that no one +should really expect their smartphone to last for a million years. +However, anyone taking too much comfort from this thought should +consider the fact that in most jurisdictions, a successful multi-year +test of a given mechanism, which might include a Linux kernel, +suffices for a number of types of safety-critical certifications. +In fact, rumor has it that the Linux kernel is already being used +in production for safety-critical applications. +I don't know about you, but I would feel quite bad if a bug in RCU +killed someone. +Which might explain my recent focus on validation and verification. + +<h2><a name="Other RCU Flavors">Other RCU Flavors</a></h2> + +<p> +One of the more surprising things about RCU is that there are now +no fewer than five <i>flavors</i>, or API families. +In addition, the primary flavor that has been the sole focus up to +this point has two different implementations, non-preemptible and +preemptible. +The other four flavors are listed below, with requirements for each +described in a separate section. + +<ol> +<li> <a href="#Bottom-Half Flavor">Bottom-Half Flavor</a> +<li> <a href="#Sched Flavor">Sched Flavor</a> +<li> <a href="#Sleepable RCU">Sleepable RCU</a> +<li> <a href="#Tasks RCU">Tasks RCU</a> +</ol> + +<h3><a name="Bottom-Half Flavor">Bottom-Half Flavor</a></h3> + +<p> +The softirq-disable (AKA “bottom-half”, +hence the “_bh” abbreviations) +flavor of RCU, or <i>RCU-bh</i>, was developed by +Dipankar Sarma to provide a flavor of RCU that could withstand the +network-based denial-of-service attacks researched by Robert +Olsson. +These attacks placed so much networking load on the system +that some of the CPUs never exited softirq execution, +which in turn prevented those CPUs from ever executing a context switch, +which, in the RCU implementation of that time, prevented grace periods +from ever ending. +The result was an out-of-memory condition and a system hang. + +<p> +The solution was the creation of RCU-bh, which does +<tt>local_bh_disable()</tt> +across its read-side critical sections, and which uses the transition +from one type of softirq processing to another as a quiescent state +in addition to context switch, idle, user mode, and offline. +This means that RCU-bh grace periods can complete even when some of +the CPUs execute in softirq indefinitely, thus allowing algorithms +based on RCU-bh to withstand network-based denial-of-service attacks. + +<p> +Because +<tt>rcu_read_lock_bh()</tt> and <tt>rcu_read_unlock_bh()</tt> +disable and re-enable softirq handlers, any attempt to start a softirq +handlers during the +RCU-bh read-side critical section will be deferred. +In this case, <tt>rcu_read_unlock_bh()</tt> +will invoke softirq processing, which can take considerable time. +One can of course argue that this softirq overhead should be associated +with the code following the RCU-bh read-side critical section rather +than <tt>rcu_read_unlock_bh()</tt>, but the fact +is that most profiling tools cannot be expected to make this sort +of fine distinction. +For example, suppose that a three-millisecond-long RCU-bh read-side +critical section executes during a time of heavy networking load. +There will very likely be an attempt to invoke at least one softirq +handler during that three milliseconds, but any such invocation will +be delayed until the time of the <tt>rcu_read_unlock_bh()</tt>. +This can of course make it appear at first glance as if +<tt>rcu_read_unlock_bh()</tt> was executing very slowly. + +<p> +The +<a href="https://lwn.net/Articles/609973/#RCU Per-Flavor API Table">RCU-bh API</a> +includes +<tt>rcu_read_lock_bh()</tt>, +<tt>rcu_read_unlock_bh()</tt>, +<tt>rcu_dereference_bh()</tt>, +<tt>rcu_dereference_bh_check()</tt>, +<tt>synchronize_rcu_bh()</tt>, +<tt>synchronize_rcu_bh_expedited()</tt>, +<tt>call_rcu_bh()</tt>, +<tt>rcu_barrier_bh()</tt>, and +<tt>rcu_read_lock_bh_held()</tt>. + +<h3><a name="Sched Flavor">Sched Flavor</a></h3> + +<p> +Before preemptible RCU, waiting for an RCU grace period had the +side effect of also waiting for all pre-existing interrupt +and NMI handlers. +However, there are legitimate preemptible-RCU implementations that +do not have this property, given that any point in the code outside +of an RCU read-side critical section can be a quiescent state. +Therefore, <i>RCU-sched</i> was created, which follows “classic” +RCU in that an RCU-sched grace period waits for for pre-existing +interrupt and NMI handlers. +In kernels built with <tt>CONFIG_PREEMPT=n</tt>, the RCU and RCU-sched +APIs have identical implementations, while kernels built with +<tt>CONFIG_PREEMPT=y</tt> provide a separate implementation for each. + +<p> +Note well that in <tt>CONFIG_PREEMPT=y</tt> kernels, +<tt>rcu_read_lock_sched()</tt> and <tt>rcu_read_unlock_sched()</tt> +disable and re-enable preemption, respectively. +This means that if there was a preemption attempt during the +RCU-sched read-side critical section, <tt>rcu_read_unlock_sched()</tt> +will enter the scheduler, with all the latency and overhead entailed. +Just as with <tt>rcu_read_unlock_bh()</tt>, this can make it look +as if <tt>rcu_read_unlock_sched()</tt> was executing very slowly. +However, the highest-priority task won't be preempted, so that task +will enjoy low-overhead <tt>rcu_read_unlock_sched()</tt> invocations. + +<p> +The +<a href="https://lwn.net/Articles/609973/#RCU Per-Flavor API Table">RCU-sched API</a> +includes +<tt>rcu_read_lock_sched()</tt>, +<tt>rcu_read_unlock_sched()</tt>, +<tt>rcu_read_lock_sched_notrace()</tt>, +<tt>rcu_read_unlock_sched_notrace()</tt>, +<tt>rcu_dereference_sched()</tt>, +<tt>rcu_dereference_sched_check()</tt>, +<tt>synchronize_sched()</tt>, +<tt>synchronize_rcu_sched_expedited()</tt>, +<tt>call_rcu_sched()</tt>, +<tt>rcu_barrier_sched()</tt>, and +<tt>rcu_read_lock_sched_held()</tt>. +However, anything that disables preemption also marks an RCU-sched +read-side critical section, including +<tt>preempt_disable()</tt> and <tt>preempt_enable()</tt>, +<tt>local_irq_save()</tt> and <tt>local_irq_restore()</tt>, +and so on. + +<h3><a name="Sleepable RCU">Sleepable RCU</a></h3> + +<p> +For well over a decade, someone saying “I need to block within +an RCU read-side critical section” was a reliable indication +that this someone did not understand RCU. +After all, if you are always blocking in an RCU read-side critical +section, you can probably afford to use a higher-overhead synchronization +mechanism. +However, that changed with the advent of the Linux kernel's notifiers, +whose RCU read-side critical +sections almost never sleep, but sometimes need to. +This resulted in the introduction of +<a href="https://lwn.net/Articles/202847/">sleepable RCU</a>, +or <i>SRCU</i>. + +<p> +SRCU allows different domains to be defined, with each such domain +defined by an instance of an <tt>srcu_struct</tt> structure. +A pointer to this structure must be passed in to each SRCU function, +for example, <tt>synchronize_srcu(&ss)</tt>, where +<tt>ss</tt> is the <tt>srcu_struct</tt> structure. +The key benefit of these domains is that a slow SRCU reader in one +domain does not delay an SRCU grace period in some other domain. +That said, one consequence of these domains is that read-side code +must pass a “cookie” from <tt>srcu_read_lock()</tt> +to <tt>srcu_read_unlock()</tt>, for example, as follows: + +<blockquote> +<pre> + 1 int idx; + 2 + 3 idx = srcu_read_lock(&ss); + 4 do_something(); + 5 srcu_read_unlock(&ss, idx); +</pre> +</blockquote> + +<p> +As noted above, it is legal to block within SRCU read-side critical sections, +however, with great power comes great responsibility. +If you block forever in one of a given domain's SRCU read-side critical +sections, then that domain's grace periods will also be blocked forever. +Of course, one good way to block forever is to deadlock, which can +happen if any operation in a given domain's SRCU read-side critical +section can block waiting, either directly or indirectly, for that domain's +grace period to elapse. +For example, this results in a self-deadlock: + +<blockquote> +<pre> + 1 int idx; + 2 + 3 idx = srcu_read_lock(&ss); + 4 do_something(); + 5 synchronize_srcu(&ss); + 6 srcu_read_unlock(&ss, idx); +</pre> +</blockquote> + +<p> +However, if line 5 acquired a mutex that was held across +a <tt>synchronize_srcu()</tt> for domain <tt>ss</tt>, +deadlock would still be possible. +Furthermore, if line 5 acquired a mutex that was held across +a <tt>synchronize_srcu()</tt> for some other domain <tt>ss1</tt>, +and if an <tt>ss1</tt>-domain SRCU read-side critical section +acquired another mutex that was held across as <tt>ss</tt>-domain +<tt>synchronize_srcu()</tt>, +deadlock would again be possible. +Such a deadlock cycle could extend across an arbitrarily large number +of different SRCU domains. +Again, with great power comes great responsibility. + +<p> +Unlike the other RCU flavors, SRCU read-side critical sections can +run on idle and even offline CPUs. +This ability requires that <tt>srcu_read_lock()</tt> and +<tt>srcu_read_unlock()</tt> contain memory barriers, which means +that SRCU readers will run a bit slower than would RCU readers. +It also motivates the <tt>smp_mb__after_srcu_read_unlock()</tt> +API, which, in combination with <tt>srcu_read_unlock()</tt>, +guarantees a full memory barrier. + +<p> +The +<a href="https://lwn.net/Articles/609973/#RCU Per-Flavor API Table">SRCU API</a> +includes +<tt>srcu_read_lock()</tt>, +<tt>srcu_read_unlock()</tt>, +<tt>srcu_dereference()</tt>, +<tt>srcu_dereference_check()</tt>, +<tt>synchronize_srcu()</tt>, +<tt>synchronize_srcu_expedited()</tt>, +<tt>call_srcu()</tt>, +<tt>srcu_barrier()</tt>, and +<tt>srcu_read_lock_held()</tt>. +It also includes +<tt>DEFINE_SRCU()</tt>, +<tt>DEFINE_STATIC_SRCU()</tt>, and +<tt>init_srcu_struct()</tt> +APIs for defining and initializing <tt>srcu_struct</tt> structures. + +<h3><a name="Tasks RCU">Tasks RCU</a></h3> + +<p> +Some forms of tracing use “tramopolines” to handle the +binary rewriting required to install different types of probes. +It would be good to be able to free old trampolines, which sounds +like a job for some form of RCU. +However, because it is necessary to be able to install a trace +anywhere in the code, it is not possible to use read-side markers +such as <tt>rcu_read_lock()</tt> and <tt>rcu_read_unlock()</tt>. +In addition, it does not work to have these markers in the trampoline +itself, because there would need to be instructions following +<tt>rcu_read_unlock()</tt>. +Although <tt>synchronize_rcu()</tt> would guarantee that execution +reached the <tt>rcu_read_unlock()</tt>, it would not be able to +guarantee that execution had completely left the trampoline. + +<p> +The solution, in the form of +<a href="https://lwn.net/Articles/607117/"><i>Tasks RCU</i></a>, +is to have implicit +read-side critical sections that are delimited by voluntary context +switches, that is, calls to <tt>schedule()</tt>, +<tt>cond_resched_rcu_qs()</tt>, and +<tt>synchronize_rcu_tasks()</tt>. +In addition, transitions to and from userspace execution also delimit +tasks-RCU read-side critical sections. + +<p> +The tasks-RCU API is quite compact, consisting only of +<tt>call_rcu_tasks()</tt>, +<tt>synchronize_rcu_tasks()</tt>, and +<tt>rcu_barrier_tasks()</tt>. + +<h2><a name="Possible Future Changes">Possible Future Changes</a></h2> + +<p> +One of the tricks that RCU uses to attain update-side scalability is +to increase grace-period latency with increasing numbers of CPUs. +If this becomes a serious problem, it will be necessary to rework the +grace-period state machine so as to avoid the need for the additional +latency. + +<p> +Expedited grace periods scan the CPUs, so their latency and overhead +increases with increasing numbers of CPUs. +If this becomes a serious problem on large systems, it will be necessary +to do some redesign to avoid this scalability problem. + +<p> +RCU disables CPU hotplug in a few places, perhaps most notably in the +expedited grace-period and <tt>rcu_barrier()</tt> operations. +If there is a strong reason to use expedited grace periods in CPU-hotplug +notifiers, it will be necessary to avoid disabling CPU hotplug. +This would introduce some complexity, so there had better be a <i>very</i> +good reason. + +<p> +The tradeoff between grace-period latency on the one hand and interruptions +of other CPUs on the other hand may need to be re-examined. +The desire is of course for zero grace-period latency as well as zero +interprocessor interrupts undertaken during an expedited grace period +operation. +While this ideal is unlikely to be achievable, it is quite possible that +further improvements can be made. + +<p> +The multiprocessor implementations of RCU use a combining tree that +groups CPUs so as to reduce lock contention and increase cache locality. +However, this combining tree does not spread its memory across NUMA +nodes nor does it align the CPU groups with hardware features such +as sockets or cores. +Such spreading and alignment is currently believed to be unnecessary +because the hotpath read-side primitives do not access the combining +tree, nor does <tt>call_rcu()</tt> in the common case. +If you believe that your architecture needs such spreading and alignment, +then your architecture should also benefit from the +<tt>rcutree.rcu_fanout_leaf</tt> boot parameter, which can be set +to the number of CPUs in a socket, NUMA node, or whatever. +If the number of CPUs is too large, use a fraction of the number of +CPUs. +If the number of CPUs is a large prime number, well, that certainly +is an “interesting” architectural choice! +More flexible arrangements might be considered, but only if +<tt>rcutree.rcu_fanout_leaf</tt> has proven inadequate, and only +if the inadequacy has been demonstrated by a carefully run and +realistic system-level workload. + +<p> +Please note that arrangements that require RCU to remap CPU numbers will +require extremely good demonstration of need and full exploration of +alternatives. + +<p> +There is an embarrassingly large number of flavors of RCU, and this +number has been increasing over time. +Perhaps it will be possible to combine some at some future date. + +<p> +RCU's various kthreads are reasonably recent additions. +It is quite likely that adjustments will be required to more gracefully +handle extreme loads. +It might also be necessary to be able to relate CPU utilization by +RCU's kthreads and softirq handlers to the code that instigated this +CPU utilization. +For example, RCU callback overhead might be charged back to the +originating <tt>call_rcu()</tt> instance, though probably not +in production kernels. + +<h2><a name="Summary">Summary</a></h2> + +<p> +This document has presented more than two decade's worth of RCU +requirements. +Given that the requirements keep changing, this will not be the last +word on this subject, but at least it serves to get an important +subset of the requirements set forth. + +<h2><a name="Acknowledgments">Acknowledgments</a></h2> + +I am grateful to Steven Rostedt, Lai Jiangshan, Ingo Molnar, +Oleg Nesterov, Borislav Petkov, Peter Zijlstra, Boqun Feng, and +Andy Lutomirski for their help in rendering +this article human readable, and to Michelle Rankin for her support +of this effort. +Other contributions are acknowledged in the Linux kernel's git archive. +The cartoon is copyright (c) 2013 by Melissa Broussard, +and is provided +under the terms of the Creative Commons Attribution-Share Alike 3.0 +United States license. + +<h3><a name="Answers to Quick Quizzes"> +Answers to Quick Quizzes</a></h3> + +<a name="qq1answer"></a> +<p><b>Quick Quiz 1</b>: +Wait a minute! +You said that updaters can make useful forward progress concurrently +with readers, but pre-existing readers will block +<tt>synchronize_rcu()</tt>!!! +Just who are you trying to fool??? + + +</p><p><b>Answer</b>: +First, if updaters do not wish to be blocked by readers, they can use +<tt>call_rcu()</tt> or <tt>kfree_rcu()</tt>, which will +be discussed later. +Second, even when using <tt>synchronize_rcu()</tt>, the other +update-side code does run concurrently with readers, whether pre-existing +or not. + + +</p><p><a href="#Quick%20Quiz%201"><b>Back to Quick Quiz 1</b>.</a> + +<a name="qq2answer"></a> +<p><b>Quick Quiz 2</b>: +Why is the <tt>synchronize_rcu()</tt> on line 28 needed? + + +</p><p><b>Answer</b>: +Without that extra grace period, memory reordering could result in +<tt>do_something_dlm()</tt> executing <tt>do_something()</tt> +concurrently with the last bits of <tt>recovery()</tt>. + + +</p><p><a href="#Quick%20Quiz%202"><b>Back to Quick Quiz 2</b>.</a> + +<a name="qq3answer"></a> +<p><b>Quick Quiz 3</b>: +But <tt>rcu_assign_pointer()</tt> does nothing to prevent the +two assignments to <tt>p->a</tt> and <tt>p->b</tt> +from being reordered. +Can't that also cause problems? + + +</p><p><b>Answer</b>: +No, it cannot. +The readers cannot see either of these two fields until +the assignment to <tt>gp</tt>, by which time both fields are +fully initialized. +So reordering the assignments +to <tt>p->a</tt> and <tt>p->b</tt> cannot possibly +cause any problems. + + +</p><p><a href="#Quick%20Quiz%203"><b>Back to Quick Quiz 3</b>.</a> + +<a name="qq4answer"></a> +<p><b>Quick Quiz 4</b>: +Without the <tt>rcu_dereference()</tt> or the +<tt>rcu_access_pointer()</tt>, what destructive optimizations +might the compiler make use of? + + +</p><p><b>Answer</b>: +Let's start with what happens to <tt>do_something_gp()</tt> +if it fails to use <tt>rcu_dereference()</tt>. +It could reuse a value formerly fetched from this same pointer. +It could also fetch the pointer from <tt>gp</tt> in a byte-at-a-time +manner, resulting in <i>load tearing</i>, in turn resulting a bytewise +mash-up of two distince pointer values. +It might even use value-speculation optimizations, where it makes a wrong +guess, but by the time it gets around to checking the value, an update +has changed the pointer to match the wrong guess. +Too bad about any dereferences that returned pre-initialization garbage +in the meantime! + +<p> +For <tt>remove_gp_synchronous()</tt>, as long as all modifications +to <tt>gp</tt> are carried out while holding <tt>gp_lock</tt>, +the above optimizations are harmless. +However, +with <tt>CONFIG_SPARSE_RCU_POINTER=y</tt>, +<tt>sparse</tt> will complain if you +define <tt>gp</tt> with <tt>__rcu</tt> and then +access it without using +either <tt>rcu_access_pointer()</tt> or <tt>rcu_dereference()</tt>. + + +</p><p><a href="#Quick%20Quiz%204"><b>Back to Quick Quiz 4</b>.</a> + +<a name="qq5answer"></a> +<p><b>Quick Quiz 5</b>: +Given that multiple CPUs can start RCU read-side critical sections +at any time without any ordering whatsoever, how can RCU possibly tell whether +or not a given RCU read-side critical section starts before a +given instance of <tt>synchronize_rcu()</tt>? + + +</p><p><b>Answer</b>: +If RCU cannot tell whether or not a given +RCU read-side critical section starts before a +given instance of <tt>synchronize_rcu()</tt>, +then it must assume that the RCU read-side critical section +started first. +In other words, a given instance of <tt>synchronize_rcu()</tt> +can avoid waiting on a given RCU read-side critical section only +if it can prove that <tt>synchronize_rcu()</tt> started first. + + +</p><p><a href="#Quick%20Quiz%205"><b>Back to Quick Quiz 5</b>.</a> + +<a name="qq6answer"></a> +<p><b>Quick Quiz 6</b>: +The first and second guarantees require unbelievably strict ordering! +Are all these memory barriers <i> really</i> required? + + +</p><p><b>Answer</b>: +Yes, they really are required. +To see why the first guarantee is required, consider the following +sequence of events: + +<ol> +<li> CPU 1: <tt>rcu_read_lock()</tt> +<li> CPU 1: <tt>q = rcu_dereference(gp); + /* Very likely to return p. */</tt> +<li> CPU 0: <tt>list_del_rcu(p);</tt> +<li> CPU 0: <tt>synchronize_rcu()</tt> starts. +<li> CPU 1: <tt>do_something_with(q->a); + /* No smp_mb(), so might happen after kfree(). */</tt> +<li> CPU 1: <tt>rcu_read_unlock()</tt> +<li> CPU 0: <tt>synchronize_rcu()</tt> returns. +<li> CPU 0: <tt>kfree(p);</tt> +</ol> + +<p> +Therefore, there absolutely must be a full memory barrier between the +end of the RCU read-side critical section and the end of the +grace period. + +<p> +The sequence of events demonstrating the necessity of the second rule +is roughly similar: + +<ol> +<li> CPU 0: <tt>list_del_rcu(p);</tt> +<li> CPU 0: <tt>synchronize_rcu()</tt> starts. +<li> CPU 1: <tt>rcu_read_lock()</tt> +<li> CPU 1: <tt>q = rcu_dereference(gp); + /* Might return p if no memory barrier. */</tt> +<li> CPU 0: <tt>synchronize_rcu()</tt> returns. +<li> CPU 0: <tt>kfree(p);</tt> +<li> CPU 1: <tt>do_something_with(q->a); /* Boom!!! */</tt> +<li> CPU 1: <tt>rcu_read_unlock()</tt> +</ol> + +<p> +And similarly, without a memory barrier between the beginning of the +grace period and the beginning of the RCU read-side critical section, +CPU 1 might end up accessing the freelist. + +<p> +The “as if” rule of course applies, so that any implementation +that acts as if the appropriate memory barriers were in place is a +correct implementation. +That said, it is much easier to fool yourself into believing that you have +adhered to the as-if rule than it is to actually adhere to it! + + +</p><p><a href="#Quick%20Quiz%206"><b>Back to Quick Quiz 6</b>.</a> + +<a name="qq7answer"></a> +<p><b>Quick Quiz 7</b>: +But how does the upgrade-to-write operation exclude other readers? + + +</p><p><b>Answer</b>: +It doesn't, just like normal RCU updates, which also do not exclude +RCU readers. + + +</p><p><a href="#Quick%20Quiz%207"><b>Back to Quick Quiz 7</b>.</a> + +<a name="qq8answer"></a> +<p><b>Quick Quiz 8</b>: +Can't the compiler also reorder this code? + + +</p><p><b>Answer</b>: +No, the volatile casts in <tt>READ_ONCE()</tt> and +<tt>WRITE_ONCE()</tt> prevent the compiler from reordering in +this particular case. + + +</p><p><a href="#Quick%20Quiz%208"><b>Back to Quick Quiz 8</b>.</a> + +<a name="qq9answer"></a> +<p><b>Quick Quiz 9</b>: +Suppose that synchronize_rcu() did wait until all readers had completed. +Would the updater be able to rely on this? + + +</p><p><b>Answer</b>: +No. +Even if <tt>synchronize_rcu()</tt> were to wait until +all readers had completed, a new reader might start immediately after +<tt>synchronize_rcu()</tt> completed. +Therefore, the code following +<tt>synchronize_rcu()</tt> cannot rely on there being no readers +in any case. + + +</p><p><a href="#Quick%20Quiz%209"><b>Back to Quick Quiz 9</b>.</a> + +<a name="qq10answer"></a> +<p><b>Quick Quiz 10</b>: +How long a sequence of grace periods, each separated by an RCU read-side +critical section, would be required to partition the RCU read-side +critical sections at the beginning and end of the chain? + + +</p><p><b>Answer</b>: +In theory, an infinite number. +In practice, an unknown number that is sensitive to both implementation +details and timing considerations. +Therefore, even in practice, RCU users must abide by the theoretical rather +than the practical answer. + + +</p><p><a href="#Quick%20Quiz%2010"><b>Back to Quick Quiz 10</b>.</a> + +<a name="qq11answer"></a> +<p><b>Quick Quiz 11</b>: +What about sleeping locks? + + +</p><p><b>Answer</b>: +These are forbidden within Linux-kernel RCU read-side critical sections +because it is not legal to place a quiescent state (in this case, +voluntary context switch) within an RCU read-side critical section. +However, sleeping locks may be used within userspace RCU read-side critical +sections, and also within Linux-kernel sleepable RCU +<a href="#Sleepable RCU">(SRCU)</a> +read-side critical sections. +In addition, the -rt patchset turns spinlocks into a sleeping locks so +that the corresponding critical sections can be preempted, which +also means that these sleeplockified spinlocks (but not other sleeping locks!) +may be acquire within -rt-Linux-kernel RCU read-side critical sections. + +<p> +Note that it <i>is</i> legal for a normal RCU read-side critical section +to conditionally acquire a sleeping locks (as in <tt>mutex_trylock()</tt>), +but only as long as it does not loop indefinitely attempting to +conditionally acquire that sleeping locks. +The key point is that things like <tt>mutex_trylock()</tt> +either return with the mutex held, or return an error indication if +the mutex was not immediately available. +Either way, <tt>mutex_trylock()</tt> returns immediately without sleeping. + + +</p><p><a href="#Quick%20Quiz%2011"><b>Back to Quick Quiz 11</b>.</a> + +<a name="qq12answer"></a> +<p><b>Quick Quiz 12</b>: +Why does line 19 use <tt>rcu_access_pointer()</tt>? +After all, <tt>call_rcu()</tt> on line 25 stores into the +structure, which would interact badly with concurrent insertions. +Doesn't this mean that <tt>rcu_dereference()</tt> is required? + + +</p><p><b>Answer</b>: +Presumably the <tt>->gp_lock</tt> acquired on line 18 excludes +any changes, including any insertions that <tt>rcu_dereference()</tt> +would protect against. +Therefore, any insertions will be delayed until after <tt>->gp_lock</tt> +is released on line 25, which in turn means that +<tt>rcu_access_pointer()</tt> suffices. + + +</p><p><a href="#Quick%20Quiz%2012"><b>Back to Quick Quiz 12</b>.</a> + +<a name="qq13answer"></a> +<p><b>Quick Quiz 13</b>: +Earlier it was claimed that <tt>call_rcu()</tt> and +<tt>kfree_rcu()</tt> allowed updaters to avoid being blocked +by readers. +But how can that be correct, given that the invocation of the callback +and the freeing of the memory (respectively) must still wait for +a grace period to elapse? + + +</p><p><b>Answer</b>: +We could define things this way, but keep in mind that this sort of +definition would say that updates in garbage-collected languages +cannot complete until the next time the garbage collector runs, +which does not seem at all reasonable. +The key point is that in most cases, an updater using either +<tt>call_rcu()</tt> or <tt>kfree_rcu()</tt> can proceed to the +next update as soon as it has invoked <tt>call_rcu()</tt> or +<tt>kfree_rcu()</tt>, without having to wait for a subsequent +grace period. + + +</p><p><a href="#Quick%20Quiz%2013"><b>Back to Quick Quiz 13</b>.</a> + +<a name="qq14answer"></a> +<p><b>Quick Quiz 14</b>: +So what happens with <tt>synchronize_rcu()</tt> during +scheduler initialization for <tt>CONFIG_PREEMPT=n</tt> +kernels? + + +</p><p><b>Answer</b>: +In <tt>CONFIG_PREEMPT=n</tt> kernel, <tt>synchronize_rcu()</tt> +maps directly to <tt>synchronize_sched()</tt>. +Therefore, <tt>synchronize_rcu()</tt> works normally throughout +boot in <tt>CONFIG_PREEMPT=n</tt> kernels. +However, your code must also work in <tt>CONFIG_PREEMPT=y</tt> kernels, +so it is still necessary to avoid invoking <tt>synchronize_rcu()</tt> +during scheduler initialization. + + +</p><p><a href="#Quick%20Quiz%2014"><b>Back to Quick Quiz 14</b>.</a> + + +</body></html> diff --git a/Documentation/RCU/Design/Requirements/Requirements.htmlx b/Documentation/RCU/Design/Requirements/Requirements.htmlx new file mode 100644 index 000000000000..3a97ba490c42 --- /dev/null +++ b/Documentation/RCU/Design/Requirements/Requirements.htmlx @@ -0,0 +1,2741 @@ +<!DOCTYPE HTML PUBLIC "-//W3C//DTD HTML 4.01 Transitional//EN" + "http://www.w3.org/TR/html4/loose.dtd"> + <html> + <head><title>A Tour Through RCU's Requirements [LWN.net]</title> + <meta HTTP-EQUIV="Content-Type" CONTENT="text/html; charset=utf-8"> + +<h1>A Tour Through RCU's Requirements</h1> + +<p>Copyright IBM Corporation, 2015</p> +<p>Author: Paul E. McKenney</p> +<p><i>The initial version of this document appeared in the +<a href="https://lwn.net/">LWN</a> articles +<a href="https://lwn.net/Articles/652156/">here</a>, +<a href="https://lwn.net/Articles/652677/">here</a>, and +<a href="https://lwn.net/Articles/653326/">here</a>.</i></p> + +<h2>Introduction</h2> + +<p> +Read-copy update (RCU) is a synchronization mechanism that is often +used as a replacement for reader-writer locking. +RCU is unusual in that updaters do not block readers, +which means that RCU's read-side primitives can be exceedingly fast +and scalable. +In addition, updaters can make useful forward progress concurrently +with readers. +However, all this concurrency between RCU readers and updaters does raise +the question of exactly what RCU readers are doing, which in turn +raises the question of exactly what RCU's requirements are. + +<p> +This document therefore summarizes RCU's requirements, and can be thought +of as an informal, high-level specification for RCU. +It is important to understand that RCU's specification is primarily +empirical in nature; +in fact, I learned about many of these requirements the hard way. +This situation might cause some consternation, however, not only +has this learning process been a lot of fun, but it has also been +a great privilege to work with so many people willing to apply +technologies in interesting new ways. + +<p> +All that aside, here are the categories of currently known RCU requirements: +</p> + +<ol> +<li> <a href="#Fundamental Requirements"> + Fundamental Requirements</a> +<li> <a href="#Fundamental Non-Requirements">Fundamental Non-Requirements</a> +<li> <a href="#Parallelism Facts of Life"> + Parallelism Facts of Life</a> +<li> <a href="#Quality-of-Implementation Requirements"> + Quality-of-Implementation Requirements</a> +<li> <a href="#Linux Kernel Complications"> + Linux Kernel Complications</a> +<li> <a href="#Software-Engineering Requirements"> + Software-Engineering Requirements</a> +<li> <a href="#Other RCU Flavors"> + Other RCU Flavors</a> +<li> <a href="#Possible Future Changes"> + Possible Future Changes</a> +</ol> + +<p> +This is followed by a <a href="#Summary">summary</a>, +which is in turn followed by the inevitable +<a href="#Answers to Quick Quizzes">answers to the quick quizzes</a>. + +<h2><a name="Fundamental Requirements">Fundamental Requirements</a></h2> + +<p> +RCU's fundamental requirements are the closest thing RCU has to hard +mathematical requirements. +These are: + +<ol> +<li> <a href="#Grace-Period Guarantee"> + Grace-Period Guarantee</a> +<li> <a href="#Publish-Subscribe Guarantee"> + Publish-Subscribe Guarantee</a> +<li> <a href="#Memory-Barrier Guarantees"> + Memory-Barrier Guarantees</a> +<li> <a href="#RCU Primitives Guaranteed to Execute Unconditionally"> + RCU Primitives Guaranteed to Execute Unconditionally</a> +<li> <a href="#Guaranteed Read-to-Write Upgrade"> + Guaranteed Read-to-Write Upgrade</a> +</ol> + +<h3><a name="Grace-Period Guarantee">Grace-Period Guarantee</a></h3> + +<p> +RCU's grace-period guarantee is unusual in being premeditated: +Jack Slingwine and I had this guarantee firmly in mind when we started +work on RCU (then called “rclock”) in the early 1990s. +That said, the past two decades of experience with RCU have produced +a much more detailed understanding of this guarantee. + +<p> +RCU's grace-period guarantee allows updaters to wait for the completion +of all pre-existing RCU read-side critical sections. +An RCU read-side critical section +begins with the marker <tt>rcu_read_lock()</tt> and ends with +the marker <tt>rcu_read_unlock()</tt>. +These markers may be nested, and RCU treats a nested set as one +big RCU read-side critical section. +Production-quality implementations of <tt>rcu_read_lock()</tt> and +<tt>rcu_read_unlock()</tt> are extremely lightweight, and in +fact have exactly zero overhead in Linux kernels built for production +use with <tt>CONFIG_PREEMPT=n</tt>. + +<p> +This guarantee allows ordering to be enforced with extremely low +overhead to readers, for example: + +<blockquote> +<pre> + 1 int x, y; + 2 + 3 void thread0(void) + 4 { + 5 rcu_read_lock(); + 6 r1 = READ_ONCE(x); + 7 r2 = READ_ONCE(y); + 8 rcu_read_unlock(); + 9 } +10 +11 void thread1(void) +12 { +13 WRITE_ONCE(x, 1); +14 synchronize_rcu(); +15 WRITE_ONCE(y, 1); +16 } +</pre> +</blockquote> + +<p> +Because the <tt>synchronize_rcu()</tt> on line 14 waits for +all pre-existing readers, any instance of <tt>thread0()</tt> that +loads a value of zero from <tt>x</tt> must complete before +<tt>thread1()</tt> stores to <tt>y</tt>, so that instance must +also load a value of zero from <tt>y</tt>. +Similarly, any instance of <tt>thread0()</tt> that loads a value of +one from <tt>y</tt> must have started after the +<tt>synchronize_rcu()</tt> started, and must therefore also load +a value of one from <tt>x</tt>. +Therefore, the outcome: +<blockquote> +<pre> +(r1 == 0 && r2 == 1) +</pre> +</blockquote> +cannot happen. + +<p>@@QQ@@ +Wait a minute! +You said that updaters can make useful forward progress concurrently +with readers, but pre-existing readers will block +<tt>synchronize_rcu()</tt>!!! +Just who are you trying to fool??? +<p>@@QQA@@ +First, if updaters do not wish to be blocked by readers, they can use +<tt>call_rcu()</tt> or <tt>kfree_rcu()</tt>, which will +be discussed later. +Second, even when using <tt>synchronize_rcu()</tt>, the other +update-side code does run concurrently with readers, whether pre-existing +or not. +<p>@@QQE@@ + +<p> +This scenario resembles one of the first uses of RCU in +<a href="https://en.wikipedia.org/wiki/DYNIX">DYNIX/ptx</a>, +which managed a distributed lock manager's transition into +a state suitable for handling recovery from node failure, +more or less as follows: + +<blockquote> +<pre> + 1 #define STATE_NORMAL 0 + 2 #define STATE_WANT_RECOVERY 1 + 3 #define STATE_RECOVERING 2 + 4 #define STATE_WANT_NORMAL 3 + 5 + 6 int state = STATE_NORMAL; + 7 + 8 void do_something_dlm(void) + 9 { +10 int state_snap; +11 +12 rcu_read_lock(); +13 state_snap = READ_ONCE(state); +14 if (state_snap == STATE_NORMAL) +15 do_something(); +16 else +17 do_something_carefully(); +18 rcu_read_unlock(); +19 } +20 +21 void start_recovery(void) +22 { +23 WRITE_ONCE(state, STATE_WANT_RECOVERY); +24 synchronize_rcu(); +25 WRITE_ONCE(state, STATE_RECOVERING); +26 recovery(); +27 WRITE_ONCE(state, STATE_WANT_NORMAL); +28 synchronize_rcu(); +29 WRITE_ONCE(state, STATE_NORMAL); +30 } +</pre> +</blockquote> + +<p> +The RCU read-side critical section in <tt>do_something_dlm()</tt> +works with the <tt>synchronize_rcu()</tt> in <tt>start_recovery()</tt> +to guarantee that <tt>do_something()</tt> never runs concurrently +with <tt>recovery()</tt>, but with little or no synchronization +overhead in <tt>do_something_dlm()</tt>. + +<p>@@QQ@@ +Why is the <tt>synchronize_rcu()</tt> on line 28 needed? +<p>@@QQA@@ +Without that extra grace period, memory reordering could result in +<tt>do_something_dlm()</tt> executing <tt>do_something()</tt> +concurrently with the last bits of <tt>recovery()</tt>. +<p>@@QQE@@ + +<p> +In order to avoid fatal problems such as deadlocks, +an RCU read-side critical section must not contain calls to +<tt>synchronize_rcu()</tt>. +Similarly, an RCU read-side critical section must not +contain anything that waits, directly or indirectly, on completion of +an invocation of <tt>synchronize_rcu()</tt>. + +<p> +Although RCU's grace-period guarantee is useful in and of itself, with +<a href="https://lwn.net/Articles/573497/">quite a few use cases</a>, +it would be good to be able to use RCU to coordinate read-side +access to linked data structures. +For this, the grace-period guarantee is not sufficient, as can +be seen in function <tt>add_gp_buggy()</tt> below. +We will look at the reader's code later, but in the meantime, just think of +the reader as locklessly picking up the <tt>gp</tt> pointer, +and, if the value loaded is non-<tt>NULL</tt>, locklessly accessing the +<tt>->a</tt> and <tt>->b</tt> fields. + +<blockquote> +<pre> + 1 bool add_gp_buggy(int a, int b) + 2 { + 3 p = kmalloc(sizeof(*p), GFP_KERNEL); + 4 if (!p) + 5 return -ENOMEM; + 6 spin_lock(&gp_lock); + 7 if (rcu_access_pointer(gp)) { + 8 spin_unlock(&gp_lock); + 9 return false; +10 } +11 p->a = a; +12 p->b = a; +13 gp = p; /* ORDERING BUG */ +14 spin_unlock(&gp_lock); +15 return true; +16 } +</pre> +</blockquote> + +<p> +The problem is that both the compiler and weakly ordered CPUs are within +their rights to reorder this code as follows: + +<blockquote> +<pre> + 1 bool add_gp_buggy_optimized(int a, int b) + 2 { + 3 p = kmalloc(sizeof(*p), GFP_KERNEL); + 4 if (!p) + 5 return -ENOMEM; + 6 spin_lock(&gp_lock); + 7 if (rcu_access_pointer(gp)) { + 8 spin_unlock(&gp_lock); + 9 return false; +10 } +<b>11 gp = p; /* ORDERING BUG */ +12 p->a = a; +13 p->b = a;</b> +14 spin_unlock(&gp_lock); +15 return true; +16 } +</pre> +</blockquote> + +<p> +If an RCU reader fetches <tt>gp</tt> just after +<tt>add_gp_buggy_optimized</tt> executes line 11, +it will see garbage in the <tt>->a</tt> and <tt>->b</tt> +fields. +And this is but one of many ways in which compiler and hardware optimizations +could cause trouble. +Therefore, we clearly need some way to prevent the compiler and the CPU from +reordering in this manner, which brings us to the publish-subscribe +guarantee discussed in the next section. + +<h3><a name="Publish-Subscribe Guarantee">Publish/Subscribe Guarantee</a></h3> + +<p> +RCU's publish-subscribe guarantee allows data to be inserted +into a linked data structure without disrupting RCU readers. +The updater uses <tt>rcu_assign_pointer()</tt> to insert the +new data, and readers use <tt>rcu_dereference()</tt> to +access data, whether new or old. +The following shows an example of insertion: + +<blockquote> +<pre> + 1 bool add_gp(int a, int b) + 2 { + 3 p = kmalloc(sizeof(*p), GFP_KERNEL); + 4 if (!p) + 5 return -ENOMEM; + 6 spin_lock(&gp_lock); + 7 if (rcu_access_pointer(gp)) { + 8 spin_unlock(&gp_lock); + 9 return false; +10 } +11 p->a = a; +12 p->b = a; +13 rcu_assign_pointer(gp, p); +14 spin_unlock(&gp_lock); +15 return true; +16 } +</pre> +</blockquote> + +<p> +The <tt>rcu_assign_pointer()</tt> on line 13 is conceptually +equivalent to a simple assignment statement, but also guarantees +that its assignment will +happen after the two assignments in lines 11 and 12, +similar to the C11 <tt>memory_order_release</tt> store operation. +It also prevents any number of “interesting” compiler +optimizations, for example, the use of <tt>gp</tt> as a scratch +location immediately preceding the assignment. + +<p>@@QQ@@ +But <tt>rcu_assign_pointer()</tt> does nothing to prevent the +two assignments to <tt>p->a</tt> and <tt>p->b</tt> +from being reordered. +Can't that also cause problems? +<p>@@QQA@@ +No, it cannot. +The readers cannot see either of these two fields until +the assignment to <tt>gp</tt>, by which time both fields are +fully initialized. +So reordering the assignments +to <tt>p->a</tt> and <tt>p->b</tt> cannot possibly +cause any problems. +<p>@@QQE@@ + +<p> +It is tempting to assume that the reader need not do anything special +to control its accesses to the RCU-protected data, +as shown in <tt>do_something_gp_buggy()</tt> below: + +<blockquote> +<pre> + 1 bool do_something_gp_buggy(void) + 2 { + 3 rcu_read_lock(); + 4 p = gp; /* OPTIMIZATIONS GALORE!!! */ + 5 if (p) { + 6 do_something(p->a, p->b); + 7 rcu_read_unlock(); + 8 return true; + 9 } +10 rcu_read_unlock(); +11 return false; +12 } +</pre> +</blockquote> + +<p> +However, this temptation must be resisted because there are a +surprisingly large number of ways that the compiler +(to say nothing of +<a href="https://h71000.www7.hp.com/wizard/wiz_2637.html">DEC Alpha CPUs</a>) +can trip this code up. +For but one example, if the compiler were short of registers, it +might choose to refetch from <tt>gp</tt> rather than keeping +a separate copy in <tt>p</tt> as follows: + +<blockquote> +<pre> + 1 bool do_something_gp_buggy_optimized(void) + 2 { + 3 rcu_read_lock(); + 4 if (gp) { /* OPTIMIZATIONS GALORE!!! */ +<b> 5 do_something(gp->a, gp->b);</b> + 6 rcu_read_unlock(); + 7 return true; + 8 } + 9 rcu_read_unlock(); +10 return false; +11 } +</pre> +</blockquote> + +<p> +If this function ran concurrently with a series of updates that +replaced the current structure with a new one, +the fetches of <tt>gp->a</tt> +and <tt>gp->b</tt> might well come from two different structures, +which could cause serious confusion. +To prevent this (and much else besides), <tt>do_something_gp()</tt> uses +<tt>rcu_dereference()</tt> to fetch from <tt>gp</tt>: + +<blockquote> +<pre> + 1 bool do_something_gp(void) + 2 { + 3 rcu_read_lock(); + 4 p = rcu_dereference(gp); + 5 if (p) { + 6 do_something(p->a, p->b); + 7 rcu_read_unlock(); + 8 return true; + 9 } +10 rcu_read_unlock(); +11 return false; +12 } +</pre> +</blockquote> + +<p> +The <tt>rcu_dereference()</tt> uses volatile casts and (for DEC Alpha) +memory barriers in the Linux kernel. +Should a +<a href="http://www.rdrop.com/users/paulmck/RCU/consume.2015.07.13a.pdf">high-quality implementation of C11 <tt>memory_order_consume</tt> [PDF]</a> +ever appear, then <tt>rcu_dereference()</tt> could be implemented +as a <tt>memory_order_consume</tt> load. +Regardless of the exact implementation, a pointer fetched by +<tt>rcu_dereference()</tt> may not be used outside of the +outermost RCU read-side critical section containing that +<tt>rcu_dereference()</tt>, unless protection of +the corresponding data element has been passed from RCU to some +other synchronization mechanism, most commonly locking or +<a href="https://www.kernel.org/doc/Documentation/RCU/rcuref.txt">reference counting</a>. + +<p> +In short, updaters use <tt>rcu_assign_pointer()</tt> and readers +use <tt>rcu_dereference()</tt>, and these two RCU API elements +work together to ensure that readers have a consistent view of +newly added data elements. + +<p> +Of course, it is also necessary to remove elements from RCU-protected +data structures, for example, using the following process: + +<ol> +<li> Remove the data element from the enclosing structure. +<li> Wait for all pre-existing RCU read-side critical sections + to complete (because only pre-existing readers can possibly have + a reference to the newly removed data element). +<li> At this point, only the updater has a reference to the + newly removed data element, so it can safely reclaim + the data element, for example, by passing it to <tt>kfree()</tt>. +</ol> + +This process is implemented by <tt>remove_gp_synchronous()</tt>: + +<blockquote> +<pre> + 1 bool remove_gp_synchronous(void) + 2 { + 3 struct foo *p; + 4 + 5 spin_lock(&gp_lock); + 6 p = rcu_access_pointer(gp); + 7 if (!p) { + 8 spin_unlock(&gp_lock); + 9 return false; +10 } +11 rcu_assign_pointer(gp, NULL); +12 spin_unlock(&gp_lock); +13 synchronize_rcu(); +14 kfree(p); +15 return true; +16 } +</pre> +</blockquote> + +<p> +This function is straightforward, with line 13 waiting for a grace +period before line 14 frees the old data element. +This waiting ensures that readers will reach line 7 of +<tt>do_something_gp()</tt> before the data element referenced by +<tt>p</tt> is freed. +The <tt>rcu_access_pointer()</tt> on line 6 is similar to +<tt>rcu_dereference()</tt>, except that: + +<ol> +<li> The value returned by <tt>rcu_access_pointer()</tt> + cannot be dereferenced. + If you want to access the value pointed to as well as + the pointer itself, use <tt>rcu_dereference()</tt> + instead of <tt>rcu_access_pointer()</tt>. +<li> The call to <tt>rcu_access_pointer()</tt> need not be + protected. + In contrast, <tt>rcu_dereference()</tt> must either be + within an RCU read-side critical section or in a code + segment where the pointer cannot change, for example, in + code protected by the corresponding update-side lock. +</ol> + +<p>@@QQ@@ +Without the <tt>rcu_dereference()</tt> or the +<tt>rcu_access_pointer()</tt>, what destructive optimizations +might the compiler make use of? +<p>@@QQA@@ +Let's start with what happens to <tt>do_something_gp()</tt> +if it fails to use <tt>rcu_dereference()</tt>. +It could reuse a value formerly fetched from this same pointer. +It could also fetch the pointer from <tt>gp</tt> in a byte-at-a-time +manner, resulting in <i>load tearing</i>, in turn resulting a bytewise +mash-up of two distince pointer values. +It might even use value-speculation optimizations, where it makes a wrong +guess, but by the time it gets around to checking the value, an update +has changed the pointer to match the wrong guess. +Too bad about any dereferences that returned pre-initialization garbage +in the meantime! + +<p> +For <tt>remove_gp_synchronous()</tt>, as long as all modifications +to <tt>gp</tt> are carried out while holding <tt>gp_lock</tt>, +the above optimizations are harmless. +However, +with <tt>CONFIG_SPARSE_RCU_POINTER=y</tt>, +<tt>sparse</tt> will complain if you +define <tt>gp</tt> with <tt>__rcu</tt> and then +access it without using +either <tt>rcu_access_pointer()</tt> or <tt>rcu_dereference()</tt>. +<p>@@QQE@@ + +<p> +In short, RCU's publish-subscribe guarantee is provided by the combination +of <tt>rcu_assign_pointer()</tt> and <tt>rcu_dereference()</tt>. +This guarantee allows data elements to be safely added to RCU-protected +linked data structures without disrupting RCU readers. +This guarantee can be used in combination with the grace-period +guarantee to also allow data elements to be removed from RCU-protected +linked data structures, again without disrupting RCU readers. + +<p> +This guarantee was only partially premeditated. +DYNIX/ptx used an explicit memory barrier for publication, but had nothing +resembling <tt>rcu_dereference()</tt> for subscription, nor did it +have anything resembling the <tt>smp_read_barrier_depends()</tt> +that was later subsumed into <tt>rcu_dereference()</tt>. +The need for these operations made itself known quite suddenly at a +late-1990s meeting with the DEC Alpha architects, back in the days when +DEC was still a free-standing company. +It took the Alpha architects a good hour to convince me that any sort +of barrier would ever be needed, and it then took me a good <i>two</i> hours +to convince them that their documentation did not make this point clear. +More recent work with the C and C++ standards committees have provided +much education on tricks and traps from the compiler. +In short, compilers were much less tricky in the early 1990s, but in +2015, don't even think about omitting <tt>rcu_dereference()</tt>! + +<h3><a name="Memory-Barrier Guarantees">Memory-Barrier Guarantees</a></h3> + +<p> +The previous section's simple linked-data-structure scenario clearly +demonstrates the need for RCU's stringent memory-ordering guarantees on +systems with more than one CPU: + +<ol> +<li> Each CPU that has an RCU read-side critical section that + begins before <tt>synchronize_rcu()</tt> starts is + guaranteed to execute a full memory barrier between the time + that the RCU read-side critical section ends and the time that + <tt>synchronize_rcu()</tt> returns. + Without this guarantee, a pre-existing RCU read-side critical section + might hold a reference to the newly removed <tt>struct foo</tt> + after the <tt>kfree()</tt> on line 14 of + <tt>remove_gp_synchronous()</tt>. +<li> Each CPU that has an RCU read-side critical section that ends + after <tt>synchronize_rcu()</tt> returns is guaranteed + to execute a full memory barrier between the time that + <tt>synchronize_rcu()</tt> begins and the time that the RCU + read-side critical section begins. + Without this guarantee, a later RCU read-side critical section + running after the <tt>kfree()</tt> on line 14 of + <tt>remove_gp_synchronous()</tt> might + later run <tt>do_something_gp()</tt> and find the + newly deleted <tt>struct foo</tt>. +<li> If the task invoking <tt>synchronize_rcu()</tt> remains + on a given CPU, then that CPU is guaranteed to execute a full + memory barrier sometime during the execution of + <tt>synchronize_rcu()</tt>. + This guarantee ensures that the <tt>kfree()</tt> on + line 14 of <tt>remove_gp_synchronous()</tt> really does + execute after the removal on line 11. +<li> If the task invoking <tt>synchronize_rcu()</tt> migrates + among a group of CPUs during that invocation, then each of the + CPUs in that group is guaranteed to execute a full memory barrier + sometime during the execution of <tt>synchronize_rcu()</tt>. + This guarantee also ensures that the <tt>kfree()</tt> on + line 14 of <tt>remove_gp_synchronous()</tt> really does + execute after the removal on + line 11, but also in the case where the thread executing the + <tt>synchronize_rcu()</tt> migrates in the meantime. +</ol> + +<p>@@QQ@@ +Given that multiple CPUs can start RCU read-side critical sections +at any time without any ordering whatsoever, how can RCU possibly tell whether +or not a given RCU read-side critical section starts before a +given instance of <tt>synchronize_rcu()</tt>? +<p>@@QQA@@ +If RCU cannot tell whether or not a given +RCU read-side critical section starts before a +given instance of <tt>synchronize_rcu()</tt>, +then it must assume that the RCU read-side critical section +started first. +In other words, a given instance of <tt>synchronize_rcu()</tt> +can avoid waiting on a given RCU read-side critical section only +if it can prove that <tt>synchronize_rcu()</tt> started first. +<p>@@QQE@@ + +<p>@@QQ@@ +The first and second guarantees require unbelievably strict ordering! +Are all these memory barriers <i> really</i> required? +<p>@@QQA@@ +Yes, they really are required. +To see why the first guarantee is required, consider the following +sequence of events: + +<ol> +<li> CPU 1: <tt>rcu_read_lock()</tt> +<li> CPU 1: <tt>q = rcu_dereference(gp); + /* Very likely to return p. */</tt> +<li> CPU 0: <tt>list_del_rcu(p);</tt> +<li> CPU 0: <tt>synchronize_rcu()</tt> starts. +<li> CPU 1: <tt>do_something_with(q->a); + /* No smp_mb(), so might happen after kfree(). */</tt> +<li> CPU 1: <tt>rcu_read_unlock()</tt> +<li> CPU 0: <tt>synchronize_rcu()</tt> returns. +<li> CPU 0: <tt>kfree(p);</tt> +</ol> + +<p> +Therefore, there absolutely must be a full memory barrier between the +end of the RCU read-side critical section and the end of the +grace period. + +<p> +The sequence of events demonstrating the necessity of the second rule +is roughly similar: + +<ol> +<li> CPU 0: <tt>list_del_rcu(p);</tt> +<li> CPU 0: <tt>synchronize_rcu()</tt> starts. +<li> CPU 1: <tt>rcu_read_lock()</tt> +<li> CPU 1: <tt>q = rcu_dereference(gp); + /* Might return p if no memory barrier. */</tt> +<li> CPU 0: <tt>synchronize_rcu()</tt> returns. +<li> CPU 0: <tt>kfree(p);</tt> +<li> CPU 1: <tt>do_something_with(q->a); /* Boom!!! */</tt> +<li> CPU 1: <tt>rcu_read_unlock()</tt> +</ol> + +<p> +And similarly, without a memory barrier between the beginning of the +grace period and the beginning of the RCU read-side critical section, +CPU 1 might end up accessing the freelist. + +<p> +The “as if” rule of course applies, so that any implementation +that acts as if the appropriate memory barriers were in place is a +correct implementation. +That said, it is much easier to fool yourself into believing that you have +adhered to the as-if rule than it is to actually adhere to it! +<p>@@QQE@@ + +<p> +Note that these memory-barrier requirements do not replace the fundamental +RCU requirement that a grace period wait for all pre-existing readers. +On the contrary, the memory barriers called out in this section must operate in +such a way as to <i>enforce</i> this fundamental requirement. +Of course, different implementations enforce this requirement in different +ways, but enforce it they must. + +<h3><a name="RCU Primitives Guaranteed to Execute Unconditionally">RCU Primitives Guaranteed to Execute Unconditionally</a></h3> + +<p> +The common-case RCU primitives are unconditional. +They are invoked, they do their job, and they return, with no possibility +of error, and no need to retry. +This is a key RCU design philosophy. + +<p> +However, this philosophy is pragmatic rather than pigheaded. +If someone comes up with a good justification for a particular conditional +RCU primitive, it might well be implemented and added. +After all, this guarantee was reverse-engineered, not premeditated. +The unconditional nature of the RCU primitives was initially an +accident of implementation, and later experience with synchronization +primitives with conditional primitives caused me to elevate this +accident to a guarantee. +Therefore, the justification for adding a conditional primitive to +RCU would need to be based on detailed and compelling use cases. + +<h3><a name="Guaranteed Read-to-Write Upgrade">Guaranteed Read-to-Write Upgrade</a></h3> + +<p> +As far as RCU is concerned, it is always possible to carry out an +update within an RCU read-side critical section. +For example, that RCU read-side critical section might search for +a given data element, and then might acquire the update-side +spinlock in order to update that element, all while remaining +in that RCU read-side critical section. +Of course, it is necessary to exit the RCU read-side critical section +before invoking <tt>synchronize_rcu()</tt>, however, this +inconvenience can be avoided through use of the +<tt>call_rcu()</tt> and <tt>kfree_rcu()</tt> API members +described later in this document. + +<p>@@QQ@@ +But how does the upgrade-to-write operation exclude other readers? +<p>@@QQA@@ +It doesn't, just like normal RCU updates, which also do not exclude +RCU readers. +<p>@@QQE@@ + +<p> +This guarantee allows lookup code to be shared between read-side +and update-side code, and was premeditated, appearing in the earliest +DYNIX/ptx RCU documentation. + +<h2><a name="Fundamental Non-Requirements">Fundamental Non-Requirements</a></h2> + +<p> +RCU provides extremely lightweight readers, and its read-side guarantees, +though quite useful, are correspondingly lightweight. +It is therefore all too easy to assume that RCU is guaranteeing more +than it really is. +Of course, the list of things that RCU does not guarantee is infinitely +long, however, the following sections list a few non-guarantees that +have caused confusion. +Except where otherwise noted, these non-guarantees were premeditated. + +<ol> +<li> <a href="#Readers Impose Minimal Ordering"> + Readers Impose Minimal Ordering</a> +<li> <a href="#Readers Do Not Exclude Updaters"> + Readers Do Not Exclude Updaters</a> +<li> <a href="#Updaters Only Wait For Old Readers"> + Updaters Only Wait For Old Readers</a> +<li> <a href="#Grace Periods Don't Partition Read-Side Critical Sections"> + Grace Periods Don't Partition Read-Side Critical Sections</a> +<li> <a href="#Read-Side Critical Sections Don't Partition Grace Periods"> + Read-Side Critical Sections Don't Partition Grace Periods</a> +<li> <a href="#Disabling Preemption Does Not Block Grace Periods"> + Disabling Preemption Does Not Block Grace Periods</a> +</ol> + +<h3><a name="Readers Impose Minimal Ordering">Readers Impose Minimal Ordering</a></h3> + +<p> +Reader-side markers such as <tt>rcu_read_lock()</tt> and +<tt>rcu_read_unlock()</tt> provide absolutely no ordering guarantees +except through their interaction with the grace-period APIs such as +<tt>synchronize_rcu()</tt>. +To see this, consider the following pair of threads: + +<blockquote> +<pre> + 1 void thread0(void) + 2 { + 3 rcu_read_lock(); + 4 WRITE_ONCE(x, 1); + 5 rcu_read_unlock(); + 6 rcu_read_lock(); + 7 WRITE_ONCE(y, 1); + 8 rcu_read_unlock(); + 9 } +10 +11 void thread1(void) +12 { +13 rcu_read_lock(); +14 r1 = READ_ONCE(y); +15 rcu_read_unlock(); +16 rcu_read_lock(); +17 r2 = READ_ONCE(x); +18 rcu_read_unlock(); +19 } +</pre> +</blockquote> + +<p> +After <tt>thread0()</tt> and <tt>thread1()</tt> execute +concurrently, it is quite possible to have + +<blockquote> +<pre> +(r1 == 1 && r2 == 0) +</pre> +</blockquote> + +(that is, <tt>y</tt> appears to have been assigned before <tt>x</tt>), +which would not be possible if <tt>rcu_read_lock()</tt> and +<tt>rcu_read_unlock()</tt> had much in the way of ordering +properties. +But they do not, so the CPU is within its rights +to do significant reordering. +This is by design: Any significant ordering constraints would slow down +these fast-path APIs. + +<p>@@QQ@@ +Can't the compiler also reorder this code? +<p>@@QQA@@ +No, the volatile casts in <tt>READ_ONCE()</tt> and +<tt>WRITE_ONCE()</tt> prevent the compiler from reordering in +this particular case. +<p>@@QQE@@ + +<h3><a name="Readers Do Not Exclude Updaters">Readers Do Not Exclude Updaters</a></h3> + +<p> +Neither <tt>rcu_read_lock()</tt> nor <tt>rcu_read_unlock()</tt> +exclude updates. +All they do is to prevent grace periods from ending. +The following example illustrates this: + +<blockquote> +<pre> + 1 void thread0(void) + 2 { + 3 rcu_read_lock(); + 4 r1 = READ_ONCE(y); + 5 if (r1) { + 6 do_something_with_nonzero_x(); + 7 r2 = READ_ONCE(x); + 8 WARN_ON(!r2); /* BUG!!! */ + 9 } +10 rcu_read_unlock(); +11 } +12 +13 void thread1(void) +14 { +15 spin_lock(&my_lock); +16 WRITE_ONCE(x, 1); +17 WRITE_ONCE(y, 1); +18 spin_unlock(&my_lock); +19 } +</pre> +</blockquote> + +<p> +If the <tt>thread0()</tt> function's <tt>rcu_read_lock()</tt> +excluded the <tt>thread1()</tt> function's update, +the <tt>WARN_ON()</tt> could never fire. +But the fact is that <tt>rcu_read_lock()</tt> does not exclude +much of anything aside from subsequent grace periods, of which +<tt>thread1()</tt> has none, so the +<tt>WARN_ON()</tt> can and does fire. + +<h3><a name="Updaters Only Wait For Old Readers">Updaters Only Wait For Old Readers</a></h3> + +<p> +It might be tempting to assume that after <tt>synchronize_rcu()</tt> +completes, there are no readers executing. +This temptation must be avoided because +new readers can start immediately after <tt>synchronize_rcu()</tt> +starts, and <tt>synchronize_rcu()</tt> is under no +obligation to wait for these new readers. + +<p>@@QQ@@ +Suppose that synchronize_rcu() did wait until all readers had completed. +Would the updater be able to rely on this? +<p>@@QQA@@ +No. +Even if <tt>synchronize_rcu()</tt> were to wait until +all readers had completed, a new reader might start immediately after +<tt>synchronize_rcu()</tt> completed. +Therefore, the code following +<tt>synchronize_rcu()</tt> cannot rely on there being no readers +in any case. +<p>@@QQE@@ + +<h3><a name="Grace Periods Don't Partition Read-Side Critical Sections"> +Grace Periods Don't Partition Read-Side Critical Sections</a></h3> + +<p> +It is tempting to assume that if any part of one RCU read-side critical +section precedes a given grace period, and if any part of another RCU +read-side critical section follows that same grace period, then all of +the first RCU read-side critical section must precede all of the second. +However, this just isn't the case: A single grace period does not +partition the set of RCU read-side critical sections. +An example of this situation can be illustrated as follows, where +<tt>x</tt>, <tt>y</tt>, and <tt>z</tt> are initially all zero: + +<blockquote> +<pre> + 1 void thread0(void) + 2 { + 3 rcu_read_lock(); + 4 WRITE_ONCE(a, 1); + 5 WRITE_ONCE(b, 1); + 6 rcu_read_unlock(); + 7 } + 8 + 9 void thread1(void) +10 { +11 r1 = READ_ONCE(a); +12 synchronize_rcu(); +13 WRITE_ONCE(c, 1); +14 } +15 +16 void thread2(void) +17 { +18 rcu_read_lock(); +19 r2 = READ_ONCE(b); +20 r3 = READ_ONCE(c); +21 rcu_read_unlock(); +22 } +</pre> +</blockquote> + +<p> +It turns out that the outcome: + +<blockquote> +<pre> +(r1 == 1 && r2 == 0 && r3 == 1) +</pre> +</blockquote> + +is entirely possible. +The following figure show how this can happen, with each circled +<tt>QS</tt> indicating the point at which RCU recorded a +<i>quiescent state</i> for each thread, that is, a state in which +RCU knows that the thread cannot be in the midst of an RCU read-side +critical section that started before the current grace period: + +<p><img src="GPpartitionReaders1.svg" alt="GPpartitionReaders1.svg" width="60%"></p> + +<p> +If it is necessary to partition RCU read-side critical sections in this +manner, it is necessary to use two grace periods, where the first +grace period is known to end before the second grace period starts: + +<blockquote> +<pre> + 1 void thread0(void) + 2 { + 3 rcu_read_lock(); + 4 WRITE_ONCE(a, 1); + 5 WRITE_ONCE(b, 1); + 6 rcu_read_unlock(); + 7 } + 8 + 9 void thread1(void) +10 { +11 r1 = READ_ONCE(a); +12 synchronize_rcu(); +13 WRITE_ONCE(c, 1); +14 } +15 +16 void thread2(void) +17 { +18 r2 = READ_ONCE(c); +19 synchronize_rcu(); +20 WRITE_ONCE(d, 1); +21 } +22 +23 void thread3(void) +24 { +25 rcu_read_lock(); +26 r3 = READ_ONCE(b); +27 r4 = READ_ONCE(d); +28 rcu_read_unlock(); +29 } +</pre> +</blockquote> + +<p> +Here, if <tt>(r1 == 1)</tt>, then +<tt>thread0()</tt>'s write to <tt>b</tt> must happen +before the end of <tt>thread1()</tt>'s grace period. +If in addition <tt>(r4 == 1)</tt>, then +<tt>thread3()</tt>'s read from <tt>b</tt> must happen +after the beginning of <tt>thread2()</tt>'s grace period. +If it is also the case that <tt>(r2 == 1)</tt>, then the +end of <tt>thread1()</tt>'s grace period must precede the +beginning of <tt>thread2()</tt>'s grace period. +This mean that the two RCU read-side critical sections cannot overlap, +guaranteeing that <tt>(r3 == 1)</tt>. +As a result, the outcome: + +<blockquote> +<pre> +(r1 == 1 && r2 == 1 && r3 == 0 && r4 == 1) +</pre> +</blockquote> + +cannot happen. + +<p> +This non-requirement was also non-premeditated, but became apparent +when studying RCU's interaction with memory ordering. + +<h3><a name="Read-Side Critical Sections Don't Partition Grace Periods"> +Read-Side Critical Sections Don't Partition Grace Periods</a></h3> + +<p> +It is also tempting to assume that if an RCU read-side critical section +happens between a pair of grace periods, then those grace periods cannot +overlap. +However, this temptation leads nowhere good, as can be illustrated by +the following, with all variables initially zero: + +<blockquote> +<pre> + 1 void thread0(void) + 2 { + 3 rcu_read_lock(); + 4 WRITE_ONCE(a, 1); + 5 WRITE_ONCE(b, 1); + 6 rcu_read_unlock(); + 7 } + 8 + 9 void thread1(void) +10 { +11 r1 = READ_ONCE(a); +12 synchronize_rcu(); +13 WRITE_ONCE(c, 1); +14 } +15 +16 void thread2(void) +17 { +18 rcu_read_lock(); +19 WRITE_ONCE(d, 1); +20 r2 = READ_ONCE(c); +21 rcu_read_unlock(); +22 } +23 +24 void thread3(void) +25 { +26 r3 = READ_ONCE(d); +27 synchronize_rcu(); +28 WRITE_ONCE(e, 1); +29 } +30 +31 void thread4(void) +32 { +33 rcu_read_lock(); +34 r4 = READ_ONCE(b); +35 r5 = READ_ONCE(e); +36 rcu_read_unlock(); +37 } +</pre> +</blockquote> + +<p> +In this case, the outcome: + +<blockquote> +<pre> +(r1 == 1 && r2 == 1 && r3 == 1 && r4 == 0 && r5 == 1) +</pre> +</blockquote> + +is entirely possible, as illustrated below: + +<p><img src="ReadersPartitionGP1.svg" alt="ReadersPartitionGP1.svg" width="100%"></p> + +<p> +Again, an RCU read-side critical section can overlap almost all of a +given grace period, just so long as it does not overlap the entire +grace period. +As a result, an RCU read-side critical section cannot partition a pair +of RCU grace periods. + +<p>@@QQ@@ +How long a sequence of grace periods, each separated by an RCU read-side +critical section, would be required to partition the RCU read-side +critical sections at the beginning and end of the chain? +<p>@@QQA@@ +In theory, an infinite number. +In practice, an unknown number that is sensitive to both implementation +details and timing considerations. +Therefore, even in practice, RCU users must abide by the theoretical rather +than the practical answer. +<p>@@QQE@@ + +<h3><a name="Disabling Preemption Does Not Block Grace Periods"> +Disabling Preemption Does Not Block Grace Periods</a></h3> + +<p> +There was a time when disabling preemption on any given CPU would block +subsequent grace periods. +However, this was an accident of implementation and is not a requirement. +And in the current Linux-kernel implementation, disabling preemption +on a given CPU in fact does not block grace periods, as Oleg Nesterov +<a href="https://lkml.kernel.org/g/20150614193825.GA19582@redhat.com">demonstrated</a>. + +<p> +If you need a preempt-disable region to block grace periods, you need to add +<tt>rcu_read_lock()</tt> and <tt>rcu_read_unlock()</tt>, for example +as follows: + +<blockquote> +<pre> + 1 preempt_disable(); + 2 rcu_read_lock(); + 3 do_something(); + 4 rcu_read_unlock(); + 5 preempt_enable(); + 6 + 7 /* Spinlocks implicitly disable preemption. */ + 8 spin_lock(&mylock); + 9 rcu_read_lock(); +10 do_something(); +11 rcu_read_unlock(); +12 spin_unlock(&mylock); +</pre> +</blockquote> + +<p> +In theory, you could enter the RCU read-side critical section first, +but it is more efficient to keep the entire RCU read-side critical +section contained in the preempt-disable region as shown above. +Of course, RCU read-side critical sections that extend outside of +preempt-disable regions will work correctly, but such critical sections +can be preempted, which forces <tt>rcu_read_unlock()</tt> to do +more work. +And no, this is <i>not</i> an invitation to enclose all of your RCU +read-side critical sections within preempt-disable regions, because +doing so would degrade real-time response. + +<p> +This non-requirement appeared with preemptible RCU. +If you need a grace period that waits on non-preemptible code regions, use +<a href="#Sched Flavor">RCU-sched</a>. + +<h2><a name="Parallelism Facts of Life">Parallelism Facts of Life</a></h2> + +<p> +These parallelism facts of life are by no means specific to RCU, but +the RCU implementation must abide by them. +They therefore bear repeating: + +<ol> +<li> Any CPU or task may be delayed at any time, + and any attempts to avoid these delays by disabling + preemption, interrupts, or whatever are completely futile. + This is most obvious in preemptible user-level + environments and in virtualized environments (where + a given guest OS's VCPUs can be preempted at any time by + the underlying hypervisor), but can also happen in bare-metal + environments due to ECC errors, NMIs, and other hardware + events. + Although a delay of more than about 20 seconds can result + in splats, the RCU implementation is obligated to use + algorithms that can tolerate extremely long delays, but where + “extremely long” is not long enough to allow + wrap-around when incrementing a 64-bit counter. +<li> Both the compiler and the CPU can reorder memory accesses. + Where it matters, RCU must use compiler directives and + memory-barrier instructions to preserve ordering. +<li> Conflicting writes to memory locations in any given cache line + will result in expensive cache misses. + Greater numbers of concurrent writes and more-frequent + concurrent writes will result in more dramatic slowdowns. + RCU is therefore obligated to use algorithms that have + sufficient locality to avoid significant performance and + scalability problems. +<li> As a rough rule of thumb, only one CPU's worth of processing + may be carried out under the protection of any given exclusive + lock. + RCU must therefore use scalable locking designs. +<li> Counters are finite, especially on 32-bit systems. + RCU's use of counters must therefore tolerate counter wrap, + or be designed such that counter wrap would take way more + time than a single system is likely to run. + An uptime of ten years is quite possible, a runtime + of a century much less so. + As an example of the latter, RCU's dyntick-idle nesting counter + allows 54 bits for interrupt nesting level (this counter + is 64 bits even on a 32-bit system). + Overflowing this counter requires 2<sup>54</sup> + half-interrupts on a given CPU without that CPU ever going idle. + If a half-interrupt happened every microsecond, it would take + 570 years of runtime to overflow this counter, which is currently + believed to be an acceptably long time. +<li> Linux systems can have thousands of CPUs running a single + Linux kernel in a single shared-memory environment. + RCU must therefore pay close attention to high-end scalability. +</ol> + +<p> +This last parallelism fact of life means that RCU must pay special +attention to the preceding facts of life. +The idea that Linux might scale to systems with thousands of CPUs would +have been met with some skepticism in the 1990s, but these requirements +would have otherwise have been unsurprising, even in the early 1990s. + +<h2><a name="Quality-of-Implementation Requirements">Quality-of-Implementation Requirements</a></h2> + +<p> +These sections list quality-of-implementation requirements. +Although an RCU implementation that ignores these requirements could +still be used, it would likely be subject to limitations that would +make it inappropriate for industrial-strength production use. +Classes of quality-of-implementation requirements are as follows: + +<ol> +<li> <a href="#Specialization">Specialization</a> +<li> <a href="#Performance and Scalability">Performance and Scalability</a> +<li> <a href="#Composability">Composability</a> +<li> <a href="#Corner Cases">Corner Cases</a> +</ol> + +<p> +These classes is covered in the following sections. + +<h3><a name="Specialization">Specialization</a></h3> + +<p> +RCU is and always has been intended primarily for read-mostly situations, as +illustrated by the following figure. +This means that RCU's read-side primitives are optimized, often at the +expense of its update-side primitives. + +<p><img src="RCUApplicability.svg" alt="RCUApplicability.svg" width="70%"></p> + +<p> +This focus on read-mostly situations means that RCU must interoperate +with other synchronization primitives. +For example, the <tt>add_gp()</tt> and <tt>remove_gp_synchronous()</tt> +examples discussed earlier use RCU to protect readers and locking to +coordinate updaters. +However, the need extends much farther, requiring that a variety of +synchronization primitives be legal within RCU read-side critical sections, +including spinlocks, sequence locks, atomic operations, reference +counters, and memory barriers. + +<p>@@QQ@@ +What about sleeping locks? +<p>@@QQA@@ +These are forbidden within Linux-kernel RCU read-side critical sections +because it is not legal to place a quiescent state (in this case, +voluntary context switch) within an RCU read-side critical section. +However, sleeping locks may be used within userspace RCU read-side critical +sections, and also within Linux-kernel sleepable RCU +<a href="#Sleepable RCU">(SRCU)</a> +read-side critical sections. +In addition, the -rt patchset turns spinlocks into a sleeping locks so +that the corresponding critical sections can be preempted, which +also means that these sleeplockified spinlocks (but not other sleeping locks!) +may be acquire within -rt-Linux-kernel RCU read-side critical sections. + +<p> +Note that it <i>is</i> legal for a normal RCU read-side critical section +to conditionally acquire a sleeping locks (as in <tt>mutex_trylock()</tt>), +but only as long as it does not loop indefinitely attempting to +conditionally acquire that sleeping locks. +The key point is that things like <tt>mutex_trylock()</tt> +either return with the mutex held, or return an error indication if +the mutex was not immediately available. +Either way, <tt>mutex_trylock()</tt> returns immediately without sleeping. +<p>@@QQE@@ + +<p> +It often comes as a surprise that many algorithms do not require a +consistent view of data, but many can function in that mode, +with network routing being the poster child. +Internet routing algorithms take significant time to propagate +updates, so that by the time an update arrives at a given system, +that system has been sending network traffic the wrong way for +a considerable length of time. +Having a few threads continue to send traffic the wrong way for a +few more milliseconds is clearly not a problem: In the worst case, +TCP retransmissions will eventually get the data where it needs to go. +In general, when tracking the state of the universe outside of the +computer, some level of inconsistency must be tolerated due to +speed-of-light delays if nothing else. + +<p> +Furthermore, uncertainty about external state is inherent in many cases. +For example, a pair of veternarians might use heartbeat to determine +whether or not a given cat was alive. +But how long should they wait after the last heartbeat to decide that +the cat is in fact dead? +Waiting less than 400 milliseconds makes no sense because this would +mean that a relaxed cat would be considered to cycle between death +and life more than 100 times per minute. +Moreover, just as with human beings, a cat's heart might stop for +some period of time, so the exact wait period is a judgment call. +One of our pair of veternarians might wait 30 seconds before pronouncing +the cat dead, while the other might insist on waiting a full minute. +The two veternarians would then disagree on the state of the cat during +the final 30 seconds of the minute following the last heartbeat, as +fancifully illustrated below: + +<p><img src="2013-08-is-it-dead.png" alt="2013-08-is-it-dead.png" width="431"></p> + +<p> +Interestingly enough, this same situation applies to hardware. +When push comes to shove, how do we tell whether or not some +external server has failed? +We send messages to it periodically, and declare it failed if we +don't receive a response within a given period of time. +Policy decisions can usually tolerate short +periods of inconsistency. +The policy was decided some time ago, and is only now being put into +effect, so a few milliseconds of delay is normally inconsequential. + +<p> +However, there are algorithms that absolutely must see consistent data. +For example, the translation between a user-level SystemV semaphore +ID to the corresponding in-kernel data structure is protected by RCU, +but it is absolutely forbidden to update a semaphore that has just been +removed. +In the Linux kernel, this need for consistency is accommodated by acquiring +spinlocks located in the in-kernel data structure from within +the RCU read-side critical section, and this is indicated by the +green box in the figure above. +Many other techniques may be used, and are in fact used within the +Linux kernel. + +<p> +In short, RCU is not required to maintain consistency, and other +mechanisms may be used in concert with RCU when consistency is required. +RCU's specialization allows it to do its job extremely well, and its +ability to interoperate with other synchronization mechanisms allows +the right mix of synchronization tools to be used for a given job. + +<h3><a name="Performance and Scalability">Performance and Scalability</a></h3> + +<p> +Energy efficiency is a critical component of performance today, +and Linux-kernel RCU implementations must therefore avoid unnecessarily +awakening idle CPUs. +I cannot claim that this requirement was premeditated. +In fact, I learned of it during a telephone conversation in which I +was given “frank and open” feedback on the importance +of energy efficiency in battery-powered systems and on specific +energy-efficiency shortcomings of the Linux-kernel RCU implementation. +In my experience, the battery-powered embedded community will consider +any unnecessary wakeups to be extremely unfriendly acts. +So much so that mere Linux-kernel-mailing-list posts are +insufficient to vent their ire. + +<p> +Memory consumption is not particularly important for in most +situations, and has become decreasingly +so as memory sizes have expanded and memory +costs have plummeted. +However, as I learned from Matt Mackall's +<a href="http://elinux.org/Linux_Tiny-FAQ">bloatwatch</a> +efforts, memory footprint is critically important on single-CPU systems with +non-preemptible (<tt>CONFIG_PREEMPT=n</tt>) kernels, and thus +<a href="https://lkml.kernel.org/g/20090113221724.GA15307@linux.vnet.ibm.com">tiny RCU</a> +was born. +Josh Triplett has since taken over the small-memory banner with his +<a href="https://tiny.wiki.kernel.org/">Linux kernel tinification</a> +project, which resulted in +<a href="#Sleepable RCU">SRCU</a> +becoming optional for those kernels not needing it. + +<p> +The remaining performance requirements are, for the most part, +unsurprising. +For example, in keeping with RCU's read-side specialization, +<tt>rcu_dereference()</tt> should have negligible overhead (for +example, suppression of a few minor compiler optimizations). +Similarly, in non-preemptible environments, <tt>rcu_read_lock()</tt> and +<tt>rcu_read_unlock()</tt> should have exactly zero overhead. + +<p> +In preemptible environments, in the case where the RCU read-side +critical section was not preempted (as will be the case for the +highest-priority real-time process), <tt>rcu_read_lock()</tt> and +<tt>rcu_read_unlock()</tt> should have minimal overhead. +In particular, they should not contain atomic read-modify-write +operations, memory-barrier instructions, preemption disabling, +interrupt disabling, or backwards branches. +However, in the case where the RCU read-side critical section was preempted, +<tt>rcu_read_unlock()</tt> may acquire spinlocks and disable interrupts. +This is why it is better to nest an RCU read-side critical section +within a preempt-disable region than vice versa, at least in cases +where that critical section is short enough to avoid unduly degrading +real-time latencies. + +<p> +The <tt>synchronize_rcu()</tt> grace-period-wait primitive is +optimized for throughput. +It may therefore incur several milliseconds of latency in addition to +the duration of the longest RCU read-side critical section. +On the other hand, multiple concurrent invocations of +<tt>synchronize_rcu()</tt> are required to use batching optimizations +so that they can be satisfied by a single underlying grace-period-wait +operation. +For example, in the Linux kernel, it is not unusual for a single +grace-period-wait operation to serve more than +<a href="https://www.usenix.org/conference/2004-usenix-annual-technical-conference/making-rcu-safe-deep-sub-millisecond-response">1,000 separate invocations</a> +of <tt>synchronize_rcu()</tt>, thus amortizing the per-invocation +overhead down to nearly zero. +However, the grace-period optimization is also required to avoid +measurable degradation of real-time scheduling and interrupt latencies. + +<p> +In some cases, the multi-millisecond <tt>synchronize_rcu()</tt> +latencies are unacceptable. +In these cases, <tt>synchronize_rcu_expedited()</tt> may be used +instead, reducing the grace-period latency down to a few tens of +microseconds on small systems, at least in cases where the RCU read-side +critical sections are short. +There are currently no special latency requirements for +<tt>synchronize_rcu_expedited()</tt> on large systems, but, +consistent with the empirical nature of the RCU specification, +that is subject to change. +However, there most definitely are scalability requirements: +A storm of <tt>synchronize_rcu_expedited()</tt> invocations on 4096 +CPUs should at least make reasonable forward progress. +In return for its shorter latencies, <tt>synchronize_rcu_expedited()</tt> +is permitted to impose modest degradation of real-time latency +on non-idle online CPUs. +That said, it will likely be necessary to take further steps to reduce this +degradation, hopefully to roughly that of a scheduling-clock interrupt. + +<p> +There are a number of situations where even +<tt>synchronize_rcu_expedited()</tt>'s reduced grace-period +latency is unacceptable. +In these situations, the asynchronous <tt>call_rcu()</tt> can be +used in place of <tt>synchronize_rcu()</tt> as follows: + +<blockquote> +<pre> + 1 struct foo { + 2 int a; + 3 int b; + 4 struct rcu_head rh; + 5 }; + 6 + 7 static void remove_gp_cb(struct rcu_head *rhp) + 8 { + 9 struct foo *p = container_of(rhp, struct foo, rh); +10 +11 kfree(p); +12 } +13 +14 bool remove_gp_asynchronous(void) +15 { +16 struct foo *p; +17 +18 spin_lock(&gp_lock); +19 p = rcu_dereference(gp); +20 if (!p) { +21 spin_unlock(&gp_lock); +22 return false; +23 } +24 rcu_assign_pointer(gp, NULL); +25 call_rcu(&p->rh, remove_gp_cb); +26 spin_unlock(&gp_lock); +27 return true; +28 } +</pre> +</blockquote> + +<p> +A definition of <tt>struct foo</tt> is finally needed, and appears +on lines 1-5. +The function <tt>remove_gp_cb()</tt> is passed to <tt>call_rcu()</tt> +on line 25, and will be invoked after the end of a subsequent +grace period. +This gets the same effect as <tt>remove_gp_synchronous()</tt>, +but without forcing the updater to wait for a grace period to elapse. +The <tt>call_rcu()</tt> function may be used in a number of +situations where neither <tt>synchronize_rcu()</tt> nor +<tt>synchronize_rcu_expedited()</tt> would be legal, +including within preempt-disable code, <tt>local_bh_disable()</tt> code, +interrupt-disable code, and interrupt handlers. +However, even <tt>call_rcu()</tt> is illegal within NMI handlers. +The callback function (<tt>remove_gp_cb()</tt> in this case) will be +executed within softirq (software interrupt) environment within the +Linux kernel, +either within a real softirq handler or under the protection +of <tt>local_bh_disable()</tt>. +In both the Linux kernel and in userspace, it is bad practice to +write an RCU callback function that takes too long. +Long-running operations should be relegated to separate threads or +(in the Linux kernel) workqueues. + +<p>@@QQ@@ +Why does line 19 use <tt>rcu_access_pointer()</tt>? +After all, <tt>call_rcu()</tt> on line 25 stores into the +structure, which would interact badly with concurrent insertions. +Doesn't this mean that <tt>rcu_dereference()</tt> is required? +<p>@@QQA@@ +Presumably the <tt>->gp_lock</tt> acquired on line 18 excludes +any changes, including any insertions that <tt>rcu_dereference()</tt> +would protect against. +Therefore, any insertions will be delayed until after <tt>->gp_lock</tt> +is released on line 25, which in turn means that +<tt>rcu_access_pointer()</tt> suffices. +<p>@@QQE@@ + +<p> +However, all that <tt>remove_gp_cb()</tt> is doing is +invoking <tt>kfree()</tt> on the data element. +This is a common idiom, and is supported by <tt>kfree_rcu()</tt>, +which allows “fire and forget” operation as shown below: + +<blockquote> +<pre> + 1 struct foo { + 2 int a; + 3 int b; + 4 struct rcu_head rh; + 5 }; + 6 + 7 bool remove_gp_faf(void) + 8 { + 9 struct foo *p; +10 +11 spin_lock(&gp_lock); +12 p = rcu_dereference(gp); +13 if (!p) { +14 spin_unlock(&gp_lock); +15 return false; +16 } +17 rcu_assign_pointer(gp, NULL); +18 kfree_rcu(p, rh); +19 spin_unlock(&gp_lock); +20 return true; +21 } +</pre> +</blockquote> + +<p> +Note that <tt>remove_gp_faf()</tt> simply invokes +<tt>kfree_rcu()</tt> and proceeds, without any need to pay any +further attention to the subsequent grace period and <tt>kfree()</tt>. +It is permissible to invoke <tt>kfree_rcu()</tt> from the same +environments as for <tt>call_rcu()</tt>. +Interestingly enough, DYNIX/ptx had the equivalents of +<tt>call_rcu()</tt> and <tt>kfree_rcu()</tt>, but not +<tt>synchronize_rcu()</tt>. +This was due to the fact that RCU was not heavily used within DYNIX/ptx, +so the very few places that needed something like +<tt>synchronize_rcu()</tt> simply open-coded it. + +<p>@@QQ@@ +Earlier it was claimed that <tt>call_rcu()</tt> and +<tt>kfree_rcu()</tt> allowed updaters to avoid being blocked +by readers. +But how can that be correct, given that the invocation of the callback +and the freeing of the memory (respectively) must still wait for +a grace period to elapse? +<p>@@QQA@@ +We could define things this way, but keep in mind that this sort of +definition would say that updates in garbage-collected languages +cannot complete until the next time the garbage collector runs, +which does not seem at all reasonable. +The key point is that in most cases, an updater using either +<tt>call_rcu()</tt> or <tt>kfree_rcu()</tt> can proceed to the +next update as soon as it has invoked <tt>call_rcu()</tt> or +<tt>kfree_rcu()</tt>, without having to wait for a subsequent +grace period. +<p>@@QQE@@ + +<p> +But what if the updater must wait for the completion of code to be +executed after the end of the grace period, but has other tasks +that can be carried out in the meantime? +The polling-style <tt>get_state_synchronize_rcu()</tt> and +<tt>cond_synchronize_rcu()</tt> functions may be used for this +purpose, as shown below: + +<blockquote> +<pre> + 1 bool remove_gp_poll(void) + 2 { + 3 struct foo *p; + 4 unsigned long s; + 5 + 6 spin_lock(&gp_lock); + 7 p = rcu_access_pointer(gp); + 8 if (!p) { + 9 spin_unlock(&gp_lock); +10 return false; +11 } +12 rcu_assign_pointer(gp, NULL); +13 spin_unlock(&gp_lock); +14 s = get_state_synchronize_rcu(); +15 do_something_while_waiting(); +16 cond_synchronize_rcu(s); +17 kfree(p); +18 return true; +19 } +</pre> +</blockquote> + +<p> +On line 14, <tt>get_state_synchronize_rcu()</tt> obtains a +“cookie” from RCU, +then line 15 carries out other tasks, +and finally, line 16 returns immediately if a grace period has +elapsed in the meantime, but otherwise waits as required. +The need for <tt>get_state_synchronize_rcu</tt> and +<tt>cond_synchronize_rcu()</tt> has appeared quite recently, +so it is too early to tell whether they will stand the test of time. + +<p> +RCU thus provides a range of tools to allow updaters to strike the +required tradeoff between latency, flexibility and CPU overhead. + +<h3><a name="Composability">Composability</a></h3> + +<p> +Composability has received much attention in recent years, perhaps in part +due to the collision of multicore hardware with object-oriented techniques +designed in single-threaded environments for single-threaded use. +And in theory, RCU read-side critical sections may be composed, and in +fact may be nested arbitrarily deeply. +In practice, as with all real-world implementations of composable +constructs, there are limitations. + +<p> +Implementations of RCU for which <tt>rcu_read_lock()</tt> +and <tt>rcu_read_unlock()</tt> generate no code, such as +Linux-kernel RCU when <tt>CONFIG_PREEMPT=n</tt>, can be +nested arbitrarily deeply. +After all, there is no overhead. +Except that if all these instances of <tt>rcu_read_lock()</tt> +and <tt>rcu_read_unlock()</tt> are visible to the compiler, +compilation will eventually fail due to exhausting memory, +mass storage, or user patience, whichever comes first. +If the nesting is not visible to the compiler, as is the case with +mutually recursive functions each in its own translation unit, +stack overflow will result. +If the nesting takes the form of loops, either the control variable +will overflow or (in the Linux kernel) you will get an RCU CPU stall warning. +Nevertheless, this class of RCU implementations is one +of the most composable constructs in existence. + +<p> +RCU implementations that explicitly track nesting depth +are limited by the nesting-depth counter. +For example, the Linux kernel's preemptible RCU limits nesting to +<tt>INT_MAX</tt>. +This should suffice for almost all practical purposes. +That said, a consecutive pair of RCU read-side critical sections +between which there is an operation that waits for a grace period +cannot be enclosed in another RCU read-side critical section. +This is because it is not legal to wait for a grace period within +an RCU read-side critical section: To do so would result either +in deadlock or +in RCU implicitly splitting the enclosing RCU read-side critical +section, neither of which is conducive to a long-lived and prosperous +kernel. + +<p> +It is worth noting that RCU is not alone in limiting composability. +For example, many transactional-memory implementations prohibit +composing a pair of transactions separated by an irrevocable +operation (for example, a network receive operation). +For another example, lock-based critical sections can be composed +surprisingly freely, but only if deadlock is avoided. + +<p> +In short, although RCU read-side critical sections are highly composable, +care is required in some situations, just as is the case for any other +composable synchronization mechanism. + +<h3><a name="Corner Cases">Corner Cases</a></h3> + +<p> +A given RCU workload might have an endless and intense stream of +RCU read-side critical sections, perhaps even so intense that there +was never a point in time during which there was not at least one +RCU read-side critical section in flight. +RCU cannot allow this situation to block grace periods: As long as +all the RCU read-side critical sections are finite, grace periods +must also be finite. + +<p> +That said, preemptible RCU implementations could potentially result +in RCU read-side critical sections being preempted for long durations, +which has the effect of creating a long-duration RCU read-side +critical section. +This situation can arise only in heavily loaded systems, but systems using +real-time priorities are of course more vulnerable. +Therefore, RCU priority boosting is provided to help deal with this +case. +That said, the exact requirements on RCU priority boosting will likely +evolve as more experience accumulates. + +<p> +Other workloads might have very high update rates. +Although one can argue that such workloads should instead use +something other than RCU, the fact remains that RCU must +handle such workloads gracefully. +This requirement is another factor driving batching of grace periods, +but it is also the driving force behind the checks for large numbers +of queued RCU callbacks in the <tt>call_rcu()</tt> code path. +Finally, high update rates should not delay RCU read-side critical +sections, although some read-side delays can occur when using +<tt>synchronize_rcu_expedited()</tt>, courtesy of this function's use +of <tt>try_stop_cpus()</tt>. +(In the future, <tt>synchronize_rcu_expedited()</tt> will be +converted to use lighter-weight inter-processor interrupts (IPIs), +but this will still disturb readers, though to a much smaller degree.) + +<p> +Although all three of these corner cases were understood in the early +1990s, a simple user-level test consisting of <tt>close(open(path))</tt> +in a tight loop +in the early 2000s suddenly provided a much deeper appreciation of the +high-update-rate corner case. +This test also motivated addition of some RCU code to react to high update +rates, for example, if a given CPU finds itself with more than 10,000 +RCU callbacks queued, it will cause RCU to take evasive action by +more aggressively starting grace periods and more aggressively forcing +completion of grace-period processing. +This evasive action causes the grace period to complete more quickly, +but at the cost of restricting RCU's batching optimizations, thus +increasing the CPU overhead incurred by that grace period. + +<h2><a name="Software-Engineering Requirements"> +Software-Engineering Requirements</a></h2> + +<p> +Between Murphy's Law and “To err is human”, it is necessary to +guard against mishaps and misuse: + +<ol> +<li> It is all too easy to forget to use <tt>rcu_read_lock()</tt> + everywhere that it is needed, so kernels built with + <tt>CONFIG_PROVE_RCU=y</tt> will spat if + <tt>rcu_dereference()</tt> is used outside of an + RCU read-side critical section. + Update-side code can use <tt>rcu_dereference_protected()</tt>, + which takes a + <a href="https://lwn.net/Articles/371986/">lockdep expression</a> + to indicate what is providing the protection. + If the indicated protection is not provided, a lockdep splat + is emitted. + + <p> + Code shared between readers and updaters can use + <tt>rcu_dereference_check()</tt>, which also takes a + lockdep expression, and emits a lockdep splat if neither + <tt>rcu_read_lock()</tt> nor the indicated protection + is in place. + In addition, <tt>rcu_dereference_raw()</tt> is used in those + (hopefully rare) cases where the required protection cannot + be easily described. + Finally, <tt>rcu_read_lock_held()</tt> is provided to + allow a function to verify that it has been invoked within + an RCU read-side critical section. + I was made aware of this set of requirements shortly after Thomas + Gleixner audited a number of RCU uses. +<li> A given function might wish to check for RCU-related preconditions + upon entry, before using any other RCU API. + The <tt>rcu_lockdep_assert()</tt> does this job, + asserting the expression in kernels having lockdep enabled + and doing nothing otherwise. +<li> It is also easy to forget to use <tt>rcu_assign_pointer()</tt> + and <tt>rcu_dereference()</tt>, perhaps (incorrectly) + substituting a simple assignment. + To catch this sort of error, a given RCU-protected pointer may be + tagged with <tt>__rcu</tt>, after which running sparse + with <tt>CONFIG_SPARSE_RCU_POINTER=y</tt> will complain + about simple-assignment accesses to that pointer. + Arnd Bergmann made me aware of this requirement, and also + supplied the needed + <a href="https://lwn.net/Articles/376011/">patch series</a>. +<li> Kernels built with <tt>CONFIG_DEBUG_OBJECTS_RCU_HEAD=y</tt> + will splat if a data element is passed to <tt>call_rcu()</tt> + twice in a row, without a grace period in between. + (This error is similar to a double free.) + The corresponding <tt>rcu_head</tt> structures that are + dynamically allocated are automatically tracked, but + <tt>rcu_head</tt> structures allocated on the stack + must be initialized with <tt>init_rcu_head_on_stack()</tt> + and cleaned up with <tt>destroy_rcu_head_on_stack()</tt>. + Similarly, statically allocated non-stack <tt>rcu_head</tt> + structures must be initialized with <tt>init_rcu_head()</tt> + and cleaned up with <tt>destroy_rcu_head()</tt>. + Mathieu Desnoyers made me aware of this requirement, and also + supplied the needed + <a href="https://lkml.kernel.org/g/20100319013024.GA28456@Krystal">patch</a>. +<li> An infinite loop in an RCU read-side critical section will + eventually trigger an RCU CPU stall warning splat, with + the duration of “eventually” being controlled by the + <tt>RCU_CPU_STALL_TIMEOUT</tt> <tt>Kconfig</tt> option, or, + alternatively, by the + <tt>rcupdate.rcu_cpu_stall_timeout</tt> boot/sysfs + parameter. + However, RCU is not obligated to produce this splat + unless there is a grace period waiting on that particular + RCU read-side critical section. + <p> + Some extreme workloads might intentionally delay + RCU grace periods, and systems running those workloads can + be booted with <tt>rcupdate.rcu_cpu_stall_suppress</tt> + to suppress the splats. + This kernel parameter may also be set via <tt>sysfs</tt>. + Furthermore, RCU CPU stall warnings are counter-productive + during sysrq dumps and during panics. + RCU therefore supplies the <tt>rcu_sysrq_start()</tt> and + <tt>rcu_sysrq_end()</tt> API members to be called before + and after long sysrq dumps. + RCU also supplies the <tt>rcu_panic()</tt> notifier that is + automatically invoked at the beginning of a panic to suppress + further RCU CPU stall warnings. + + <p> + This requirement made itself known in the early 1990s, pretty + much the first time that it was necessary to debug a CPU stall. + That said, the initial implementation in DYNIX/ptx was quite + generic in comparison with that of Linux. +<li> Although it would be very good to detect pointers leaking out + of RCU read-side critical sections, there is currently no + good way of doing this. + One complication is the need to distinguish between pointers + leaking and pointers that have been handed off from RCU to + some other synchronization mechanism, for example, reference + counting. +<li> In kernels built with <tt>CONFIG_RCU_TRACE=y</tt>, RCU-related + information is provided via both debugfs and event tracing. +<li> Open-coded use of <tt>rcu_assign_pointer()</tt> and + <tt>rcu_dereference()</tt> to create typical linked + data structures can be surprisingly error-prone. + Therefore, RCU-protected + <a href="https://lwn.net/Articles/609973/#RCU List APIs">linked lists</a> + and, more recently, RCU-protected + <a href="https://lwn.net/Articles/612100/">hash tables</a> + are available. + Many other special-purpose RCU-protected data structures are + available in the Linux kernel and the userspace RCU library. +<li> Some linked structures are created at compile time, but still + require <tt>__rcu</tt> checking. + The <tt>RCU_POINTER_INITIALIZER()</tt> macro serves this + purpose. +<li> It is not necessary to use <tt>rcu_assign_pointer()</tt> + when creating linked structures that are to be published via + a single external pointer. + The <tt>RCU_INIT_POINTER()</tt> macro is provided for + this task and also for assigning <tt>NULL</tt> pointers + at runtime. +</ol> + +<p> +This not a hard-and-fast list: RCU's diagnostic capabilities will +continue to be guided by the number and type of usage bugs found +in real-world RCU usage. + +<h2><a name="Linux Kernel Complications">Linux Kernel Complications</a></h2> + +<p> +The Linux kernel provides an interesting environment for all kinds of +software, including RCU. +Some of the relevant points of interest are as follows: + +<ol> +<li> <a href="#Configuration">Configuration</a>. +<li> <a href="#Firmware Interface">Firmware Interface</a>. +<li> <a href="#Early Boot">Early Boot</a>. +<li> <a href="#Interrupts and NMIs"> + Interrupts and non-maskable interrupts (NMIs)</a>. +<li> <a href="#Loadable Modules">Loadable Modules</a>. +<li> <a href="#Hotplug CPU">Hotplug CPU</a>. +<li> <a href="#Scheduler and RCU">Scheduler and RCU</a>. +<li> <a href="#Tracing and RCU">Tracing and RCU</a>. +<li> <a href="#Energy Efficiency">Energy Efficiency</a>. +<li> <a href="#Memory Efficiency">Memory Efficiency</a>. +<li> <a href="#Performance, Scalability, Response Time, and Reliability"> + Performance, Scalability, Response Time, and Reliability</a>. +</ol> + +<p> +This list is probably incomplete, but it does give a feel for the +most notable Linux-kernel complications. +Each of the following sections covers one of the above topics. + +<h3><a name="Configuration">Configuration</a></h3> + +<p> +RCU's goal is automatic configuration, so that almost nobody +needs to worry about RCU's <tt>Kconfig</tt> options. +And for almost all users, RCU does in fact work well +“out of the box.” + +<p> +However, there are specialized use cases that are handled by +kernel boot parameters and <tt>Kconfig</tt> options. +Unfortunately, the <tt>Kconfig</tt> system will explicitly ask users +about new <tt>Kconfig</tt> options, which requires almost all of them +be hidden behind a <tt>CONFIG_RCU_EXPERT</tt> <tt>Kconfig</tt> option. + +<p> +This all should be quite obvious, but the fact remains that +Linus Torvalds recently had to +<a href="https://lkml.kernel.org/g/CA+55aFy4wcCwaL4okTs8wXhGZ5h-ibecy_Meg9C4MNQrUnwMcg@mail.gmail.com">remind</a> +me of this requirement. + +<h3><a name="Firmware Interface">Firmware Interface</a></h3> + +<p> +In many cases, kernel obtains information about the system from the +firmware, and sometimes things are lost in translation. +Or the translation is accurate, but the original message is bogus. + +<p> +For example, some systems' firmware overreports the number of CPUs, +sometimes by a large factor. +If RCU naively believed the firmware, as it used to do, +it would create too many per-CPU kthreads. +Although the resulting system will still run correctly, the extra +kthreads needlessly consume memory and can cause confusion +when they show up in <tt>ps</tt> listings. + +<p> +RCU must therefore wait for a given CPU to actually come online before +it can allow itself to believe that the CPU actually exists. +The resulting “ghost CPUs” (which are never going to +come online) cause a number of +<a href="https://paulmck.livejournal.com/37494.html">interesting complications</a>. + +<h3><a name="Early Boot">Early Boot</a></h3> + +<p> +The Linux kernel's boot sequence is an interesting process, +and RCU is used early, even before <tt>rcu_init()</tt> +is invoked. +In fact, a number of RCU's primitives can be used as soon as the +initial task's <tt>task_struct</tt> is available and the +boot CPU's per-CPU variables are set up. +The read-side primitives (<tt>rcu_read_lock()</tt>, +<tt>rcu_read_unlock()</tt>, <tt>rcu_dereference()</tt>, +and <tt>rcu_access_pointer()</tt>) will operate normally very early on, +as will <tt>rcu_assign_pointer()</tt>. + +<p> +Although <tt>call_rcu()</tt> may be invoked at any +time during boot, callbacks are not guaranteed to be invoked until after +the scheduler is fully up and running. +This delay in callback invocation is due to the fact that RCU does not +invoke callbacks until it is fully initialized, and this full initialization +cannot occur until after the scheduler has initialized itself to the +point where RCU can spawn and run its kthreads. +In theory, it would be possible to invoke callbacks earlier, +however, this is not a panacea because there would be severe restrictions +on what operations those callbacks could invoke. + +<p> +Perhaps surprisingly, <tt>synchronize_rcu()</tt>, +<a href="#Bottom-Half Flavor"><tt>synchronize_rcu_bh()</tt></a> +(<a href="#Bottom-Half Flavor">discussed below</a>), +and +<a href="#Sched Flavor"><tt>synchronize_sched()</tt></a> +will all operate normally +during very early boot, the reason being that there is only one CPU +and preemption is disabled. +This means that the call <tt>synchronize_rcu()</tt> (or friends) +itself is a quiescent +state and thus a grace period, so the early-boot implementation can +be a no-op. + +<p> +Both <tt>synchronize_rcu_bh()</tt> and <tt>synchronize_sched()</tt> +continue to operate normally through the remainder of boot, courtesy +of the fact that preemption is disabled across their RCU read-side +critical sections and also courtesy of the fact that there is still +only one CPU. +However, once the scheduler starts initializing, preemption is enabled. +There is still only a single CPU, but the fact that preemption is enabled +means that the no-op implementation of <tt>synchronize_rcu()</tt> no +longer works in <tt>CONFIG_PREEMPT=y</tt> kernels. +Therefore, as soon as the scheduler starts initializing, the early-boot +fastpath is disabled. +This means that <tt>synchronize_rcu()</tt> switches to its runtime +mode of operation where it posts callbacks, which in turn means that +any call to <tt>synchronize_rcu()</tt> will block until the corresponding +callback is invoked. +Unfortunately, the callback cannot be invoked until RCU's runtime +grace-period machinery is up and running, which cannot happen until +the scheduler has initialized itself sufficiently to allow RCU's +kthreads to be spawned. +Therefore, invoking <tt>synchronize_rcu()</tt> during scheduler +initialization can result in deadlock. + +<p>@@QQ@@ +So what happens with <tt>synchronize_rcu()</tt> during +scheduler initialization for <tt>CONFIG_PREEMPT=n</tt> +kernels? +<p>@@QQA@@ +In <tt>CONFIG_PREEMPT=n</tt> kernel, <tt>synchronize_rcu()</tt> +maps directly to <tt>synchronize_sched()</tt>. +Therefore, <tt>synchronize_rcu()</tt> works normally throughout +boot in <tt>CONFIG_PREEMPT=n</tt> kernels. +However, your code must also work in <tt>CONFIG_PREEMPT=y</tt> kernels, +so it is still necessary to avoid invoking <tt>synchronize_rcu()</tt> +during scheduler initialization. +<p>@@QQE@@ + +<p> +I learned of these boot-time requirements as a result of a series of +system hangs. + +<h3><a name="Interrupts and NMIs">Interrupts and NMIs</a></h3> + +<p> +The Linux kernel has interrupts, and RCU read-side critical sections are +legal within interrupt handlers and within interrupt-disabled regions +of code, as are invocations of <tt>call_rcu()</tt>. + +<p> +Some Linux-kernel architectures can enter an interrupt handler from +non-idle process context, and then just never leave it, instead stealthily +transitioning back to process context. +This trick is sometimes used to invoke system calls from inside the kernel. +These “half-interrupts” mean that RCU has to be very careful +about how it counts interrupt nesting levels. +I learned of this requirement the hard way during a rewrite +of RCU's dyntick-idle code. + +<p> +The Linux kernel has non-maskable interrupts (NMIs), and +RCU read-side critical sections are legal within NMI handlers. +Thankfully, RCU update-side primitives, including +<tt>call_rcu()</tt>, are prohibited within NMI handlers. + +<p> +The name notwithstanding, some Linux-kernel architectures +can have nested NMIs, which RCU must handle correctly. +Andy Lutomirski +<a href="https://lkml.kernel.org/g/CALCETrXLq1y7e_dKFPgou-FKHB6Pu-r8+t-6Ds+8=va7anBWDA@mail.gmail.com">surprised me</a> +with this requirement; +he also kindly surprised me with +<a href="https://lkml.kernel.org/g/CALCETrXSY9JpW3uE6H8WYk81sg56qasA2aqmjMPsq5dOtzso=g@mail.gmail.com">an algorithm</a> +that meets this requirement. + +<h3><a name="Loadable Modules">Loadable Modules</a></h3> + +<p> +The Linux kernel has loadable modules, and these modules can +also be unloaded. +After a given module has been unloaded, any attempt to call +one of its functions results in a segmentation fault. +The module-unload functions must therefore cancel any +delayed calls to loadable-module functions, for example, +any outstanding <tt>mod_timer()</tt> must be dealt with +via <tt>del_timer_sync()</tt> or similar. + +<p> +Unfortunately, there is no way to cancel an RCU callback; +once you invoke <tt>call_rcu()</tt>, the callback function is +going to eventually be invoked, unless the system goes down first. +Because it is normally considered socially irresponsible to crash the system +in response to a module unload request, we need some other way +to deal with in-flight RCU callbacks. + +<p> +RCU therefore provides +<tt><a href="https://lwn.net/Articles/217484/">rcu_barrier()</a></tt>, +which waits until all in-flight RCU callbacks have been invoked. +If a module uses <tt>call_rcu()</tt>, its exit function should therefore +prevent any future invocation of <tt>call_rcu()</tt>, then invoke +<tt>rcu_barrier()</tt>. +In theory, the underlying module-unload code could invoke +<tt>rcu_barrier()</tt> unconditionally, but in practice this would +incur unacceptable latencies. + +<p> +Nikita Danilov noted this requirement for an analogous filesystem-unmount +situation, and Dipankar Sarma incorporated <tt>rcu_barrier()</tt> into RCU. +The need for <tt>rcu_barrier()</tt> for module unloading became +apparent later. + +<h3><a name="Hotplug CPU">Hotplug CPU</a></h3> + +<p> +The Linux kernel supports CPU hotplug, which means that CPUs +can come and go. +It is of course illegal to use any RCU API member from an offline CPU. +This requirement was present from day one in DYNIX/ptx, but +on the other hand, the Linux kernel's CPU-hotplug implementation +is “interesting.” + +<p> +The Linux-kernel CPU-hotplug implementation has notifiers that +are used to allow the various kernel subsystems (including RCU) +to respond appropriately to a given CPU-hotplug operation. +Most RCU operations may be invoked from CPU-hotplug notifiers, +including even normal synchronous grace-period operations +such as <tt>synchronize_rcu()</tt>. +However, expedited grace-period operations such as +<tt>synchronize_rcu_expedited()</tt> are not supported, +due to the fact that current implementations block CPU-hotplug +operations, which could result in deadlock. + +<p> +In addition, all-callback-wait operations such as +<tt>rcu_barrier()</tt> are also not supported, due to the +fact that there are phases of CPU-hotplug operations where +the outgoing CPU's callbacks will not be invoked until after +the CPU-hotplug operation ends, which could also result in deadlock. + +<h3><a name="Scheduler and RCU">Scheduler and RCU</a></h3> + +<p> +RCU depends on the scheduler, and the scheduler uses RCU to +protect some of its data structures. +This means the scheduler is forbidden from acquiring +the runqueue locks and the priority-inheritance locks +in the middle of an outermost RCU read-side critical section unless either +(1) it releases them before exiting that same +RCU read-side critical section, or +(2) interrupts are disabled across +that entire RCU read-side critical section. +This same prohibition also applies (recursively!) to any lock that is acquired +while holding any lock to which this prohibition applies. +Adhering to this rule prevents preemptible RCU from invoking +<tt>rcu_read_unlock_special()</tt> while either runqueue or +priority-inheritance locks are held, thus avoiding deadlock. + +<p> +Prior to v4.4, it was only necessary to disable preemption across +RCU read-side critical sections that acquired scheduler locks. +In v4.4, expedited grace periods started using IPIs, and these +IPIs could force a <tt>rcu_read_unlock()</tt> to take the slowpath. +Therefore, this expedited-grace-period change required disabling of +interrupts, not just preemption. + +<p> +For RCU's part, the preemptible-RCU <tt>rcu_read_unlock()</tt> +implementation must be written carefully to avoid similar deadlocks. +In particular, <tt>rcu_read_unlock()</tt> must tolerate an +interrupt where the interrupt handler invokes both +<tt>rcu_read_lock()</tt> and <tt>rcu_read_unlock()</tt>. +This possibility requires <tt>rcu_read_unlock()</tt> to use +negative nesting levels to avoid destructive recursion via +interrupt handler's use of RCU. + +<p> +This pair of mutual scheduler-RCU requirements came as a +<a href="https://lwn.net/Articles/453002/">complete surprise</a>. + +<p> +As noted above, RCU makes use of kthreads, and it is necessary to +avoid excessive CPU-time accumulation by these kthreads. +This requirement was no surprise, but RCU's violation of it +when running context-switch-heavy workloads when built with +<tt>CONFIG_NO_HZ_FULL=y</tt> +<a href="http://www.rdrop.com/users/paulmck/scalability/paper/BareMetal.2015.01.15b.pdf">did come as a surprise [PDF]</a>. +RCU has made good progress towards meeting this requirement, even +for context-switch-have <tt>CONFIG_NO_HZ_FULL=y</tt> workloads, +but there is room for further improvement. + +<h3><a name="Tracing and RCU">Tracing and RCU</a></h3> + +<p> +It is possible to use tracing on RCU code, but tracing itself +uses RCU. +For this reason, <tt>rcu_dereference_raw_notrace()</tt> +is provided for use by tracing, which avoids the destructive +recursion that could otherwise ensue. +This API is also used by virtualization in some architectures, +where RCU readers execute in environments in which tracing +cannot be used. +The tracing folks both located the requirement and provided the +needed fix, so this surprise requirement was relatively painless. + +<h3><a name="Energy Efficiency">Energy Efficiency</a></h3> + +<p> +Interrupting idle CPUs is considered socially unacceptable, +especially by people with battery-powered embedded systems. +RCU therefore conserves energy by detecting which CPUs are +idle, including tracking CPUs that have been interrupted from idle. +This is a large part of the energy-efficiency requirement, +so I learned of this via an irate phone call. + +<p> +Because RCU avoids interrupting idle CPUs, it is illegal to +execute an RCU read-side critical section on an idle CPU. +(Kernels built with <tt>CONFIG_PROVE_RCU=y</tt> will splat +if you try it.) +The <tt>RCU_NONIDLE()</tt> macro and <tt>_rcuidle</tt> +event tracing is provided to work around this restriction. +In addition, <tt>rcu_is_watching()</tt> may be used to +test whether or not it is currently legal to run RCU read-side +critical sections on this CPU. +I learned of the need for diagnostics on the one hand +and <tt>RCU_NONIDLE()</tt> on the other while inspecting +idle-loop code. +Steven Rostedt supplied <tt>_rcuidle</tt> event tracing, +which is used quite heavily in the idle loop. + +<p> +It is similarly socially unacceptable to interrupt an +<tt>nohz_full</tt> CPU running in userspace. +RCU must therefore track <tt>nohz_full</tt> userspace +execution. +And in +<a href="https://lwn.net/Articles/558284/"><tt>CONFIG_NO_HZ_FULL_SYSIDLE=y</tt></a> +kernels, RCU must separately track idle CPUs on the one hand and +CPUs that are either idle or executing in userspace on the other. +In both cases, RCU must be able to sample state at two points in +time, and be able to determine whether or not some other CPU spent +any time idle and/or executing in userspace. + +<p> +These energy-efficiency requirements have proven quite difficult to +understand and to meet, for example, there have been more than five +clean-sheet rewrites of RCU's energy-efficiency code, the last of +which was finally able to demonstrate +<a href="http://www.rdrop.com/users/paulmck/realtime/paper/AMPenergy.2013.04.19a.pdf">real energy savings running on real hardware [PDF]</a>. +As noted earlier, +I learned of many of these requirements via angry phone calls: +Flaming me on the Linux-kernel mailing list was apparently not +sufficient to fully vent their ire at RCU's energy-efficiency bugs! + +<h3><a name="Memory Efficiency">Memory Efficiency</a></h3> + +<p> +Although small-memory non-realtime systems can simply use Tiny RCU, +code size is only one aspect of memory efficiency. +Another aspect is the size of the <tt>rcu_head</tt> structure +used by <tt>call_rcu()</tt> and <tt>kfree_rcu()</tt>. +Although this structure contains nothing more than a pair of pointers, +it does appear in many RCU-protected data structures, including +some that are size critical. +The <tt>page</tt> structure is a case in point, as evidenced by +the many occurrences of the <tt>union</tt> keyword within that structure. + +<p> +This need for memory efficiency is one reason that RCU uses hand-crafted +singly linked lists to track the <tt>rcu_head</tt> structures that +are waiting for a grace period to elapse. +It is also the reason why <tt>rcu_head</tt> structures do not contain +debug information, such as fields tracking the file and line of the +<tt>call_rcu()</tt> or <tt>kfree_rcu()</tt> that posted them. +Although this information might appear in debug-only kernel builds at some +point, in the meantime, the <tt>->func</tt> field will often provide +the needed debug information. + +<p> +However, in some cases, the need for memory efficiency leads to even +more extreme measures. +Returning to the <tt>page</tt> structure, the <tt>rcu_head</tt> field +shares storage with a great many other structures that are used at +various points in the corresponding page's lifetime. +In order to correctly resolve certain +<a href="https://lkml.kernel.org/g/1439976106-137226-1-git-send-email-kirill.shutemov@linux.intel.com">race conditions</a>, +the Linux kernel's memory-management subsystem needs a particular bit +to remain zero during all phases of grace-period processing, +and that bit happens to map to the bottom bit of the +<tt>rcu_head</tt> structure's <tt>->next</tt> field. +RCU makes this guarantee as long as <tt>call_rcu()</tt> +is used to post the callback, as opposed to <tt>kfree_rcu()</tt> +or some future “lazy” +variant of <tt>call_rcu()</tt> that might one day be created for +energy-efficiency purposes. + +<h3><a name="Performance, Scalability, Response Time, and Reliability"> +Performance, Scalability, Response Time, and Reliability</a></h3> + +<p> +Expanding on the +<a href="#Performance and Scalability">earlier discussion</a>, +RCU is used heavily by hot code paths in performance-critical +portions of the Linux kernel's networking, security, virtualization, +and scheduling code paths. +RCU must therefore use efficient implementations, especially in its +read-side primitives. +To that end, it would be good if preemptible RCU's implementation +of <tt>rcu_read_lock()</tt> could be inlined, however, doing +this requires resolving <tt>#include</tt> issues with the +<tt>task_struct</tt> structure. + +<p> +The Linux kernel supports hardware configurations with up to +4096 CPUs, which means that RCU must be extremely scalable. +Algorithms that involve frequent acquisitions of global locks or +frequent atomic operations on global variables simply cannot be +tolerated within the RCU implementation. +RCU therefore makes heavy use of a combining tree based on the +<tt>rcu_node</tt> structure. +RCU is required to tolerate all CPUs continuously invoking any +combination of RCU's runtime primitives with minimal per-operation +overhead. +In fact, in many cases, increasing load must <i>decrease</i> the +per-operation overhead, witness the batching optimizations for +<tt>synchronize_rcu()</tt>, <tt>call_rcu()</tt>, +<tt>synchronize_rcu_expedited()</tt>, and <tt>rcu_barrier()</tt>. +As a general rule, RCU must cheerfully accept whatever the +rest of the Linux kernel decides to throw at it. + +<p> +The Linux kernel is used for real-time workloads, especially +in conjunction with the +<a href="https://rt.wiki.kernel.org/index.php/Main_Page">-rt patchset</a>. +The real-time-latency response requirements are such that the +traditional approach of disabling preemption across RCU +read-side critical sections is inappropriate. +Kernels built with <tt>CONFIG_PREEMPT=y</tt> therefore +use an RCU implementation that allows RCU read-side critical +sections to be preempted. +This requirement made its presence known after users made it +clear that an earlier +<a href="https://lwn.net/Articles/107930/">real-time patch</a> +did not meet their needs, in conjunction with some +<a href="https://lkml.kernel.org/g/20050318002026.GA2693@us.ibm.com">RCU issues</a> +encountered by a very early version of the -rt patchset. + +<p> +In addition, RCU must make do with a sub-100-microsecond real-time latency +budget. +In fact, on smaller systems with the -rt patchset, the Linux kernel +provides sub-20-microsecond real-time latencies for the whole kernel, +including RCU. +RCU's scalability and latency must therefore be sufficient for +these sorts of configurations. +To my surprise, the sub-100-microsecond real-time latency budget +<a href="http://www.rdrop.com/users/paulmck/realtime/paper/bigrt.2013.01.31a.LCA.pdf"> +applies to even the largest systems [PDF]</a>, +up to and including systems with 4096 CPUs. +This real-time requirement motivated the grace-period kthread, which +also simplified handling of a number of race conditions. + +<p> +Finally, RCU's status as a synchronization primitive means that +any RCU failure can result in arbitrary memory corruption that can be +extremely difficult to debug. +This means that RCU must be extremely reliable, which in +practice also means that RCU must have an aggressive stress-test +suite. +This stress-test suite is called <tt>rcutorture</tt>. + +<p> +Although the need for <tt>rcutorture</tt> was no surprise, +the current immense popularity of the Linux kernel is posing +interesting—and perhaps unprecedented—validation +challenges. +To see this, keep in mind that there are well over one billion +instances of the Linux kernel running today, given Android +smartphones, Linux-powered televisions, and servers. +This number can be expected to increase sharply with the advent of +the celebrated Internet of Things. + +<p> +Suppose that RCU contains a race condition that manifests on average +once per million years of runtime. +This bug will be occurring about three times per <i>day</i> across +the installed base. +RCU could simply hide behind hardware error rates, given that no one +should really expect their smartphone to last for a million years. +However, anyone taking too much comfort from this thought should +consider the fact that in most jurisdictions, a successful multi-year +test of a given mechanism, which might include a Linux kernel, +suffices for a number of types of safety-critical certifications. +In fact, rumor has it that the Linux kernel is already being used +in production for safety-critical applications. +I don't know about you, but I would feel quite bad if a bug in RCU +killed someone. +Which might explain my recent focus on validation and verification. + +<h2><a name="Other RCU Flavors">Other RCU Flavors</a></h2> + +<p> +One of the more surprising things about RCU is that there are now +no fewer than five <i>flavors</i>, or API families. +In addition, the primary flavor that has been the sole focus up to +this point has two different implementations, non-preemptible and +preemptible. +The other four flavors are listed below, with requirements for each +described in a separate section. + +<ol> +<li> <a href="#Bottom-Half Flavor">Bottom-Half Flavor</a> +<li> <a href="#Sched Flavor">Sched Flavor</a> +<li> <a href="#Sleepable RCU">Sleepable RCU</a> +<li> <a href="#Tasks RCU">Tasks RCU</a> +</ol> + +<h3><a name="Bottom-Half Flavor">Bottom-Half Flavor</a></h3> + +<p> +The softirq-disable (AKA “bottom-half”, +hence the “_bh” abbreviations) +flavor of RCU, or <i>RCU-bh</i>, was developed by +Dipankar Sarma to provide a flavor of RCU that could withstand the +network-based denial-of-service attacks researched by Robert +Olsson. +These attacks placed so much networking load on the system +that some of the CPUs never exited softirq execution, +which in turn prevented those CPUs from ever executing a context switch, +which, in the RCU implementation of that time, prevented grace periods +from ever ending. +The result was an out-of-memory condition and a system hang. + +<p> +The solution was the creation of RCU-bh, which does +<tt>local_bh_disable()</tt> +across its read-side critical sections, and which uses the transition +from one type of softirq processing to another as a quiescent state +in addition to context switch, idle, user mode, and offline. +This means that RCU-bh grace periods can complete even when some of +the CPUs execute in softirq indefinitely, thus allowing algorithms +based on RCU-bh to withstand network-based denial-of-service attacks. + +<p> +Because +<tt>rcu_read_lock_bh()</tt> and <tt>rcu_read_unlock_bh()</tt> +disable and re-enable softirq handlers, any attempt to start a softirq +handlers during the +RCU-bh read-side critical section will be deferred. +In this case, <tt>rcu_read_unlock_bh()</tt> +will invoke softirq processing, which can take considerable time. +One can of course argue that this softirq overhead should be associated +with the code following the RCU-bh read-side critical section rather +than <tt>rcu_read_unlock_bh()</tt>, but the fact +is that most profiling tools cannot be expected to make this sort +of fine distinction. +For example, suppose that a three-millisecond-long RCU-bh read-side +critical section executes during a time of heavy networking load. +There will very likely be an attempt to invoke at least one softirq +handler during that three milliseconds, but any such invocation will +be delayed until the time of the <tt>rcu_read_unlock_bh()</tt>. +This can of course make it appear at first glance as if +<tt>rcu_read_unlock_bh()</tt> was executing very slowly. + +<p> +The +<a href="https://lwn.net/Articles/609973/#RCU Per-Flavor API Table">RCU-bh API</a> +includes +<tt>rcu_read_lock_bh()</tt>, +<tt>rcu_read_unlock_bh()</tt>, +<tt>rcu_dereference_bh()</tt>, +<tt>rcu_dereference_bh_check()</tt>, +<tt>synchronize_rcu_bh()</tt>, +<tt>synchronize_rcu_bh_expedited()</tt>, +<tt>call_rcu_bh()</tt>, +<tt>rcu_barrier_bh()</tt>, and +<tt>rcu_read_lock_bh_held()</tt>. + +<h3><a name="Sched Flavor">Sched Flavor</a></h3> + +<p> +Before preemptible RCU, waiting for an RCU grace period had the +side effect of also waiting for all pre-existing interrupt +and NMI handlers. +However, there are legitimate preemptible-RCU implementations that +do not have this property, given that any point in the code outside +of an RCU read-side critical section can be a quiescent state. +Therefore, <i>RCU-sched</i> was created, which follows “classic” +RCU in that an RCU-sched grace period waits for for pre-existing +interrupt and NMI handlers. +In kernels built with <tt>CONFIG_PREEMPT=n</tt>, the RCU and RCU-sched +APIs have identical implementations, while kernels built with +<tt>CONFIG_PREEMPT=y</tt> provide a separate implementation for each. + +<p> +Note well that in <tt>CONFIG_PREEMPT=y</tt> kernels, +<tt>rcu_read_lock_sched()</tt> and <tt>rcu_read_unlock_sched()</tt> +disable and re-enable preemption, respectively. +This means that if there was a preemption attempt during the +RCU-sched read-side critical section, <tt>rcu_read_unlock_sched()</tt> +will enter the scheduler, with all the latency and overhead entailed. +Just as with <tt>rcu_read_unlock_bh()</tt>, this can make it look +as if <tt>rcu_read_unlock_sched()</tt> was executing very slowly. +However, the highest-priority task won't be preempted, so that task +will enjoy low-overhead <tt>rcu_read_unlock_sched()</tt> invocations. + +<p> +The +<a href="https://lwn.net/Articles/609973/#RCU Per-Flavor API Table">RCU-sched API</a> +includes +<tt>rcu_read_lock_sched()</tt>, +<tt>rcu_read_unlock_sched()</tt>, +<tt>rcu_read_lock_sched_notrace()</tt>, +<tt>rcu_read_unlock_sched_notrace()</tt>, +<tt>rcu_dereference_sched()</tt>, +<tt>rcu_dereference_sched_check()</tt>, +<tt>synchronize_sched()</tt>, +<tt>synchronize_rcu_sched_expedited()</tt>, +<tt>call_rcu_sched()</tt>, +<tt>rcu_barrier_sched()</tt>, and +<tt>rcu_read_lock_sched_held()</tt>. +However, anything that disables preemption also marks an RCU-sched +read-side critical section, including +<tt>preempt_disable()</tt> and <tt>preempt_enable()</tt>, +<tt>local_irq_save()</tt> and <tt>local_irq_restore()</tt>, +and so on. + +<h3><a name="Sleepable RCU">Sleepable RCU</a></h3> + +<p> +For well over a decade, someone saying “I need to block within +an RCU read-side critical section” was a reliable indication +that this someone did not understand RCU. +After all, if you are always blocking in an RCU read-side critical +section, you can probably afford to use a higher-overhead synchronization +mechanism. +However, that changed with the advent of the Linux kernel's notifiers, +whose RCU read-side critical +sections almost never sleep, but sometimes need to. +This resulted in the introduction of +<a href="https://lwn.net/Articles/202847/">sleepable RCU</a>, +or <i>SRCU</i>. + +<p> +SRCU allows different domains to be defined, with each such domain +defined by an instance of an <tt>srcu_struct</tt> structure. +A pointer to this structure must be passed in to each SRCU function, +for example, <tt>synchronize_srcu(&ss)</tt>, where +<tt>ss</tt> is the <tt>srcu_struct</tt> structure. +The key benefit of these domains is that a slow SRCU reader in one +domain does not delay an SRCU grace period in some other domain. +That said, one consequence of these domains is that read-side code +must pass a “cookie” from <tt>srcu_read_lock()</tt> +to <tt>srcu_read_unlock()</tt>, for example, as follows: + +<blockquote> +<pre> + 1 int idx; + 2 + 3 idx = srcu_read_lock(&ss); + 4 do_something(); + 5 srcu_read_unlock(&ss, idx); +</pre> +</blockquote> + +<p> +As noted above, it is legal to block within SRCU read-side critical sections, +however, with great power comes great responsibility. +If you block forever in one of a given domain's SRCU read-side critical +sections, then that domain's grace periods will also be blocked forever. +Of course, one good way to block forever is to deadlock, which can +happen if any operation in a given domain's SRCU read-side critical +section can block waiting, either directly or indirectly, for that domain's +grace period to elapse. +For example, this results in a self-deadlock: + +<blockquote> +<pre> + 1 int idx; + 2 + 3 idx = srcu_read_lock(&ss); + 4 do_something(); + 5 synchronize_srcu(&ss); + 6 srcu_read_unlock(&ss, idx); +</pre> +</blockquote> + +<p> +However, if line 5 acquired a mutex that was held across +a <tt>synchronize_srcu()</tt> for domain <tt>ss</tt>, +deadlock would still be possible. +Furthermore, if line 5 acquired a mutex that was held across +a <tt>synchronize_srcu()</tt> for some other domain <tt>ss1</tt>, +and if an <tt>ss1</tt>-domain SRCU read-side critical section +acquired another mutex that was held across as <tt>ss</tt>-domain +<tt>synchronize_srcu()</tt>, +deadlock would again be possible. +Such a deadlock cycle could extend across an arbitrarily large number +of different SRCU domains. +Again, with great power comes great responsibility. + +<p> +Unlike the other RCU flavors, SRCU read-side critical sections can +run on idle and even offline CPUs. +This ability requires that <tt>srcu_read_lock()</tt> and +<tt>srcu_read_unlock()</tt> contain memory barriers, which means +that SRCU readers will run a bit slower than would RCU readers. +It also motivates the <tt>smp_mb__after_srcu_read_unlock()</tt> +API, which, in combination with <tt>srcu_read_unlock()</tt>, +guarantees a full memory barrier. + +<p> +The +<a href="https://lwn.net/Articles/609973/#RCU Per-Flavor API Table">SRCU API</a> +includes +<tt>srcu_read_lock()</tt>, +<tt>srcu_read_unlock()</tt>, +<tt>srcu_dereference()</tt>, +<tt>srcu_dereference_check()</tt>, +<tt>synchronize_srcu()</tt>, +<tt>synchronize_srcu_expedited()</tt>, +<tt>call_srcu()</tt>, +<tt>srcu_barrier()</tt>, and +<tt>srcu_read_lock_held()</tt>. +It also includes +<tt>DEFINE_SRCU()</tt>, +<tt>DEFINE_STATIC_SRCU()</tt>, and +<tt>init_srcu_struct()</tt> +APIs for defining and initializing <tt>srcu_struct</tt> structures. + +<h3><a name="Tasks RCU">Tasks RCU</a></h3> + +<p> +Some forms of tracing use “tramopolines” to handle the +binary rewriting required to install different types of probes. +It would be good to be able to free old trampolines, which sounds +like a job for some form of RCU. +However, because it is necessary to be able to install a trace +anywhere in the code, it is not possible to use read-side markers +such as <tt>rcu_read_lock()</tt> and <tt>rcu_read_unlock()</tt>. +In addition, it does not work to have these markers in the trampoline +itself, because there would need to be instructions following +<tt>rcu_read_unlock()</tt>. +Although <tt>synchronize_rcu()</tt> would guarantee that execution +reached the <tt>rcu_read_unlock()</tt>, it would not be able to +guarantee that execution had completely left the trampoline. + +<p> +The solution, in the form of +<a href="https://lwn.net/Articles/607117/"><i>Tasks RCU</i></a>, +is to have implicit +read-side critical sections that are delimited by voluntary context +switches, that is, calls to <tt>schedule()</tt>, +<tt>cond_resched_rcu_qs()</tt>, and +<tt>synchronize_rcu_tasks()</tt>. +In addition, transitions to and from userspace execution also delimit +tasks-RCU read-side critical sections. + +<p> +The tasks-RCU API is quite compact, consisting only of +<tt>call_rcu_tasks()</tt>, +<tt>synchronize_rcu_tasks()</tt>, and +<tt>rcu_barrier_tasks()</tt>. + +<h2><a name="Possible Future Changes">Possible Future Changes</a></h2> + +<p> +One of the tricks that RCU uses to attain update-side scalability is +to increase grace-period latency with increasing numbers of CPUs. +If this becomes a serious problem, it will be necessary to rework the +grace-period state machine so as to avoid the need for the additional +latency. + +<p> +Expedited grace periods scan the CPUs, so their latency and overhead +increases with increasing numbers of CPUs. +If this becomes a serious problem on large systems, it will be necessary +to do some redesign to avoid this scalability problem. + +<p> +RCU disables CPU hotplug in a few places, perhaps most notably in the +expedited grace-period and <tt>rcu_barrier()</tt> operations. +If there is a strong reason to use expedited grace periods in CPU-hotplug +notifiers, it will be necessary to avoid disabling CPU hotplug. +This would introduce some complexity, so there had better be a <i>very</i> +good reason. + +<p> +The tradeoff between grace-period latency on the one hand and interruptions +of other CPUs on the other hand may need to be re-examined. +The desire is of course for zero grace-period latency as well as zero +interprocessor interrupts undertaken during an expedited grace period +operation. +While this ideal is unlikely to be achievable, it is quite possible that +further improvements can be made. + +<p> +The multiprocessor implementations of RCU use a combining tree that +groups CPUs so as to reduce lock contention and increase cache locality. +However, this combining tree does not spread its memory across NUMA +nodes nor does it align the CPU groups with hardware features such +as sockets or cores. +Such spreading and alignment is currently believed to be unnecessary +because the hotpath read-side primitives do not access the combining +tree, nor does <tt>call_rcu()</tt> in the common case. +If you believe that your architecture needs such spreading and alignment, +then your architecture should also benefit from the +<tt>rcutree.rcu_fanout_leaf</tt> boot parameter, which can be set +to the number of CPUs in a socket, NUMA node, or whatever. +If the number of CPUs is too large, use a fraction of the number of +CPUs. +If the number of CPUs is a large prime number, well, that certainly +is an “interesting” architectural choice! +More flexible arrangements might be considered, but only if +<tt>rcutree.rcu_fanout_leaf</tt> has proven inadequate, and only +if the inadequacy has been demonstrated by a carefully run and +realistic system-level workload. + +<p> +Please note that arrangements that require RCU to remap CPU numbers will +require extremely good demonstration of need and full exploration of +alternatives. + +<p> +There is an embarrassingly large number of flavors of RCU, and this +number has been increasing over time. +Perhaps it will be possible to combine some at some future date. + +<p> +RCU's various kthreads are reasonably recent additions. +It is quite likely that adjustments will be required to more gracefully +handle extreme loads. +It might also be necessary to be able to relate CPU utilization by +RCU's kthreads and softirq handlers to the code that instigated this +CPU utilization. +For example, RCU callback overhead might be charged back to the +originating <tt>call_rcu()</tt> instance, though probably not +in production kernels. + +<h2><a name="Summary">Summary</a></h2> + +<p> +This document has presented more than two decade's worth of RCU +requirements. +Given that the requirements keep changing, this will not be the last +word on this subject, but at least it serves to get an important +subset of the requirements set forth. + +<h2><a name="Acknowledgments">Acknowledgments</a></h2> + +I am grateful to Steven Rostedt, Lai Jiangshan, Ingo Molnar, +Oleg Nesterov, Borislav Petkov, Peter Zijlstra, Boqun Feng, and +Andy Lutomirski for their help in rendering +this article human readable, and to Michelle Rankin for her support +of this effort. +Other contributions are acknowledged in the Linux kernel's git archive. +The cartoon is copyright (c) 2013 by Melissa Broussard, +and is provided +under the terms of the Creative Commons Attribution-Share Alike 3.0 +United States license. + +<p>@@QQAL@@ + +</body></html> diff --git a/Documentation/RCU/Design/htmlqqz.sh b/Documentation/RCU/Design/htmlqqz.sh new file mode 100755 index 000000000000..d354f069559b --- /dev/null +++ b/Documentation/RCU/Design/htmlqqz.sh @@ -0,0 +1,108 @@ +#!/bin/sh +# +# Usage: sh htmlqqz.sh file +# +# Extracts and converts quick quizzes in a proto-HTML document file.htmlx. +# Commands, all of which must be on a line by themselves: +# +# "<p>@@QQ@@": Start of a quick quiz. +# "<p>@@QQA@@": Start of a quick-quiz answer. +# "<p>@@QQE@@": End of a quick-quiz answer, and thus of the quick quiz. +# "<p>@@QQAL@@": Place to put quick-quiz answer list. +# +# Places the result in file.html. +# +# This program is free software; you can redistribute it and/or modify +# it under the terms of the GNU General Public License as published by +# the Free Software Foundation; either version 2 of the License, or +# (at your option) any later version. +# +# This program is distributed in the hope that it will be useful, +# but WITHOUT ANY WARRANTY; without even the implied warranty of +# MERCHANTABILITY or FITNESS FOR A PARTICULAR PURPOSE. See the +# GNU General Public License for more details. +# +# You should have received a copy of the GNU General Public License +# along with this program; if not, you can access it online at +# http://www.gnu.org/licenses/gpl-2.0.html. +# +# Copyright (c) 2013 Paul E. McKenney, IBM Corporation. + +fn=$1 +if test ! -r $fn.htmlx +then + echo "Error: $fn.htmlx unreadable." + exit 1 +fi + +echo "<!-- DO NOT HAND EDIT. -->" > $fn.html +echo "<!-- Instead, edit $fn.htmlx and run 'sh htmlqqz.sh $fn' -->" >> $fn.html +awk < $fn.htmlx >> $fn.html ' + +state == "" && $1 != "<p>@@QQ@@" && $1 != "<p>@@QQAL@@" { + print $0; + if ($0 ~ /^<p>@@QQ/) + print "Bad Quick Quiz command: " NR " (expected <p>@@QQ@@ or <p>@@QQAL@@)." > "/dev/stderr" + next; +} + +state == "" && $1 == "<p>@@QQ@@" { + qqn++; + qqlineno = NR; + haveqq = 1; + state = "qq"; + print "<p><a name=\"Quick Quiz " qqn "\"><b>Quick Quiz " qqn "</b>:</a>" + next; +} + +state == "qq" && $1 != "<p>@@QQA@@" { + qq[qqn] = qq[qqn] $0 "\n"; + print $0 + if ($0 ~ /^<p>@@QQ/) + print "Bad Quick Quiz command: " NR ". (expected <p>@@QQA@@)" > "/dev/stderr" + next; +} + +state == "qq" && $1 == "<p>@@QQA@@" { + state = "qqa"; + print "<br><a href=\"#qq" qqn "answer\">Answer</a>" + next; +} + +state == "qqa" && $1 != "<p>@@QQE@@" { + qqa[qqn] = qqa[qqn] $0 "\n"; + if ($0 ~ /^<p>@@QQ/) + print "Bad Quick Quiz command: " NR " (expected <p>@@QQE@@)." > "/dev/stderr" + next; +} + +state == "qqa" && $1 == "<p>@@QQE@@" { + state = ""; + next; +} + +state == "" && $1 == "<p>@@QQAL@@" { + haveqq = ""; + print "<h3><a name=\"Answers to Quick Quizzes\">" + print "Answers to Quick Quizzes</a></h3>" + print ""; + for (i = 1; i <= qqn; i++) { + print "<a name=\"qq" i "answer\"></a>" + print "<p><b>Quick Quiz " i "</b>:" + print qq[i]; + print ""; + print "</p><p><b>Answer</b>:" + print qqa[i]; + print ""; + print "</p><p><a href=\"#Quick%20Quiz%20" i "\"><b>Back to Quick Quiz " i "</b>.</a>" + print ""; + } + next; +} + +END { + if (state != "") + print "Unterminated Quick Quiz: " qqlineno "." > "/dev/stderr" + else if (haveqq) + print "Missing \"<p>@@QQAL@@\", no Quick Quiz." > "/dev/stderr" +}' diff --git a/Documentation/arm64/silicon-errata.txt b/Documentation/arm64/silicon-errata.txt new file mode 100644 index 000000000000..58b71ddf9b60 --- /dev/null +++ b/Documentation/arm64/silicon-errata.txt @@ -0,0 +1,58 @@ + Silicon Errata and Software Workarounds + ======================================= + +Author: Will Deacon <will.deacon@arm.com> +Date : 27 November 2015 + +It is an unfortunate fact of life that hardware is often produced with +so-called "errata", which can cause it to deviate from the architecture +under specific circumstances. For hardware produced by ARM, these +errata are broadly classified into the following categories: + + Category A: A critical error without a viable workaround. + Category B: A significant or critical error with an acceptable + workaround. + Category C: A minor error that is not expected to occur under normal + operation. + +For more information, consult one of the "Software Developers Errata +Notice" documents available on infocenter.arm.com (registration +required). + +As far as Linux is concerned, Category B errata may require some special +treatment in the operating system. For example, avoiding a particular +sequence of code, or configuring the processor in a particular way. A +less common situation may require similar actions in order to declassify +a Category A erratum into a Category C erratum. These are collectively +known as "software workarounds" and are only required in the minority of +cases (e.g. those cases that both require a non-secure workaround *and* +can be triggered by Linux). + +For software workarounds that may adversely impact systems unaffected by +the erratum in question, a Kconfig entry is added under "Kernel +Features" -> "ARM errata workarounds via the alternatives framework". +These are enabled by default and patched in at runtime when an affected +CPU is detected. For less-intrusive workarounds, a Kconfig option is not +available and the code is structured (preferably with a comment) in such +a way that the erratum will not be hit. + +This approach can make it slightly onerous to determine exactly which +errata are worked around in an arbitrary kernel source tree, so this +file acts as a registry of software workarounds in the Linux Kernel and +will be updated when new workarounds are committed and backported to +stable kernels. + +| Implementor | Component | Erratum ID | Kconfig | ++----------------+-----------------+-----------------+-------------------------+ +| ARM | Cortex-A53 | #826319 | ARM64_ERRATUM_826319 | +| ARM | Cortex-A53 | #827319 | ARM64_ERRATUM_827319 | +| ARM | Cortex-A53 | #824069 | ARM64_ERRATUM_824069 | +| ARM | Cortex-A53 | #819472 | ARM64_ERRATUM_819472 | +| ARM | Cortex-A53 | #845719 | ARM64_ERRATUM_845719 | +| ARM | Cortex-A53 | #843419 | ARM64_ERRATUM_843419 | +| ARM | Cortex-A57 | #832075 | ARM64_ERRATUM_832075 | +| ARM | Cortex-A57 | #852523 | N/A | +| ARM | Cortex-A57 | #834220 | ARM64_ERRATUM_834220 | +| | | | | +| Cavium | ThunderX ITS | #22375, #24313 | CAVIUM_ERRATUM_22375 | +| Cavium | ThunderX GICv3 | #23154 | CAVIUM_ERRATUM_23154 | diff --git a/Documentation/cgroups/00-INDEX b/Documentation/cgroup-v1/00-INDEX index 3f5a40f57d4a..6ad425f7cf56 100644 --- a/Documentation/cgroups/00-INDEX +++ b/Documentation/cgroup-v1/00-INDEX @@ -24,7 +24,5 @@ net_prio.txt - Network priority cgroups details and usages. pids.txt - Process number cgroups details and usages. -resource_counter.txt - - Resource Counter API. unified-hierarchy.txt - Description the new/next cgroup interface. diff --git a/Documentation/cgroups/blkio-controller.txt b/Documentation/cgroup-v1/blkio-controller.txt index 52fa9f353342..673dc34d3f78 100644 --- a/Documentation/cgroups/blkio-controller.txt +++ b/Documentation/cgroup-v1/blkio-controller.txt @@ -84,8 +84,7 @@ Throttling/Upper Limit policy - Run dd to read a file and see if rate is throttled to 1MB/s or not. - # dd if=/mnt/common/zerofile of=/dev/null bs=4K count=1024 - # iflag=direct + # dd iflag=direct if=/mnt/common/zerofile of=/dev/null bs=4K count=1024 1024+0 records in 1024+0 records out 4194304 bytes (4.2 MB) copied, 4.0001 s, 1.0 MB/s @@ -374,82 +373,3 @@ One can experience an overall throughput drop if you have created multiple groups and put applications in that group which are not driving enough IO to keep disk busy. In that case set group_idle=0, and CFQ will not idle on individual groups and throughput should improve. - -Writeback -========= - -Page cache is dirtied through buffered writes and shared mmaps and -written asynchronously to the backing filesystem by the writeback -mechanism. Writeback sits between the memory and IO domains and -regulates the proportion of dirty memory by balancing dirtying and -write IOs. - -On traditional cgroup hierarchies, relationships between different -controllers cannot be established making it impossible for writeback -to operate accounting for cgroup resource restrictions and all -writeback IOs are attributed to the root cgroup. - -If both the blkio and memory controllers are used on the v2 hierarchy -and the filesystem supports cgroup writeback, writeback operations -correctly follow the resource restrictions imposed by both memory and -blkio controllers. - -Writeback examines both system-wide and per-cgroup dirty memory status -and enforces the more restrictive of the two. Also, writeback control -parameters which are absolute values - vm.dirty_bytes and -vm.dirty_background_bytes - are distributed across cgroups according -to their current writeback bandwidth. - -There's a peculiarity stemming from the discrepancy in ownership -granularity between memory controller and writeback. While memory -controller tracks ownership per page, writeback operates on inode -basis. cgroup writeback bridges the gap by tracking ownership by -inode but migrating ownership if too many foreign pages, pages which -don't match the current inode ownership, have been encountered while -writing back the inode. - -This is a conscious design choice as writeback operations are -inherently tied to inodes making strictly following page ownership -complicated and inefficient. The only use case which suffers from -this compromise is multiple cgroups concurrently dirtying disjoint -regions of the same inode, which is an unlikely use case and decided -to be unsupported. Note that as memory controller assigns page -ownership on the first use and doesn't update it until the page is -released, even if cgroup writeback strictly follows page ownership, -multiple cgroups dirtying overlapping areas wouldn't work as expected. -In general, write-sharing an inode across multiple cgroups is not well -supported. - -Filesystem support for cgroup writeback ---------------------------------------- - -A filesystem can make writeback IOs cgroup-aware by updating -address_space_operations->writepage[s]() to annotate bio's using the -following two functions. - -* wbc_init_bio(@wbc, @bio) - - Should be called for each bio carrying writeback data and associates - the bio with the inode's owner cgroup. Can be called anytime - between bio allocation and submission. - -* wbc_account_io(@wbc, @page, @bytes) - - Should be called for each data segment being written out. While - this function doesn't care exactly when it's called during the - writeback session, it's the easiest and most natural to call it as - data segments are added to a bio. - -With writeback bio's annotated, cgroup support can be enabled per -super_block by setting MS_CGROUPWB in ->s_flags. This allows for -selective disabling of cgroup writeback support which is helpful when -certain filesystem features, e.g. journaled data mode, are -incompatible. - -wbc_init_bio() binds the specified bio to its cgroup. Depending on -the configuration, the bio may be executed at a lower priority and if -the writeback session is holding shared resources, e.g. a journal -entry, may lead to priority inversion. There is no one easy solution -for the problem. Filesystems can try to work around specific problem -cases by skipping wbc_init_bio() or using bio_associate_blkcg() -directly. diff --git a/Documentation/cgroups/cgroups.txt b/Documentation/cgroup-v1/cgroups.txt index c6256ae9885b..c6256ae9885b 100644 --- a/Documentation/cgroups/cgroups.txt +++ b/Documentation/cgroup-v1/cgroups.txt diff --git a/Documentation/cgroups/cpuacct.txt b/Documentation/cgroup-v1/cpuacct.txt index 9d73cc0cadb9..9d73cc0cadb9 100644 --- a/Documentation/cgroups/cpuacct.txt +++ b/Documentation/cgroup-v1/cpuacct.txt diff --git a/Documentation/cgroups/cpusets.txt b/Documentation/cgroup-v1/cpusets.txt index fdf7dff3f607..fdf7dff3f607 100644 --- a/Documentation/cgroups/cpusets.txt +++ b/Documentation/cgroup-v1/cpusets.txt diff --git a/Documentation/cgroups/devices.txt b/Documentation/cgroup-v1/devices.txt index 3c1095ca02ea..3c1095ca02ea 100644 --- a/Documentation/cgroups/devices.txt +++ b/Documentation/cgroup-v1/devices.txt diff --git a/Documentation/cgroups/freezer-subsystem.txt b/Documentation/cgroup-v1/freezer-subsystem.txt index e831cb2b8394..e831cb2b8394 100644 --- a/Documentation/cgroups/freezer-subsystem.txt +++ b/Documentation/cgroup-v1/freezer-subsystem.txt diff --git a/Documentation/cgroups/hugetlb.txt b/Documentation/cgroup-v1/hugetlb.txt index 106245c3aecc..106245c3aecc 100644 --- a/Documentation/cgroups/hugetlb.txt +++ b/Documentation/cgroup-v1/hugetlb.txt diff --git a/Documentation/cgroups/memcg_test.txt b/Documentation/cgroup-v1/memcg_test.txt index 8870b0212150..8870b0212150 100644 --- a/Documentation/cgroups/memcg_test.txt +++ b/Documentation/cgroup-v1/memcg_test.txt diff --git a/Documentation/cgroups/memory.txt b/Documentation/cgroup-v1/memory.txt index ff71e16cc752..ff71e16cc752 100644 --- a/Documentation/cgroups/memory.txt +++ b/Documentation/cgroup-v1/memory.txt diff --git a/Documentation/cgroups/net_cls.txt b/Documentation/cgroup-v1/net_cls.txt index ec182346dea2..ec182346dea2 100644 --- a/Documentation/cgroups/net_cls.txt +++ b/Documentation/cgroup-v1/net_cls.txt diff --git a/Documentation/cgroups/net_prio.txt b/Documentation/cgroup-v1/net_prio.txt index a82cbd28ea8a..a82cbd28ea8a 100644 --- a/Documentation/cgroups/net_prio.txt +++ b/Documentation/cgroup-v1/net_prio.txt diff --git a/Documentation/cgroups/pids.txt b/Documentation/cgroup-v1/pids.txt index 1a078b5d281a..1a078b5d281a 100644 --- a/Documentation/cgroups/pids.txt +++ b/Documentation/cgroup-v1/pids.txt diff --git a/Documentation/cgroup-v2.txt b/Documentation/cgroup-v2.txt new file mode 100644 index 000000000000..31d1f7bf12a1 --- /dev/null +++ b/Documentation/cgroup-v2.txt @@ -0,0 +1,1293 @@ + +Control Group v2 + +October, 2015 Tejun Heo <tj@kernel.org> + +This is the authoritative documentation on the design, interface and +conventions of cgroup v2. It describes all userland-visible aspects +of cgroup including core and specific controller behaviors. All +future changes must be reflected in this document. Documentation for +v1 is available under Documentation/cgroup-legacy/. + +CONTENTS + +1. Introduction + 1-1. Terminology + 1-2. What is cgroup? +2. Basic Operations + 2-1. Mounting + 2-2. Organizing Processes + 2-3. [Un]populated Notification + 2-4. Controlling Controllers + 2-4-1. Enabling and Disabling + 2-4-2. Top-down Constraint + 2-4-3. No Internal Process Constraint + 2-5. Delegation + 2-5-1. Model of Delegation + 2-5-2. Delegation Containment + 2-6. Guidelines + 2-6-1. Organize Once and Control + 2-6-2. Avoid Name Collisions +3. Resource Distribution Models + 3-1. Weights + 3-2. Limits + 3-3. Protections + 3-4. Allocations +4. Interface Files + 4-1. Format + 4-2. Conventions + 4-3. Core Interface Files +5. Controllers + 5-1. CPU + 5-1-1. CPU Interface Files + 5-2. Memory + 5-2-1. Memory Interface Files + 5-2-2. Usage Guidelines + 5-2-3. Memory Ownership + 5-3. IO + 5-3-1. IO Interface Files + 5-3-2. Writeback +P. Information on Kernel Programming + P-1. Filesystem Support for Writeback +D. Deprecated v1 Core Features +R. Issues with v1 and Rationales for v2 + R-1. Multiple Hierarchies + R-2. Thread Granularity + R-3. Competition Between Inner Nodes and Threads + R-4. Other Interface Issues + R-5. Controller Issues and Remedies + R-5-1. Memory + + +1. Introduction + +1-1. Terminology + +"cgroup" stands for "control group" and is never capitalized. The +singular form is used to designate the whole feature and also as a +qualifier as in "cgroup controllers". When explicitly referring to +multiple individual control groups, the plural form "cgroups" is used. + + +1-2. What is cgroup? + +cgroup is a mechanism to organize processes hierarchically and +distribute system resources along the hierarchy in a controlled and +configurable manner. + +cgroup is largely composed of two parts - the core and controllers. +cgroup core is primarily responsible for hierarchically organizing +processes. A cgroup controller is usually responsible for +distributing a specific type of system resource along the hierarchy +although there are utility controllers which serve purposes other than +resource distribution. + +cgroups form a tree structure and every process in the system belongs +to one and only one cgroup. All threads of a process belong to the +same cgroup. On creation, all processes are put in the cgroup that +the parent process belongs to at the time. A process can be migrated +to another cgroup. Migration of a process doesn't affect already +existing descendant processes. + +Following certain structural constraints, controllers may be enabled or +disabled selectively on a cgroup. All controller behaviors are +hierarchical - if a controller is enabled on a cgroup, it affects all +processes which belong to the cgroups consisting the inclusive +sub-hierarchy of the cgroup. When a controller is enabled on a nested +cgroup, it always restricts the resource distribution further. The +restrictions set closer to the root in the hierarchy can not be +overridden from further away. + + +2. Basic Operations + +2-1. Mounting + +Unlike v1, cgroup v2 has only single hierarchy. The cgroup v2 +hierarchy can be mounted with the following mount command. + + # mount -t cgroup2 none $MOUNT_POINT + +cgroup2 filesystem has the magic number 0x63677270 ("cgrp"). All +controllers which support v2 and are not bound to a v1 hierarchy are +automatically bound to the v2 hierarchy and show up at the root. +Controllers which are not in active use in the v2 hierarchy can be +bound to other hierarchies. This allows mixing v2 hierarchy with the +legacy v1 multiple hierarchies in a fully backward compatible way. + +A controller can be moved across hierarchies only after the controller +is no longer referenced in its current hierarchy. Because per-cgroup +controller states are destroyed asynchronously and controllers may +have lingering references, a controller may not show up immediately on +the v2 hierarchy after the final umount of the previous hierarchy. +Similarly, a controller should be fully disabled to be moved out of +the unified hierarchy and it may take some time for the disabled +controller to become available for other hierarchies; furthermore, due +to inter-controller dependencies, other controllers may need to be +disabled too. + +While useful for development and manual configurations, moving +controllers dynamically between the v2 and other hierarchies is +strongly discouraged for production use. It is recommended to decide +the hierarchies and controller associations before starting using the +controllers after system boot. + + +2-2. Organizing Processes + +Initially, only the root cgroup exists to which all processes belong. +A child cgroup can be created by creating a sub-directory. + + # mkdir $CGROUP_NAME + +A given cgroup may have multiple child cgroups forming a tree +structure. Each cgroup has a read-writable interface file +"cgroup.procs". When read, it lists the PIDs of all processes which +belong to the cgroup one-per-line. The PIDs are not ordered and the +same PID may show up more than once if the process got moved to +another cgroup and then back or the PID got recycled while reading. + +A process can be migrated into a cgroup by writing its PID to the +target cgroup's "cgroup.procs" file. Only one process can be migrated +on a single write(2) call. If a process is composed of multiple +threads, writing the PID of any thread migrates all threads of the +process. + +When a process forks a child process, the new process is born into the +cgroup that the forking process belongs to at the time of the +operation. After exit, a process stays associated with the cgroup +that it belonged to at the time of exit until it's reaped; however, a +zombie process does not appear in "cgroup.procs" and thus can't be +moved to another cgroup. + +A cgroup which doesn't have any children or live processes can be +destroyed by removing the directory. Note that a cgroup which doesn't +have any children and is associated only with zombie processes is +considered empty and can be removed. + + # rmdir $CGROUP_NAME + +"/proc/$PID/cgroup" lists a process's cgroup membership. If legacy +cgroup is in use in the system, this file may contain multiple lines, +one for each hierarchy. The entry for cgroup v2 is always in the +format "0::$PATH". + + # cat /proc/842/cgroup + ... + 0::/test-cgroup/test-cgroup-nested + +If the process becomes a zombie and the cgroup it was associated with +is removed subsequently, " (deleted)" is appended to the path. + + # cat /proc/842/cgroup + ... + 0::/test-cgroup/test-cgroup-nested (deleted) + + +2-3. [Un]populated Notification + +Each non-root cgroup has a "cgroup.events" file which contains +"populated" field indicating whether the cgroup's sub-hierarchy has +live processes in it. Its value is 0 if there is no live process in +the cgroup and its descendants; otherwise, 1. poll and [id]notify +events are triggered when the value changes. This can be used, for +example, to start a clean-up operation after all processes of a given +sub-hierarchy have exited. The populated state updates and +notifications are recursive. Consider the following sub-hierarchy +where the numbers in the parentheses represent the numbers of processes +in each cgroup. + + A(4) - B(0) - C(1) + \ D(0) + +A, B and C's "populated" fields would be 1 while D's 0. After the one +process in C exits, B and C's "populated" fields would flip to "0" and +file modified events will be generated on the "cgroup.events" files of +both cgroups. + + +2-4. Controlling Controllers + +2-4-1. Enabling and Disabling + +Each cgroup has a "cgroup.controllers" file which lists all +controllers available for the cgroup to enable. + + # cat cgroup.controllers + cpu io memory + +No controller is enabled by default. Controllers can be enabled and +disabled by writing to the "cgroup.subtree_control" file. + + # echo "+cpu +memory -io" > cgroup.subtree_control + +Only controllers which are listed in "cgroup.controllers" can be +enabled. When multiple operations are specified as above, either they +all succeed or fail. If multiple operations on the same controller +are specified, the last one is effective. + +Enabling a controller in a cgroup indicates that the distribution of +the target resource across its immediate children will be controlled. +Consider the following sub-hierarchy. The enabled controllers are +listed in parentheses. + + A(cpu,memory) - B(memory) - C() + \ D() + +As A has "cpu" and "memory" enabled, A will control the distribution +of CPU cycles and memory to its children, in this case, B. As B has +"memory" enabled but not "CPU", C and D will compete freely on CPU +cycles but their division of memory available to B will be controlled. + +As a controller regulates the distribution of the target resource to +the cgroup's children, enabling it creates the controller's interface +files in the child cgroups. In the above example, enabling "cpu" on B +would create the "cpu." prefixed controller interface files in C and +D. Likewise, disabling "memory" from B would remove the "memory." +prefixed controller interface files from C and D. This means that the +controller interface files - anything which doesn't start with +"cgroup." are owned by the parent rather than the cgroup itself. + + +2-4-2. Top-down Constraint + +Resources are distributed top-down and a cgroup can further distribute +a resource only if the resource has been distributed to it from the +parent. This means that all non-root "cgroup.subtree_control" files +can only contain controllers which are enabled in the parent's +"cgroup.subtree_control" file. A controller can be enabled only if +the parent has the controller enabled and a controller can't be +disabled if one or more children have it enabled. + + +2-4-3. No Internal Process Constraint + +Non-root cgroups can only distribute resources to their children when +they don't have any processes of their own. In other words, only +cgroups which don't contain any processes can have controllers enabled +in their "cgroup.subtree_control" files. + +This guarantees that, when a controller is looking at the part of the +hierarchy which has it enabled, processes are always only on the +leaves. This rules out situations where child cgroups compete against +internal processes of the parent. + +The root cgroup is exempt from this restriction. Root contains +processes and anonymous resource consumption which can't be associated +with any other cgroups and requires special treatment from most +controllers. How resource consumption in the root cgroup is governed +is up to each controller. + +Note that the restriction doesn't get in the way if there is no +enabled controller in the cgroup's "cgroup.subtree_control". This is +important as otherwise it wouldn't be possible to create children of a +populated cgroup. To control resource distribution of a cgroup, the +cgroup must create children and transfer all its processes to the +children before enabling controllers in its "cgroup.subtree_control" +file. + + +2-5. Delegation + +2-5-1. Model of Delegation + +A cgroup can be delegated to a less privileged user by granting write +access of the directory and its "cgroup.procs" file to the user. Note +that resource control interface files in a given directory control the +distribution of the parent's resources and thus must not be delegated +along with the directory. + +Once delegated, the user can build sub-hierarchy under the directory, +organize processes as it sees fit and further distribute the resources +it received from the parent. The limits and other settings of all +resource controllers are hierarchical and regardless of what happens +in the delegated sub-hierarchy, nothing can escape the resource +restrictions imposed by the parent. + +Currently, cgroup doesn't impose any restrictions on the number of +cgroups in or nesting depth of a delegated sub-hierarchy; however, +this may be limited explicitly in the future. + + +2-5-2. Delegation Containment + +A delegated sub-hierarchy is contained in the sense that processes +can't be moved into or out of the sub-hierarchy by the delegatee. For +a process with a non-root euid to migrate a target process into a +cgroup by writing its PID to the "cgroup.procs" file, the following +conditions must be met. + +- The writer's euid must match either uid or suid of the target process. + +- The writer must have write access to the "cgroup.procs" file. + +- The writer must have write access to the "cgroup.procs" file of the + common ancestor of the source and destination cgroups. + +The above three constraints ensure that while a delegatee may migrate +processes around freely in the delegated sub-hierarchy it can't pull +in from or push out to outside the sub-hierarchy. + +For an example, let's assume cgroups C0 and C1 have been delegated to +user U0 who created C00, C01 under C0 and C10 under C1 as follows and +all processes under C0 and C1 belong to U0. + + ~~~~~~~~~~~~~ - C0 - C00 + ~ cgroup ~ \ C01 + ~ hierarchy ~ + ~~~~~~~~~~~~~ - C1 - C10 + +Let's also say U0 wants to write the PID of a process which is +currently in C10 into "C00/cgroup.procs". U0 has write access to the +file and uid match on the process; however, the common ancestor of the +source cgroup C10 and the destination cgroup C00 is above the points +of delegation and U0 would not have write access to its "cgroup.procs" +files and thus the write will be denied with -EACCES. + + +2-6. Guidelines + +2-6-1. Organize Once and Control + +Migrating a process across cgroups is a relatively expensive operation +and stateful resources such as memory are not moved together with the +process. This is an explicit design decision as there often exist +inherent trade-offs between migration and various hot paths in terms +of synchronization cost. + +As such, migrating processes across cgroups frequently as a means to +apply different resource restrictions is discouraged. A workload +should be assigned to a cgroup according to the system's logical and +resource structure once on start-up. Dynamic adjustments to resource +distribution can be made by changing controller configuration through +the interface files. + + +2-6-2. Avoid Name Collisions + +Interface files for a cgroup and its children cgroups occupy the same +directory and it is possible to create children cgroups which collide +with interface files. + +All cgroup core interface files are prefixed with "cgroup." and each +controller's interface files are prefixed with the controller name and +a dot. A controller's name is composed of lower case alphabets and +'_'s but never begins with an '_' so it can be used as the prefix +character for collision avoidance. Also, interface file names won't +start or end with terms which are often used in categorizing workloads +such as job, service, slice, unit or workload. + +cgroup doesn't do anything to prevent name collisions and it's the +user's responsibility to avoid them. + + +3. Resource Distribution Models + +cgroup controllers implement several resource distribution schemes +depending on the resource type and expected use cases. This section +describes major schemes in use along with their expected behaviors. + + +3-1. Weights + +A parent's resource is distributed by adding up the weights of all +active children and giving each the fraction matching the ratio of its +weight against the sum. As only children which can make use of the +resource at the moment participate in the distribution, this is +work-conserving. Due to the dynamic nature, this model is usually +used for stateless resources. + +All weights are in the range [1, 10000] with the default at 100. This +allows symmetric multiplicative biases in both directions at fine +enough granularity while staying in the intuitive range. + +As long as the weight is in range, all configuration combinations are +valid and there is no reason to reject configuration changes or +process migrations. + +"cpu.weight" proportionally distributes CPU cycles to active children +and is an example of this type. + + +3-2. Limits + +A child can only consume upto the configured amount of the resource. +Limits can be over-committed - the sum of the limits of children can +exceed the amount of resource available to the parent. + +Limits are in the range [0, max] and defaults to "max", which is noop. + +As limits can be over-committed, all configuration combinations are +valid and there is no reason to reject configuration changes or +process migrations. + +"io.max" limits the maximum BPS and/or IOPS that a cgroup can consume +on an IO device and is an example of this type. + + +3-3. Protections + +A cgroup is protected to be allocated upto the configured amount of +the resource if the usages of all its ancestors are under their +protected levels. Protections can be hard guarantees or best effort +soft boundaries. Protections can also be over-committed in which case +only upto the amount available to the parent is protected among +children. + +Protections are in the range [0, max] and defaults to 0, which is +noop. + +As protections can be over-committed, all configuration combinations +are valid and there is no reason to reject configuration changes or +process migrations. + +"memory.low" implements best-effort memory protection and is an +example of this type. + + +3-4. Allocations + +A cgroup is exclusively allocated a certain amount of a finite +resource. Allocations can't be over-committed - the sum of the +allocations of children can not exceed the amount of resource +available to the parent. + +Allocations are in the range [0, max] and defaults to 0, which is no +resource. + +As allocations can't be over-committed, some configuration +combinations are invalid and should be rejected. Also, if the +resource is mandatory for execution of processes, process migrations +may be rejected. + +"cpu.rt.max" hard-allocates realtime slices and is an example of this +type. + + +4. Interface Files + +4-1. Format + +All interface files should be in one of the following formats whenever +possible. + + New-line separated values + (when only one value can be written at once) + + VAL0\n + VAL1\n + ... + + Space separated values + (when read-only or multiple values can be written at once) + + VAL0 VAL1 ...\n + + Flat keyed + + KEY0 VAL0\n + KEY1 VAL1\n + ... + + Nested keyed + + KEY0 SUB_KEY0=VAL00 SUB_KEY1=VAL01... + KEY1 SUB_KEY0=VAL10 SUB_KEY1=VAL11... + ... + +For a writable file, the format for writing should generally match +reading; however, controllers may allow omitting later fields or +implement restricted shortcuts for most common use cases. + +For both flat and nested keyed files, only the values for a single key +can be written at a time. For nested keyed files, the sub key pairs +may be specified in any order and not all pairs have to be specified. + + +4-2. Conventions + +- Settings for a single feature should be contained in a single file. + +- The root cgroup should be exempt from resource control and thus + shouldn't have resource control interface files. Also, + informational files on the root cgroup which end up showing global + information available elsewhere shouldn't exist. + +- If a controller implements weight based resource distribution, its + interface file should be named "weight" and have the range [1, + 10000] with 100 as the default. The values are chosen to allow + enough and symmetric bias in both directions while keeping it + intuitive (the default is 100%). + +- If a controller implements an absolute resource guarantee and/or + limit, the interface files should be named "min" and "max" + respectively. If a controller implements best effort resource + guarantee and/or limit, the interface files should be named "low" + and "high" respectively. + + In the above four control files, the special token "max" should be + used to represent upward infinity for both reading and writing. + +- If a setting has a configurable default value and keyed specific + overrides, the default entry should be keyed with "default" and + appear as the first entry in the file. + + The default value can be updated by writing either "default $VAL" or + "$VAL". + + When writing to update a specific override, "default" can be used as + the value to indicate removal of the override. Override entries + with "default" as the value must not appear when read. + + For example, a setting which is keyed by major:minor device numbers + with integer values may look like the following. + + # cat cgroup-example-interface-file + default 150 + 8:0 300 + + The default value can be updated by + + # echo 125 > cgroup-example-interface-file + + or + + # echo "default 125" > cgroup-example-interface-file + + An override can be set by + + # echo "8:16 170" > cgroup-example-interface-file + + and cleared by + + # echo "8:0 default" > cgroup-example-interface-file + # cat cgroup-example-interface-file + default 125 + 8:16 170 + +- For events which are not very high frequency, an interface file + "events" should be created which lists event key value pairs. + Whenever a notifiable event happens, file modified event should be + generated on the file. + + +4-3. Core Interface Files + +All cgroup core files are prefixed with "cgroup." + + cgroup.procs + + A read-write new-line separated values file which exists on + all cgroups. + + When read, it lists the PIDs of all processes which belong to + the cgroup one-per-line. The PIDs are not ordered and the + same PID may show up more than once if the process got moved + to another cgroup and then back or the PID got recycled while + reading. + + A PID can be written to migrate the process associated with + the PID to the cgroup. The writer should match all of the + following conditions. + + - Its euid is either root or must match either uid or suid of + the target process. + + - It must have write access to the "cgroup.procs" file. + + - It must have write access to the "cgroup.procs" file of the + common ancestor of the source and destination cgroups. + + When delegating a sub-hierarchy, write access to this file + should be granted along with the containing directory. + + cgroup.controllers + + A read-only space separated values file which exists on all + cgroups. + + It shows space separated list of all controllers available to + the cgroup. The controllers are not ordered. + + cgroup.subtree_control + + A read-write space separated values file which exists on all + cgroups. Starts out empty. + + When read, it shows space separated list of the controllers + which are enabled to control resource distribution from the + cgroup to its children. + + Space separated list of controllers prefixed with '+' or '-' + can be written to enable or disable controllers. A controller + name prefixed with '+' enables the controller and '-' + disables. If a controller appears more than once on the list, + the last one is effective. When multiple enable and disable + operations are specified, either all succeed or all fail. + + cgroup.events + + A read-only flat-keyed file which exists on non-root cgroups. + The following entries are defined. Unless specified + otherwise, a value change in this file generates a file + modified event. + + populated + + 1 if the cgroup or its descendants contains any live + processes; otherwise, 0. + + +5. Controllers + +5-1. CPU + +[NOTE: The interface for the cpu controller hasn't been merged yet] + +The "cpu" controllers regulates distribution of CPU cycles. This +controller implements weight and absolute bandwidth limit models for +normal scheduling policy and absolute bandwidth allocation model for +realtime scheduling policy. + + +5-1-1. CPU Interface Files + +All time durations are in microseconds. + + cpu.stat + + A read-only flat-keyed file which exists on non-root cgroups. + + It reports the following six stats. + + usage_usec + user_usec + system_usec + nr_periods + nr_throttled + throttled_usec + + cpu.weight + + A read-write single value file which exists on non-root + cgroups. The default is "100". + + The weight in the range [1, 10000]. + + cpu.max + + A read-write two value file which exists on non-root cgroups. + The default is "max 100000". + + The maximum bandwidth limit. It's in the following format. + + $MAX $PERIOD + + which indicates that the group may consume upto $MAX in each + $PERIOD duration. "max" for $MAX indicates no limit. If only + one number is written, $MAX is updated. + + cpu.rt.max + + [NOTE: The semantics of this file is still under discussion and the + interface hasn't been merged yet] + + A read-write two value file which exists on all cgroups. + The default is "0 100000". + + The maximum realtime runtime allocation. Over-committing + configurations are disallowed and process migrations are + rejected if not enough bandwidth is available. It's in the + following format. + + $MAX $PERIOD + + which indicates that the group may consume upto $MAX in each + $PERIOD duration. If only one number is written, $MAX is + updated. + + +5-2. Memory + +The "memory" controller regulates distribution of memory. Memory is +stateful and implements both limit and protection models. Due to the +intertwining between memory usage and reclaim pressure and the +stateful nature of memory, the distribution model is relatively +complex. + +While not completely water-tight, all major memory usages by a given +cgroup are tracked so that the total memory consumption can be +accounted and controlled to a reasonable extent. Currently, the +following types of memory usages are tracked. + +- Userland memory - page cache and anonymous memory. + +- Kernel data structures such as dentries and inodes. + +- TCP socket buffers. + +The above list may expand in the future for better coverage. + + +5-2-1. Memory Interface Files + +All memory amounts are in bytes. If a value which is not aligned to +PAGE_SIZE is written, the value may be rounded up to the closest +PAGE_SIZE multiple when read back. + + memory.current + + A read-only single value file which exists on non-root + cgroups. + + The total amount of memory currently being used by the cgroup + and its descendants. + + memory.low + + A read-write single value file which exists on non-root + cgroups. The default is "0". + + Best-effort memory protection. If the memory usages of a + cgroup and all its ancestors are below their low boundaries, + the cgroup's memory won't be reclaimed unless memory can be + reclaimed from unprotected cgroups. + + Putting more memory than generally available under this + protection is discouraged. + + memory.high + + A read-write single value file which exists on non-root + cgroups. The default is "max". + + Memory usage throttle limit. This is the main mechanism to + control memory usage of a cgroup. If a cgroup's usage goes + over the high boundary, the processes of the cgroup are + throttled and put under heavy reclaim pressure. + + Going over the high limit never invokes the OOM killer and + under extreme conditions the limit may be breached. + + memory.max + + A read-write single value file which exists on non-root + cgroups. The default is "max". + + Memory usage hard limit. This is the final protection + mechanism. If a cgroup's memory usage reaches this limit and + can't be reduced, the OOM killer is invoked in the cgroup. + Under certain circumstances, the usage may go over the limit + temporarily. + + This is the ultimate protection mechanism. As long as the + high limit is used and monitored properly, this limit's + utility is limited to providing the final safety net. + + memory.events + + A read-only flat-keyed file which exists on non-root cgroups. + The following entries are defined. Unless specified + otherwise, a value change in this file generates a file + modified event. + + low + + The number of times the cgroup is reclaimed due to + high memory pressure even though its usage is under + the low boundary. This usually indicates that the low + boundary is over-committed. + + high + + The number of times processes of the cgroup are + throttled and routed to perform direct memory reclaim + because the high memory boundary was exceeded. For a + cgroup whose memory usage is capped by the high limit + rather than global memory pressure, this event's + occurrences are expected. + + max + + The number of times the cgroup's memory usage was + about to go over the max boundary. If direct reclaim + fails to bring it down, the OOM killer is invoked. + + oom + + The number of times the OOM killer has been invoked in + the cgroup. This may not exactly match the number of + processes killed but should generally be close. + + +5-2-2. General Usage + +"memory.high" is the main mechanism to control memory usage. +Over-committing on high limit (sum of high limits > available memory) +and letting global memory pressure to distribute memory according to +usage is a viable strategy. + +Because breach of the high limit doesn't trigger the OOM killer but +throttles the offending cgroup, a management agent has ample +opportunities to monitor and take appropriate actions such as granting +more memory or terminating the workload. + +Determining whether a cgroup has enough memory is not trivial as +memory usage doesn't indicate whether the workload can benefit from +more memory. For example, a workload which writes data received from +network to a file can use all available memory but can also operate as +performant with a small amount of memory. A measure of memory +pressure - how much the workload is being impacted due to lack of +memory - is necessary to determine whether a workload needs more +memory; unfortunately, memory pressure monitoring mechanism isn't +implemented yet. + + +5-2-3. Memory Ownership + +A memory area is charged to the cgroup which instantiated it and stays +charged to the cgroup until the area is released. Migrating a process +to a different cgroup doesn't move the memory usages that it +instantiated while in the previous cgroup to the new cgroup. + +A memory area may be used by processes belonging to different cgroups. +To which cgroup the area will be charged is in-deterministic; however, +over time, the memory area is likely to end up in a cgroup which has +enough memory allowance to avoid high reclaim pressure. + +If a cgroup sweeps a considerable amount of memory which is expected +to be accessed repeatedly by other cgroups, it may make sense to use +POSIX_FADV_DONTNEED to relinquish the ownership of memory areas +belonging to the affected files to ensure correct memory ownership. + + +5-3. IO + +The "io" controller regulates the distribution of IO resources. This +controller implements both weight based and absolute bandwidth or IOPS +limit distribution; however, weight based distribution is available +only if cfq-iosched is in use and neither scheme is available for +blk-mq devices. + + +5-3-1. IO Interface Files + + io.stat + + A read-only nested-keyed file which exists on non-root + cgroups. + + Lines are keyed by $MAJ:$MIN device numbers and not ordered. + The following nested keys are defined. + + rbytes Bytes read + wbytes Bytes written + rios Number of read IOs + wios Number of write IOs + + An example read output follows. + + 8:16 rbytes=1459200 wbytes=314773504 rios=192 wios=353 + 8:0 rbytes=90430464 wbytes=299008000 rios=8950 wios=1252 + + io.weight + + A read-write flat-keyed file which exists on non-root cgroups. + The default is "default 100". + + The first line is the default weight applied to devices + without specific override. The rest are overrides keyed by + $MAJ:$MIN device numbers and not ordered. The weights are in + the range [1, 10000] and specifies the relative amount IO time + the cgroup can use in relation to its siblings. + + The default weight can be updated by writing either "default + $WEIGHT" or simply "$WEIGHT". Overrides can be set by writing + "$MAJ:$MIN $WEIGHT" and unset by writing "$MAJ:$MIN default". + + An example read output follows. + + default 100 + 8:16 200 + 8:0 50 + + io.max + + A read-write nested-keyed file which exists on non-root + cgroups. + + BPS and IOPS based IO limit. Lines are keyed by $MAJ:$MIN + device numbers and not ordered. The following nested keys are + defined. + + rbps Max read bytes per second + wbps Max write bytes per second + riops Max read IO operations per second + wiops Max write IO operations per second + + When writing, any number of nested key-value pairs can be + specified in any order. "max" can be specified as the value + to remove a specific limit. If the same key is specified + multiple times, the outcome is undefined. + + BPS and IOPS are measured in each IO direction and IOs are + delayed if limit is reached. Temporary bursts are allowed. + + Setting read limit at 2M BPS and write at 120 IOPS for 8:16. + + echo "8:16 rbps=2097152 wiops=120" > io.max + + Reading returns the following. + + 8:16 rbps=2097152 wbps=max riops=max wiops=120 + + Write IOPS limit can be removed by writing the following. + + echo "8:16 wiops=max" > io.max + + Reading now returns the following. + + 8:16 rbps=2097152 wbps=max riops=max wiops=max + + +5-3-2. Writeback + +Page cache is dirtied through buffered writes and shared mmaps and +written asynchronously to the backing filesystem by the writeback +mechanism. Writeback sits between the memory and IO domains and +regulates the proportion of dirty memory by balancing dirtying and +write IOs. + +The io controller, in conjunction with the memory controller, +implements control of page cache writeback IOs. The memory controller +defines the memory domain that dirty memory ratio is calculated and +maintained for and the io controller defines the io domain which +writes out dirty pages for the memory domain. Both system-wide and +per-cgroup dirty memory states are examined and the more restrictive +of the two is enforced. + +cgroup writeback requires explicit support from the underlying +filesystem. Currently, cgroup writeback is implemented on ext2, ext4 +and btrfs. On other filesystems, all writeback IOs are attributed to +the root cgroup. + +There are inherent differences in memory and writeback management +which affects how cgroup ownership is tracked. Memory is tracked per +page while writeback per inode. For the purpose of writeback, an +inode is assigned to a cgroup and all IO requests to write dirty pages +from the inode are attributed to that cgroup. + +As cgroup ownership for memory is tracked per page, there can be pages +which are associated with different cgroups than the one the inode is +associated with. These are called foreign pages. The writeback +constantly keeps track of foreign pages and, if a particular foreign +cgroup becomes the majority over a certain period of time, switches +the ownership of the inode to that cgroup. + +While this model is enough for most use cases where a given inode is +mostly dirtied by a single cgroup even when the main writing cgroup +changes over time, use cases where multiple cgroups write to a single +inode simultaneously are not supported well. In such circumstances, a +significant portion of IOs are likely to be attributed incorrectly. +As memory controller assigns page ownership on the first use and +doesn't update it until the page is released, even if writeback +strictly follows page ownership, multiple cgroups dirtying overlapping +areas wouldn't work as expected. It's recommended to avoid such usage +patterns. + +The sysctl knobs which affect writeback behavior are applied to cgroup +writeback as follows. + + vm.dirty_background_ratio + vm.dirty_ratio + + These ratios apply the same to cgroup writeback with the + amount of available memory capped by limits imposed by the + memory controller and system-wide clean memory. + + vm.dirty_background_bytes + vm.dirty_bytes + + For cgroup writeback, this is calculated into ratio against + total available memory and applied the same way as + vm.dirty[_background]_ratio. + + +P. Information on Kernel Programming + +This section contains kernel programming information in the areas +where interacting with cgroup is necessary. cgroup core and +controllers are not covered. + + +P-1. Filesystem Support for Writeback + +A filesystem can support cgroup writeback by updating +address_space_operations->writepage[s]() to annotate bio's using the +following two functions. + + wbc_init_bio(@wbc, @bio) + + Should be called for each bio carrying writeback data and + associates the bio with the inode's owner cgroup. Can be + called anytime between bio allocation and submission. + + wbc_account_io(@wbc, @page, @bytes) + + Should be called for each data segment being written out. + While this function doesn't care exactly when it's called + during the writeback session, it's the easiest and most + natural to call it as data segments are added to a bio. + +With writeback bio's annotated, cgroup support can be enabled per +super_block by setting SB_I_CGROUPWB in ->s_iflags. This allows for +selective disabling of cgroup writeback support which is helpful when +certain filesystem features, e.g. journaled data mode, are +incompatible. + +wbc_init_bio() binds the specified bio to its cgroup. Depending on +the configuration, the bio may be executed at a lower priority and if +the writeback session is holding shared resources, e.g. a journal +entry, may lead to priority inversion. There is no one easy solution +for the problem. Filesystems can try to work around specific problem +cases by skipping wbc_init_bio() or using bio_associate_blkcg() +directly. + + +D. Deprecated v1 Core Features + +- Multiple hierarchies including named ones are not supported. + +- All mount options and remounting are not supported. + +- The "tasks" file is removed and "cgroup.procs" is not sorted. + +- "cgroup.clone_children" is removed. + +- /proc/cgroups is meaningless for v2. Use "cgroup.controllers" file + at the root instead. + + +R. Issues with v1 and Rationales for v2 + +R-1. Multiple Hierarchies + +cgroup v1 allowed an arbitrary number of hierarchies and each +hierarchy could host any number of controllers. While this seemed to +provide a high level of flexibility, it wasn't useful in practice. + +For example, as there is only one instance of each controller, utility +type controllers such as freezer which can be useful in all +hierarchies could only be used in one. The issue is exacerbated by +the fact that controllers couldn't be moved to another hierarchy once +hierarchies were populated. Another issue was that all controllers +bound to a hierarchy were forced to have exactly the same view of the +hierarchy. It wasn't possible to vary the granularity depending on +the specific controller. + +In practice, these issues heavily limited which controllers could be +put on the same hierarchy and most configurations resorted to putting +each controller on its own hierarchy. Only closely related ones, such +as the cpu and cpuacct controllers, made sense to be put on the same +hierarchy. This often meant that userland ended up managing multiple +similar hierarchies repeating the same steps on each hierarchy +whenever a hierarchy management operation was necessary. + +Furthermore, support for multiple hierarchies came at a steep cost. +It greatly complicated cgroup core implementation but more importantly +the support for multiple hierarchies restricted how cgroup could be +used in general and what controllers was able to do. + +There was no limit on how many hierarchies there might be, which meant +that a thread's cgroup membership couldn't be described in finite +length. The key might contain any number of entries and was unlimited +in length, which made it highly awkward to manipulate and led to +addition of controllers which existed only to identify membership, +which in turn exacerbated the original problem of proliferating number +of hierarchies. + +Also, as a controller couldn't have any expectation regarding the +topologies of hierarchies other controllers might be on, each +controller had to assume that all other controllers were attached to +completely orthogonal hierarchies. This made it impossible, or at +least very cumbersome, for controllers to cooperate with each other. + +In most use cases, putting controllers on hierarchies which are +completely orthogonal to each other isn't necessary. What usually is +called for is the ability to have differing levels of granularity +depending on the specific controller. In other words, hierarchy may +be collapsed from leaf towards root when viewed from specific +controllers. For example, a given configuration might not care about +how memory is distributed beyond a certain level while still wanting +to control how CPU cycles are distributed. + + +R-2. Thread Granularity + +cgroup v1 allowed threads of a process to belong to different cgroups. +This didn't make sense for some controllers and those controllers +ended up implementing different ways to ignore such situations but +much more importantly it blurred the line between API exposed to +individual applications and system management interface. + +Generally, in-process knowledge is available only to the process +itself; thus, unlike service-level organization of processes, +categorizing threads of a process requires active participation from +the application which owns the target process. + +cgroup v1 had an ambiguously defined delegation model which got abused +in combination with thread granularity. cgroups were delegated to +individual applications so that they can create and manage their own +sub-hierarchies and control resource distributions along them. This +effectively raised cgroup to the status of a syscall-like API exposed +to lay programs. + +First of all, cgroup has a fundamentally inadequate interface to be +exposed this way. For a process to access its own knobs, it has to +extract the path on the target hierarchy from /proc/self/cgroup, +construct the path by appending the name of the knob to the path, open +and then read and/or write to it. This is not only extremely clunky +and unusual but also inherently racy. There is no conventional way to +define transaction across the required steps and nothing can guarantee +that the process would actually be operating on its own sub-hierarchy. + +cgroup controllers implemented a number of knobs which would never be +accepted as public APIs because they were just adding control knobs to +system-management pseudo filesystem. cgroup ended up with interface +knobs which were not properly abstracted or refined and directly +revealed kernel internal details. These knobs got exposed to +individual applications through the ill-defined delegation mechanism +effectively abusing cgroup as a shortcut to implementing public APIs +without going through the required scrutiny. + +This was painful for both userland and kernel. Userland ended up with +misbehaving and poorly abstracted interfaces and kernel exposing and +locked into constructs inadvertently. + + +R-3. Competition Between Inner Nodes and Threads + +cgroup v1 allowed threads to be in any cgroups which created an +interesting problem where threads belonging to a parent cgroup and its +children cgroups competed for resources. This was nasty as two +different types of entities competed and there was no obvious way to +settle it. Different controllers did different things. + +The cpu controller considered threads and cgroups as equivalents and +mapped nice levels to cgroup weights. This worked for some cases but +fell flat when children wanted to be allocated specific ratios of CPU +cycles and the number of internal threads fluctuated - the ratios +constantly changed as the number of competing entities fluctuated. +There also were other issues. The mapping from nice level to weight +wasn't obvious or universal, and there were various other knobs which +simply weren't available for threads. + +The io controller implicitly created a hidden leaf node for each +cgroup to host the threads. The hidden leaf had its own copies of all +the knobs with "leaf_" prefixed. While this allowed equivalent +control over internal threads, it was with serious drawbacks. It +always added an extra layer of nesting which wouldn't be necessary +otherwise, made the interface messy and significantly complicated the +implementation. + +The memory controller didn't have a way to control what happened +between internal tasks and child cgroups and the behavior was not +clearly defined. There were attempts to add ad-hoc behaviors and +knobs to tailor the behavior to specific workloads which would have +led to problems extremely difficult to resolve in the long term. + +Multiple controllers struggled with internal tasks and came up with +different ways to deal with it; unfortunately, all the approaches were +severely flawed and, furthermore, the widely different behaviors +made cgroup as a whole highly inconsistent. + +This clearly is a problem which needs to be addressed from cgroup core +in a uniform way. + + +R-4. Other Interface Issues + +cgroup v1 grew without oversight and developed a large number of +idiosyncrasies and inconsistencies. One issue on the cgroup core side +was how an empty cgroup was notified - a userland helper binary was +forked and executed for each event. The event delivery wasn't +recursive or delegatable. The limitations of the mechanism also led +to in-kernel event delivery filtering mechanism further complicating +the interface. + +Controller interfaces were problematic too. An extreme example is +controllers completely ignoring hierarchical organization and treating +all cgroups as if they were all located directly under the root +cgroup. Some controllers exposed a large amount of inconsistent +implementation details to userland. + +There also was no consistency across controllers. When a new cgroup +was created, some controllers defaulted to not imposing extra +restrictions while others disallowed any resource usage until +explicitly configured. Configuration knobs for the same type of +control used widely differing naming schemes and formats. Statistics +and information knobs were named arbitrarily and used different +formats and units even in the same controller. + +cgroup v2 establishes common conventions where appropriate and updates +controllers so that they expose minimal and consistent interfaces. + + +R-5. Controller Issues and Remedies + +R-5-1. Memory + +The original lower boundary, the soft limit, is defined as a limit +that is per default unset. As a result, the set of cgroups that +global reclaim prefers is opt-in, rather than opt-out. The costs for +optimizing these mostly negative lookups are so high that the +implementation, despite its enormous size, does not even provide the +basic desirable behavior. First off, the soft limit has no +hierarchical meaning. All configured groups are organized in a global +rbtree and treated like equal peers, regardless where they are located +in the hierarchy. This makes subtree delegation impossible. Second, +the soft limit reclaim pass is so aggressive that it not just +introduces high allocation latencies into the system, but also impacts +system performance due to overreclaim, to the point where the feature +becomes self-defeating. + +The memory.low boundary on the other hand is a top-down allocated +reserve. A cgroup enjoys reclaim protection when it and all its +ancestors are below their low boundaries, which makes delegation of +subtrees possible. Secondly, new cgroups have no reserve per default +and in the common case most cgroups are eligible for the preferred +reclaim pass. This allows the new low boundary to be efficiently +implemented with just a minor addition to the generic reclaim code, +without the need for out-of-band data structures and reclaim passes. +Because the generic reclaim code considers all cgroups except for the +ones running low in the preferred first reclaim pass, overreclaim of +individual groups is eliminated as well, resulting in much better +overall workload performance. + +The original high boundary, the hard limit, is defined as a strict +limit that can not budge, even if the OOM killer has to be called. +But this generally goes against the goal of making the most out of the +available memory. The memory consumption of workloads varies during +runtime, and that requires users to overcommit. But doing that with a +strict upper limit requires either a fairly accurate prediction of the +working set size or adding slack to the limit. Since working set size +estimation is hard and error prone, and getting it wrong results in +OOM kills, most users tend to err on the side of a looser limit and +end up wasting precious resources. + +The memory.high boundary on the other hand can be set much more +conservatively. When hit, it throttles allocations by forcing them +into direct reclaim to work off the excess, but it never invokes the +OOM killer. As a result, a high boundary that is chosen too +aggressively will not terminate the processes, but instead it will +lead to gradual performance degradation. The user can monitor this +and make corrections until the minimal memory footprint that still +gives acceptable performance is found. + +In extreme cases, with many concurrent allocations and a complete +breakdown of reclaim progress within the group, the high boundary can +be exceeded. But even then it's mostly better to satisfy the +allocation from the slack available in other groups or the rest of the +system than killing the group. Otherwise, memory.max is there to +limit this type of spillover and ultimately contain buggy or even +malicious applications. diff --git a/Documentation/cgroups/unified-hierarchy.txt b/Documentation/cgroups/unified-hierarchy.txt deleted file mode 100644 index 781b1d475bcf..000000000000 --- a/Documentation/cgroups/unified-hierarchy.txt +++ /dev/null @@ -1,647 +0,0 @@ - -Cgroup unified hierarchy - -April, 2014 Tejun Heo <tj@kernel.org> - -This document describes the changes made by unified hierarchy and -their rationales. It will eventually be merged into the main cgroup -documentation. - -CONTENTS - -1. Background -2. Basic Operation - 2-1. Mounting - 2-2. cgroup.subtree_control - 2-3. cgroup.controllers -3. Structural Constraints - 3-1. Top-down - 3-2. No internal tasks -4. Delegation - 4-1. Model of delegation - 4-2. Common ancestor rule -5. Other Changes - 5-1. [Un]populated Notification - 5-2. Other Core Changes - 5-3. Controller File Conventions - 5-3-1. Format - 5-3-2. Control Knobs - 5-4. Per-Controller Changes - 5-4-1. io - 5-4-2. cpuset - 5-4-3. memory -6. Planned Changes - 6-1. CAP for resource control - - -1. Background - -cgroup allows an arbitrary number of hierarchies and each hierarchy -can host any number of controllers. While this seems to provide a -high level of flexibility, it isn't quite useful in practice. - -For example, as there is only one instance of each controller, utility -type controllers such as freezer which can be useful in all -hierarchies can only be used in one. The issue is exacerbated by the -fact that controllers can't be moved around once hierarchies are -populated. Another issue is that all controllers bound to a hierarchy -are forced to have exactly the same view of the hierarchy. It isn't -possible to vary the granularity depending on the specific controller. - -In practice, these issues heavily limit which controllers can be put -on the same hierarchy and most configurations resort to putting each -controller on its own hierarchy. Only closely related ones, such as -the cpu and cpuacct controllers, make sense to put on the same -hierarchy. This often means that userland ends up managing multiple -similar hierarchies repeating the same steps on each hierarchy -whenever a hierarchy management operation is necessary. - -Unfortunately, support for multiple hierarchies comes at a steep cost. -Internal implementation in cgroup core proper is dazzlingly -complicated but more importantly the support for multiple hierarchies -restricts how cgroup is used in general and what controllers can do. - -There's no limit on how many hierarchies there may be, which means -that a task's cgroup membership can't be described in finite length. -The key may contain any varying number of entries and is unlimited in -length, which makes it highly awkward to handle and leads to addition -of controllers which exist only to identify membership, which in turn -exacerbates the original problem. - -Also, as a controller can't have any expectation regarding what shape -of hierarchies other controllers would be on, each controller has to -assume that all other controllers are operating on completely -orthogonal hierarchies. This makes it impossible, or at least very -cumbersome, for controllers to cooperate with each other. - -In most use cases, putting controllers on hierarchies which are -completely orthogonal to each other isn't necessary. What usually is -called for is the ability to have differing levels of granularity -depending on the specific controller. In other words, hierarchy may -be collapsed from leaf towards root when viewed from specific -controllers. For example, a given configuration might not care about -how memory is distributed beyond a certain level while still wanting -to control how CPU cycles are distributed. - -Unified hierarchy is the next version of cgroup interface. It aims to -address the aforementioned issues by having more structure while -retaining enough flexibility for most use cases. Various other -general and controller-specific interface issues are also addressed in -the process. - - -2. Basic Operation - -2-1. Mounting - -Currently, unified hierarchy can be mounted with the following mount -command. Note that this is still under development and scheduled to -change soon. - - mount -t cgroup -o __DEVEL__sane_behavior cgroup $MOUNT_POINT - -All controllers which support the unified hierarchy and are not bound -to other hierarchies are automatically bound to unified hierarchy and -show up at the root of it. Controllers which are enabled only in the -root of unified hierarchy can be bound to other hierarchies. This -allows mixing unified hierarchy with the traditional multiple -hierarchies in a fully backward compatible way. - -A controller can be moved across hierarchies only after the controller -is no longer referenced in its current hierarchy. Because per-cgroup -controller states are destroyed asynchronously and controllers may -have lingering references, a controller may not show up immediately on -the unified hierarchy after the final umount of the previous -hierarchy. Similarly, a controller should be fully disabled to be -moved out of the unified hierarchy and it may take some time for the -disabled controller to become available for other hierarchies; -furthermore, due to dependencies among controllers, other controllers -may need to be disabled too. - -While useful for development and manual configurations, dynamically -moving controllers between the unified and other hierarchies is -strongly discouraged for production use. It is recommended to decide -the hierarchies and controller associations before starting using the -controllers. - - -2-2. cgroup.subtree_control - -All cgroups on unified hierarchy have a "cgroup.subtree_control" file -which governs which controllers are enabled on the children of the -cgroup. Let's assume a hierarchy like the following. - - root - A - B - C - \ D - -root's "cgroup.subtree_control" file determines which controllers are -enabled on A. A's on B. B's on C and D. This coincides with the -fact that controllers on the immediate sub-level are used to -distribute the resources of the parent. In fact, it's natural to -assume that resource control knobs of a child belong to its parent. -Enabling a controller in a "cgroup.subtree_control" file declares that -distribution of the respective resources of the cgroup will be -controlled. Note that this means that controller enable states are -shared among siblings. - -When read, the file contains a space-separated list of currently -enabled controllers. A write to the file should contain a -space-separated list of controllers with '+' or '-' prefixed (without -the quotes). Controllers prefixed with '+' are enabled and '-' -disabled. If a controller is listed multiple times, the last entry -wins. The specific operations are executed atomically - either all -succeed or fail. - - -2-3. cgroup.controllers - -Read-only "cgroup.controllers" file contains a space-separated list of -controllers which can be enabled in the cgroup's -"cgroup.subtree_control" file. - -In the root cgroup, this lists controllers which are not bound to -other hierarchies and the content changes as controllers are bound to -and unbound from other hierarchies. - -In non-root cgroups, the content of this file equals that of the -parent's "cgroup.subtree_control" file as only controllers enabled -from the parent can be used in its children. - - -3. Structural Constraints - -3-1. Top-down - -As it doesn't make sense to nest control of an uncontrolled resource, -all non-root "cgroup.subtree_control" files can only contain -controllers which are enabled in the parent's "cgroup.subtree_control" -file. A controller can be enabled only if the parent has the -controller enabled and a controller can't be disabled if one or more -children have it enabled. - - -3-2. No internal tasks - -One long-standing issue that cgroup faces is the competition between -tasks belonging to the parent cgroup and its children cgroups. This -is inherently nasty as two different types of entities compete and -there is no agreed-upon obvious way to handle it. Different -controllers are doing different things. - -The cpu controller considers tasks and cgroups as equivalents and maps -nice levels to cgroup weights. This works for some cases but falls -flat when children should be allocated specific ratios of CPU cycles -and the number of internal tasks fluctuates - the ratios constantly -change as the number of competing entities fluctuates. There also are -other issues. The mapping from nice level to weight isn't obvious or -universal, and there are various other knobs which simply aren't -available for tasks. - -The io controller implicitly creates a hidden leaf node for each -cgroup to host the tasks. The hidden leaf has its own copies of all -the knobs with "leaf_" prefixed. While this allows equivalent control -over internal tasks, it's with serious drawbacks. It always adds an -extra layer of nesting which may not be necessary, makes the interface -messy and significantly complicates the implementation. - -The memory controller currently doesn't have a way to control what -happens between internal tasks and child cgroups and the behavior is -not clearly defined. There have been attempts to add ad-hoc behaviors -and knobs to tailor the behavior to specific workloads. Continuing -this direction will lead to problems which will be extremely difficult -to resolve in the long term. - -Multiple controllers struggle with internal tasks and came up with -different ways to deal with it; unfortunately, all the approaches in -use now are severely flawed and, furthermore, the widely different -behaviors make cgroup as whole highly inconsistent. - -It is clear that this is something which needs to be addressed from -cgroup core proper in a uniform way so that controllers don't need to -worry about it and cgroup as a whole shows a consistent and logical -behavior. To achieve that, unified hierarchy enforces the following -structural constraint: - - Except for the root, only cgroups which don't contain any task may - have controllers enabled in their "cgroup.subtree_control" files. - -Combined with other properties, this guarantees that, when a -controller is looking at the part of the hierarchy which has it -enabled, tasks are always only on the leaves. This rules out -situations where child cgroups compete against internal tasks of the -parent. - -There are two things to note. Firstly, the root cgroup is exempt from -the restriction. Root contains tasks and anonymous resource -consumption which can't be associated with any other cgroup and -requires special treatment from most controllers. How resource -consumption in the root cgroup is governed is up to each controller. - -Secondly, the restriction doesn't take effect if there is no enabled -controller in the cgroup's "cgroup.subtree_control" file. This is -important as otherwise it wouldn't be possible to create children of a -populated cgroup. To control resource distribution of a cgroup, the -cgroup must create children and transfer all its tasks to the children -before enabling controllers in its "cgroup.subtree_control" file. - - -4. Delegation - -4-1. Model of delegation - -A cgroup can be delegated to a less privileged user by granting write -access of the directory and its "cgroup.procs" file to the user. Note -that the resource control knobs in a given directory concern the -resources of the parent and thus must not be delegated along with the -directory. - -Once delegated, the user can build sub-hierarchy under the directory, -organize processes as it sees fit and further distribute the resources -it got from the parent. The limits and other settings of all resource -controllers are hierarchical and regardless of what happens in the -delegated sub-hierarchy, nothing can escape the resource restrictions -imposed by the parent. - -Currently, cgroup doesn't impose any restrictions on the number of -cgroups in or nesting depth of a delegated sub-hierarchy; however, -this may in the future be limited explicitly. - - -4-2. Common ancestor rule - -On the unified hierarchy, to write to a "cgroup.procs" file, in -addition to the usual write permission to the file and uid match, the -writer must also have write access to the "cgroup.procs" file of the -common ancestor of the source and destination cgroups. This prevents -delegatees from smuggling processes across disjoint sub-hierarchies. - -Let's say cgroups C0 and C1 have been delegated to user U0 who created -C00, C01 under C0 and C10 under C1 as follows. - - ~~~~~~~~~~~~~ - C0 - C00 - ~ cgroup ~ \ C01 - ~ hierarchy ~ - ~~~~~~~~~~~~~ - C1 - C10 - -C0 and C1 are separate entities in terms of resource distribution -regardless of their relative positions in the hierarchy. The -resources the processes under C0 are entitled to are controlled by -C0's ancestors and may be completely different from C1. It's clear -that the intention of delegating C0 to U0 is allowing U0 to organize -the processes under C0 and further control the distribution of C0's -resources. - -On traditional hierarchies, if a task has write access to "tasks" or -"cgroup.procs" file of a cgroup and its uid agrees with the target, it -can move the target to the cgroup. In the above example, U0 will not -only be able to move processes in each sub-hierarchy but also across -the two sub-hierarchies, effectively allowing it to violate the -organizational and resource restrictions implied by the hierarchical -structure above C0 and C1. - -On the unified hierarchy, let's say U0 wants to write the pid of a -process which has a matching uid and is currently in C10 into -"C00/cgroup.procs". U0 obviously has write access to the file and -migration permission on the process; however, the common ancestor of -the source cgroup C10 and the destination cgroup C00 is above the -points of delegation and U0 would not have write access to its -"cgroup.procs" and thus be denied with -EACCES. - - -5. Other Changes - -5-1. [Un]populated Notification - -cgroup users often need a way to determine when a cgroup's -subhierarchy becomes empty so that it can be cleaned up. cgroup -currently provides release_agent for it; unfortunately, this mechanism -is riddled with issues. - -- It delivers events by forking and execing a userland binary - specified as the release_agent. This is a long deprecated method of - notification delivery. It's extremely heavy, slow and cumbersome to - integrate with larger infrastructure. - -- There is single monitoring point at the root. There's no way to - delegate management of a subtree. - -- The event isn't recursive. It triggers when a cgroup doesn't have - any tasks or child cgroups. Events for internal nodes trigger only - after all children are removed. This again makes it impossible to - delegate management of a subtree. - -- Events are filtered from the kernel side. A "notify_on_release" - file is used to subscribe to or suppress release events. This is - unnecessarily complicated and probably done this way because event - delivery itself was expensive. - -Unified hierarchy implements "populated" field in "cgroup.events" -interface file which can be used to monitor whether the cgroup's -subhierarchy has tasks in it or not. Its value is 0 if there is no -task in the cgroup and its descendants; otherwise, 1. poll and -[id]notify events are triggered when the value changes. - -This is significantly lighter and simpler and trivially allows -delegating management of subhierarchy - subhierarchy monitoring can -block further propagation simply by putting itself or another process -in the subhierarchy and monitor events that it's interested in from -there without interfering with monitoring higher in the tree. - -In unified hierarchy, the release_agent mechanism is no longer -supported and the interface files "release_agent" and -"notify_on_release" do not exist. - - -5-2. Other Core Changes - -- None of the mount options is allowed. - -- remount is disallowed. - -- rename(2) is disallowed. - -- The "tasks" file is removed. Everything should at process - granularity. Use the "cgroup.procs" file instead. - -- The "cgroup.procs" file is not sorted. pids will be unique unless - they got recycled in-between reads. - -- The "cgroup.clone_children" file is removed. - -- /proc/PID/cgroup keeps reporting the cgroup that a zombie belonged - to before exiting. If the cgroup is removed before the zombie is - reaped, " (deleted)" is appeneded to the path. - - -5-3. Controller File Conventions - -5-3-1. Format - -In general, all controller files should be in one of the following -formats whenever possible. - -- Values only files - - VAL0 VAL1...\n - -- Flat keyed files - - KEY0 VAL0\n - KEY1 VAL1\n - ... - -- Nested keyed files - - KEY0 SUB_KEY0=VAL00 SUB_KEY1=VAL01... - KEY1 SUB_KEY0=VAL10 SUB_KEY1=VAL11... - ... - -For a writeable file, the format for writing should generally match -reading; however, controllers may allow omitting later fields or -implement restricted shortcuts for most common use cases. - -For both flat and nested keyed files, only the values for a single key -can be written at a time. For nested keyed files, the sub key pairs -may be specified in any order and not all pairs have to be specified. - - -5-3-2. Control Knobs - -- Settings for a single feature should generally be implemented in a - single file. - -- In general, the root cgroup should be exempt from resource control - and thus shouldn't have resource control knobs. - -- If a controller implements ratio based resource distribution, the - control knob should be named "weight" and have the range [1, 10000] - and 100 should be the default value. The values are chosen to allow - enough and symmetric bias in both directions while keeping it - intuitive (the default is 100%). - -- If a controller implements an absolute resource guarantee and/or - limit, the control knobs should be named "min" and "max" - respectively. If a controller implements best effort resource - gurantee and/or limit, the control knobs should be named "low" and - "high" respectively. - - In the above four control files, the special token "max" should be - used to represent upward infinity for both reading and writing. - -- If a setting has configurable default value and specific overrides, - the default settings should be keyed with "default" and appear as - the first entry in the file. Specific entries can use "default" as - its value to indicate inheritance of the default value. - -- For events which are not very high frequency, an interface file - "events" should be created which lists event key value pairs. - Whenever a notifiable event happens, file modified event should be - generated on the file. - - -5-4. Per-Controller Changes - -5-4-1. io - -- blkio is renamed to io. The interface is overhauled anyway. The - new name is more in line with the other two major controllers, cpu - and memory, and better suited given that it may be used for cgroup - writeback without involving block layer. - -- Everything including stat is always hierarchical making separate - recursive stat files pointless and, as no internal node can have - tasks, leaf weights are meaningless. The operation model is - simplified and the interface is overhauled accordingly. - - io.stat - - The stat file. The reported stats are from the point where - bio's are issued to request_queue. The stats are counted - independent of which policies are enabled. Each line in the - file follows the following format. More fields may later be - added at the end. - - $MAJ:$MIN rbytes=$RBYTES wbytes=$WBYTES rios=$RIOS wrios=$WIOS - - io.weight - - The weight setting, currently only available and effective if - cfq-iosched is in use for the target device. The weight is - between 1 and 10000 and defaults to 100. The first line - always contains the default weight in the following format to - use when per-device setting is missing. - - default $WEIGHT - - Subsequent lines list per-device weights of the following - format. - - $MAJ:$MIN $WEIGHT - - Writing "$WEIGHT" or "default $WEIGHT" changes the default - setting. Writing "$MAJ:$MIN $WEIGHT" sets per-device weight - while "$MAJ:$MIN default" clears it. - - This file is available only on non-root cgroups. - - io.max - - The maximum bandwidth and/or iops setting, only available if - blk-throttle is enabled. The file is of the following format. - - $MAJ:$MIN rbps=$RBPS wbps=$WBPS riops=$RIOPS wiops=$WIOPS - - ${R|W}BPS are read/write bytes per second and ${R|W}IOPS are - read/write IOs per second. "max" indicates no limit. Writing - to the file follows the same format but the individual - settings may be omitted or specified in any order. - - This file is available only on non-root cgroups. - - -5-4-2. cpuset - -- Tasks are kept in empty cpusets after hotplug and take on the masks - of the nearest non-empty ancestor, instead of being moved to it. - -- A task can be moved into an empty cpuset, and again it takes on the - masks of the nearest non-empty ancestor. - - -5-4-3. memory - -- use_hierarchy is on by default and the cgroup file for the flag is - not created. - -- The original lower boundary, the soft limit, is defined as a limit - that is per default unset. As a result, the set of cgroups that - global reclaim prefers is opt-in, rather than opt-out. The costs - for optimizing these mostly negative lookups are so high that the - implementation, despite its enormous size, does not even provide the - basic desirable behavior. First off, the soft limit has no - hierarchical meaning. All configured groups are organized in a - global rbtree and treated like equal peers, regardless where they - are located in the hierarchy. This makes subtree delegation - impossible. Second, the soft limit reclaim pass is so aggressive - that it not just introduces high allocation latencies into the - system, but also impacts system performance due to overreclaim, to - the point where the feature becomes self-defeating. - - The memory.low boundary on the other hand is a top-down allocated - reserve. A cgroup enjoys reclaim protection when it and all its - ancestors are below their low boundaries, which makes delegation of - subtrees possible. Secondly, new cgroups have no reserve per - default and in the common case most cgroups are eligible for the - preferred reclaim pass. This allows the new low boundary to be - efficiently implemented with just a minor addition to the generic - reclaim code, without the need for out-of-band data structures and - reclaim passes. Because the generic reclaim code considers all - cgroups except for the ones running low in the preferred first - reclaim pass, overreclaim of individual groups is eliminated as - well, resulting in much better overall workload performance. - -- The original high boundary, the hard limit, is defined as a strict - limit that can not budge, even if the OOM killer has to be called. - But this generally goes against the goal of making the most out of - the available memory. The memory consumption of workloads varies - during runtime, and that requires users to overcommit. But doing - that with a strict upper limit requires either a fairly accurate - prediction of the working set size or adding slack to the limit. - Since working set size estimation is hard and error prone, and - getting it wrong results in OOM kills, most users tend to err on the - side of a looser limit and end up wasting precious resources. - - The memory.high boundary on the other hand can be set much more - conservatively. When hit, it throttles allocations by forcing them - into direct reclaim to work off the excess, but it never invokes the - OOM killer. As a result, a high boundary that is chosen too - aggressively will not terminate the processes, but instead it will - lead to gradual performance degradation. The user can monitor this - and make corrections until the minimal memory footprint that still - gives acceptable performance is found. - - In extreme cases, with many concurrent allocations and a complete - breakdown of reclaim progress within the group, the high boundary - can be exceeded. But even then it's mostly better to satisfy the - allocation from the slack available in other groups or the rest of - the system than killing the group. Otherwise, memory.max is there - to limit this type of spillover and ultimately contain buggy or even - malicious applications. - -- The original control file names are unwieldy and inconsistent in - many different ways. For example, the upper boundary hit count is - exported in the memory.failcnt file, but an OOM event count has to - be manually counted by listening to memory.oom_control events, and - lower boundary / soft limit events have to be counted by first - setting a threshold for that value and then counting those events. - Also, usage and limit files encode their units in the filename. - That makes the filenames very long, even though this is not - information that a user needs to be reminded of every time they type - out those names. - - To address these naming issues, as well as to signal clearly that - the new interface carries a new configuration model, the naming - conventions in it necessarily differ from the old interface. - -- The original limit files indicate the state of an unset limit with a - Very High Number, and a configured limit can be unset by echoing -1 - into those files. But that very high number is implementation and - architecture dependent and not very descriptive. And while -1 can - be understood as an underflow into the highest possible value, -2 or - -10M etc. do not work, so it's not consistent. - - memory.low, memory.high, and memory.max will use the string "max" to - indicate and set the highest possible value. - -6. Planned Changes - -6-1. CAP for resource control - -Unified hierarchy will require one of the capabilities(7), which is -yet to be decided, for all resource control related knobs. Process -organization operations - creation of sub-cgroups and migration of -processes in sub-hierarchies may be delegated by changing the -ownership and/or permissions on the cgroup directory and -"cgroup.procs" interface file; however, all operations which affect -resource control - writes to a "cgroup.subtree_control" file or any -controller-specific knobs - will require an explicit CAP privilege. - -This, in part, is to prevent the cgroup interface from being -inadvertently promoted to programmable API used by non-privileged -binaries. cgroup exposes various aspects of the system in ways which -aren't properly abstracted for direct consumption by regular programs. -This is an administration interface much closer to sysctl knobs than -system calls. Even the basic access model, being filesystem path -based, isn't suitable for direct consumption. There's no way to -access "my cgroup" in a race-free way or make multiple operations -atomic against migration to another cgroup. - -Another aspect is that, for better or for worse, the cgroup interface -goes through far less scrutiny than regular interfaces for -unprivileged userland. The upside is that cgroup is able to expose -useful features which may not be suitable for general consumption in a -reasonable time frame. It provides a relatively short path between -internal details and userland-visible interface. Of course, this -shortcut comes with high risk. We go through what we go through for -general kernel APIs for good reasons. It may end up leaking internal -details in a way which can exert significant pain by locking the -kernel into a contract that can't be maintained in a reasonable -manner. - -Also, due to the specific nature, cgroup and its controllers don't -tend to attract attention from a wide scope of developers. cgroup's -short history is already fraught with severely mis-designed -interfaces, unnecessary commitments to and exposing of internal -details, broken and dangerous implementations of various features. - -Keeping cgroup as an administration interface is both advantageous for -its role and imperative given its nature. Some of the cgroup features -may make sense for unprivileged access. If deemed justified, those -must be further abstracted and implemented as a different interface, -be it a system call or process-private filesystem, and survive through -the scrutiny that any interface for general consumption is required to -go through. - -Requiring CAP is not a complete solution but should serve as a -significant deterrent against spraying cgroup usages in non-privileged -programs. diff --git a/Documentation/cpu-freq/intel-pstate.txt b/Documentation/cpu-freq/intel-pstate.txt index be8d4006bf76..f7b12c071d53 100644 --- a/Documentation/cpu-freq/intel-pstate.txt +++ b/Documentation/cpu-freq/intel-pstate.txt @@ -1,61 +1,131 @@ -Intel P-state driver +Intel P-State driver -------------------- -This driver provides an interface to control the P state selection for -SandyBridge+ Intel processors. The driver can operate two different -modes based on the processor model, legacy mode and Hardware P state (HWP) -mode. - -In legacy mode, the Intel P-state implements two internal governors, -performance and powersave, that differ from the general cpufreq governors of -the same name (the general cpufreq governors implement target(), whereas the -internal Intel P-state governors implement setpolicy()). The internal -performance governor sets the max_perf_pct and min_perf_pct to 100; that is, -the governor selects the highest available P state to maximize the performance -of the core. The internal powersave governor selects the appropriate P state -based on the current load on the CPU. - -In HWP mode P state selection is implemented in the processor -itself. The driver provides the interfaces between the cpufreq core and -the processor to control P state selection based on user preferences -and reporting frequency to the cpufreq core. In this mode the -internal Intel P-state governor code is disabled. - -In addition to the interfaces provided by the cpufreq core for -controlling frequency the driver provides sysfs files for -controlling P state selection. These files have been added to -/sys/devices/system/cpu/intel_pstate/ - - max_perf_pct: limits the maximum P state that will be requested by - the driver stated as a percentage of the available performance. The - available (P states) performance may be reduced by the no_turbo +This driver provides an interface to control the P-State selection for the +SandyBridge+ Intel processors. + +The following document explains P-States: +http://events.linuxfoundation.org/sites/events/files/slides/LinuxConEurope_2015.pdf +As stated in the document, P-State doesn’t exactly mean a frequency. However, for +the sake of the relationship with cpufreq, P-State and frequency are used +interchangeably. + +Understanding the cpufreq core governors and policies are important before +discussing more details about the Intel P-State driver. Based on what callbacks +a cpufreq driver provides to the cpufreq core, it can support two types of +drivers: +- with target_index() callback: In this mode, the drivers using cpufreq core +simply provide the minimum and maximum frequency limits and an additional +interface target_index() to set the current frequency. The cpufreq subsystem +has a number of scaling governors ("performance", "powersave", "ondemand", +etc.). Depending on which governor is in use, cpufreq core will call for +transitions to a specific frequency using target_index() callback. +- setpolicy() callback: In this mode, drivers do not provide target_index() +callback, so cpufreq core can't request a transition to a specific frequency. +The driver provides minimum and maximum frequency limits and callbacks to set a +policy. The policy in cpufreq sysfs is referred to as the "scaling governor". +The cpufreq core can request the driver to operate in any of the two policies: +"performance: and "powersave". The driver decides which frequency to use based +on the above policy selection considering minimum and maximum frequency limits. + +The Intel P-State driver falls under the latter category, which implements the +setpolicy() callback. This driver decides what P-State to use based on the +requested policy from the cpufreq core. If the processor is capable of +selecting its next P-State internally, then the driver will offload this +responsibility to the processor (aka HWP: Hardware P-States). If not, the +driver implements algorithms to select the next P-State. + +Since these policies are implemented in the driver, they are not same as the +cpufreq scaling governors implementation, even if they have the same name in +the cpufreq sysfs (scaling_governors). For example the "performance" policy is +similar to cpufreq’s "performance" governor, but "powersave" is completely +different than the cpufreq "powersave" governor. The strategy here is similar +to cpufreq "ondemand", where the requested P-State is related to the system load. + +Sysfs Interface + +In addition to the frequency-controlling interfaces provided by the cpufreq +core, the driver provides its own sysfs files to control the P-State selection. +These files have been added to /sys/devices/system/cpu/intel_pstate/. +Any changes made to these files are applicable to all CPUs (even in a +multi-package system). + + max_perf_pct: Limits the maximum P-State that will be requested by + the driver. It states it as a percentage of the available performance. The + available (P-State) performance may be reduced by the no_turbo setting described below. - min_perf_pct: limits the minimum P state that will be requested by - the driver stated as a percentage of the max (non-turbo) + min_perf_pct: Limits the minimum P-State that will be requested by + the driver. It states it as a percentage of the max (non-turbo) performance level. - no_turbo: limits the driver to selecting P states below the turbo + no_turbo: Limits the driver to selecting P-State below the turbo frequency range. - turbo_pct: displays the percentage of the total performance that - is supported by hardware that is in the turbo range. This number + turbo_pct: Displays the percentage of the total performance that + is supported by hardware that is in the turbo range. This number is independent of whether turbo has been disabled or not. - num_pstates: displays the number of pstates that are supported - by hardware. This number is independent of whether turbo has + num_pstates: Displays the number of P-States that are supported + by hardware. This number is independent of whether turbo has been disabled or not. +For example, if a system has these parameters: + Max 1 core turbo ratio: 0x21 (Max 1 core ratio is the maximum P-State) + Max non turbo ratio: 0x17 + Minimum ratio : 0x08 (Here the ratio is called max efficiency ratio) + +Sysfs will show : + max_perf_pct:100, which corresponds to 1 core ratio + min_perf_pct:24, max_efficiency_ratio / max 1 Core ratio + no_turbo:0, turbo is not disabled + num_pstates:26 = (max 1 Core ratio - Max Efficiency Ratio + 1) + turbo_pct:39 = (max 1 core ratio - max non turbo ratio) / num_pstates + +Refer to "Intel® 64 and IA-32 Architectures Software Developer’s Manual +Volume 3: System Programming Guide" to understand ratios. + +cpufreq sysfs for Intel P-State + +Since this driver registers with cpufreq, cpufreq sysfs is also presented. +There are some important differences, which need to be considered. + +scaling_cur_freq: This displays the real frequency which was used during +the last sample period instead of what is requested. Some other cpufreq driver, +like acpi-cpufreq, displays what is requested (Some changes are on the +way to fix this for acpi-cpufreq driver). The same is true for frequencies +displayed at /proc/cpuinfo. + +scaling_governor: This displays current active policy. Since each CPU has a +cpufreq sysfs, it is possible to set a scaling governor to each CPU. But this +is not possible with Intel P-States, as there is one common policy for all +CPUs. Here, the last requested policy will be applicable to all CPUs. It is +suggested that one use the cpupower utility to change policy to all CPUs at the +same time. + +scaling_setspeed: This attribute can never be used with Intel P-State. + +scaling_max_freq/scaling_min_freq: This interface can be used similarly to +the max_perf_pct/min_perf_pct of Intel P-State sysfs. However since frequencies +are converted to nearest possible P-State, this is prone to rounding errors. +This method is not preferred to limit performance. + +affected_cpus: Not used +related_cpus: Not used + For contemporary Intel processors, the frequency is controlled by the -processor itself and the P-states exposed to software are related to +processor itself and the P-State exposed to software is related to performance levels. The idea that frequency can be set to a single -frequency is fiction for Intel Core processors. Even if the scaling -driver selects a single P state the actual frequency the processor +frequency is fictional for Intel Core processors. Even if the scaling +driver selects a single P-State, the actual frequency the processor will run at is selected by the processor itself. -For legacy mode debugfs files have also been added to allow tuning of -the internal governor algorythm. These files are located at -/sys/kernel/debug/pstate_snb/ These files are NOT present in HWP mode. +Tuning Intel P-State driver + +When HWP mode is not used, debugfs files have also been added to allow the +tuning of the internal governor algorithm. These files are located at +/sys/kernel/debug/pstate_snb/. The algorithm uses a PID (Proportional +Integral Derivative) controller. The PID tunable parameters are: deadband d_gain_pct @@ -63,3 +133,90 @@ the internal governor algorythm. These files are located at p_gain_pct sample_rate_ms setpoint + +To adjust these parameters, some understanding of driver implementation is +necessary. There are some tweeks described here, but be very careful. Adjusting +them requires expert level understanding of power and performance relationship. +These limits are only useful when the "powersave" policy is active. + +-To make the system more responsive to load changes, sample_rate_ms can +be adjusted (current default is 10ms). +-To make the system use higher performance, even if the load is lower, setpoint +can be adjusted to a lower number. This will also lead to faster ramp up time +to reach the maximum P-State. +If there are no derivative and integral coefficients, The next P-State will be +equal to: + current P-State - ((setpoint - current cpu load) * p_gain_pct) + +For example, if the current PID parameters are (Which are defaults for the core +processors like SandyBridge): + deadband = 0 + d_gain_pct = 0 + i_gain_pct = 0 + p_gain_pct = 20 + sample_rate_ms = 10 + setpoint = 97 + +If the current P-State = 0x08 and current load = 100, this will result in the +next P-State = 0x08 - ((97 - 100) * 0.2) = 8.6 (rounded to 9). Here the P-State +goes up by only 1. If during next sample interval the current load doesn't +change and still 100, then P-State goes up by one again. This process will +continue as long as the load is more than the setpoint until the maximum P-State +is reached. + +For the same load at setpoint = 60, this will result in the next P-State += 0x08 - ((60 - 100) * 0.2) = 16 +So by changing the setpoint from 97 to 60, there is an increase of the +next P-State from 9 to 16. So this will make processor execute at higher +P-State for the same CPU load. If the load continues to be more than the +setpoint during next sample intervals, then P-State will go up again till the +maximum P-State is reached. But the ramp up time to reach the maximum P-State +will be much faster when the setpoint is 60 compared to 97. + +Debugging Intel P-State driver + +Event tracing +To debug P-State transition, the Linux event tracing interface can be used. +There are two specific events, which can be enabled (Provided the kernel +configs related to event tracing are enabled). + +# cd /sys/kernel/debug/tracing/ +# echo 1 > events/power/pstate_sample/enable +# echo 1 > events/power/cpu_frequency/enable +# cat trace +gnome-terminal--4510 [001] ..s. 1177.680733: pstate_sample: core_busy=107 + scaled=94 from=26 to=26 mperf=1143818 aperf=1230607 tsc=29838618 + freq=2474476 +cat-5235 [002] ..s. 1177.681723: cpu_frequency: state=2900000 cpu_id=2 + + +Using ftrace + +If function level tracing is required, the Linux ftrace interface can be used. +For example if we want to check how often a function to set a P-State is +called, we can set ftrace filter to intel_pstate_set_pstate. + +# cd /sys/kernel/debug/tracing/ +# cat available_filter_functions | grep -i pstate +intel_pstate_set_pstate +intel_pstate_cpu_init +... + +# echo intel_pstate_set_pstate > set_ftrace_filter +# echo function > current_tracer +# cat trace | head -15 +# tracer: function +# +# entries-in-buffer/entries-written: 80/80 #P:4 +# +# _-----=> irqs-off +# / _----=> need-resched +# | / _---=> hardirq/softirq +# || / _--=> preempt-depth +# ||| / delay +# TASK-PID CPU# |||| TIMESTAMP FUNCTION +# | | | |||| | | + Xorg-3129 [000] ..s. 2537.644844: intel_pstate_set_pstate <-intel_pstate_timer_func + gnome-terminal--4510 [002] ..s. 2537.649844: intel_pstate_set_pstate <-intel_pstate_timer_func + gnome-shell-3409 [001] ..s. 2537.650850: intel_pstate_set_pstate <-intel_pstate_timer_func + <idle>-0 [000] ..s. 2537.654843: intel_pstate_set_pstate <-intel_pstate_timer_func diff --git a/Documentation/cpu-freq/pcc-cpufreq.txt b/Documentation/cpu-freq/pcc-cpufreq.txt index 9e3c3b33514c..0a94224ad296 100644 --- a/Documentation/cpu-freq/pcc-cpufreq.txt +++ b/Documentation/cpu-freq/pcc-cpufreq.txt @@ -159,8 +159,8 @@ to be strictly associated with a P-state. 2.2 cpuinfo_transition_latency: ------------------------------- -The cpuinfo_transition_latency field is 0. The PCC specification does -not include a field to expose this value currently. +The cpuinfo_transition_latency field is CPUFREQ_ETERNAL. The PCC specification +does not include a field to expose this value currently. 2.3 cpuinfo_cur_freq: --------------------- diff --git a/Documentation/device-mapper/verity.txt b/Documentation/device-mapper/verity.txt index e15bc1a0fb98..89fd8f9a259f 100644 --- a/Documentation/device-mapper/verity.txt +++ b/Documentation/device-mapper/verity.txt @@ -18,11 +18,11 @@ Construction Parameters 0 is the original format used in the Chromium OS. The salt is appended when hashing, digests are stored continuously and - the rest of the block is padded with zeros. + the rest of the block is padded with zeroes. 1 is the current format that should be used for new devices. The salt is prepended when hashing and each digest is - padded with zeros to the power of two. + padded with zeroes to the power of two. <dev> This is the device containing data, the integrity of which needs to be @@ -79,6 +79,37 @@ restart_on_corruption not compatible with ignore_corruption and requires user space support to avoid restart loops. +ignore_zero_blocks + Do not verify blocks that are expected to contain zeroes and always return + zeroes instead. This may be useful if the partition contains unused blocks + that are not guaranteed to contain zeroes. + +use_fec_from_device <fec_dev> + Use forward error correction (FEC) to recover from corruption if hash + verification fails. Use encoding data from the specified device. This + may be the same device where data and hash blocks reside, in which case + fec_start must be outside data and hash areas. + + If the encoding data covers additional metadata, it must be accessible + on the hash device after the hash blocks. + + Note: block sizes for data and hash devices must match. Also, if the + verity <dev> is encrypted the <fec_dev> should be too. + +fec_roots <num> + Number of generator roots. This equals to the number of parity bytes in + the encoding data. For example, in RS(M, N) encoding, the number of roots + is M-N. + +fec_blocks <num> + The number of encoding data blocks on the FEC device. The block size for + the FEC device is <data_block_size>. + +fec_start <offset> + This is the offset, in <data_block_size> blocks, from the start of the + FEC device to the beginning of the encoding data. + + Theory of operation =================== @@ -98,6 +129,11 @@ per-block basis. This allows for a lightweight hash computation on first read into the page cache. Block hashes are stored linearly, aligned to the nearest block size. +If forward error correction (FEC) support is enabled any recovery of +corrupted data will be verified using the cryptographic hash of the +corresponding data. This is why combining error correction with +integrity checking is essential. + Hash Tree --------- diff --git a/Documentation/devicetree/bindings/arm/cpus.txt b/Documentation/devicetree/bindings/arm/cpus.txt index 3a07a87fef20..6aca64f289b6 100644 --- a/Documentation/devicetree/bindings/arm/cpus.txt +++ b/Documentation/devicetree/bindings/arm/cpus.txt @@ -242,6 +242,23 @@ nodes to be present and contain the properties described below. Definition: Specifies the syscon node controlling the cpu core power domains. + - dynamic-power-coefficient + Usage: optional + Value type: <prop-encoded-array> + Definition: A u32 value that represents the running time dynamic + power coefficient in units of mW/MHz/uVolt^2. The + coefficient can either be calculated from power + measurements or derived by analysis. + + The dynamic power consumption of the CPU is + proportional to the square of the Voltage (V) and + the clock frequency (f). The coefficient is used to + calculate the dynamic power as below - + + Pdyn = dynamic-power-coefficient * V^2 * f + + where voltage is in uV, frequency is in MHz. + Example 1 (dual-cluster big.LITTLE system 32-bit): cpus { diff --git a/Documentation/devicetree/bindings/arm/l2cc.txt b/Documentation/devicetree/bindings/arm/l2c2x0.txt index 06c88a4d28ac..fe0398c5c77b 100644 --- a/Documentation/devicetree/bindings/arm/l2cc.txt +++ b/Documentation/devicetree/bindings/arm/l2c2x0.txt @@ -1,7 +1,8 @@ * ARM L2 Cache Controller -ARM cores often have a separate level 2 cache controller. There are various -implementations of the L2 cache controller with compatible programming models. +ARM cores often have a separate L2C210/L2C220/L2C310 (also known as PL210/PL220/ +PL310 and variants) based level 2 cache controller. All these various implementations +of the L2 cache controller have compatible programming models (Note 1). Some of the properties that are just prefixed "cache-*" are taken from section 3.7.3 of the ePAPR v1.1 specification which can be found at: https://www.power.org/wp-content/uploads/2012/06/Power_ePAPR_APPROVED_v1.1.pdf @@ -67,12 +68,17 @@ Optional properties: disable if zero. - arm,prefetch-offset : Override prefetch offset value. Valid values are 0-7, 15, 23, and 31. -- arm,shared-override : The default behavior of the pl310 cache controller with - respect to the shareable attribute is to transform "normal memory - non-cacheable transactions" into "cacheable no allocate" (for reads) or - "write through no write allocate" (for writes). +- arm,shared-override : The default behavior of the L220 or PL310 cache + controllers with respect to the shareable attribute is to transform "normal + memory non-cacheable transactions" into "cacheable no allocate" (for reads) + or "write through no write allocate" (for writes). On systems where this may cause DMA buffer corruption, this property must be specified to indicate that such transforms are precluded. +- arm,parity-enable : enable parity checking on the L2 cache (L220 or PL310). +- arm,parity-disable : disable parity checking on the L2 cache (L220 or PL310). +- arm,outer-sync-disable : disable the outer sync operation on the L2 cache. + Some core tiles, especially ARM PB11MPCore have a faulty L220 cache that + will randomly hang unless outer sync operations are disabled. - prefetch-data : Data prefetch. Value: <0> (forcibly disable), <1> (forcibly enable), property absent (retain settings set by firmware) - prefetch-instr : Instruction prefetch. Value: <0> (forcibly disable), @@ -91,3 +97,9 @@ L2: cache-controller { cache-level = <2>; interrupts = <45>; }; + +Note 1: The description in this document doesn't apply to integrated L2 + cache controllers as found in e.g. Cortex-A15/A7/A57/A53. These + integrated L2 controllers are assumed to be all preconfigured by + early secure boot code. Thus no need to deal with their configuration + in the kernel at all. diff --git a/Documentation/devicetree/bindings/arm/pmu.txt b/Documentation/devicetree/bindings/arm/pmu.txt index 97ba45af04fc..56518839f52a 100644 --- a/Documentation/devicetree/bindings/arm/pmu.txt +++ b/Documentation/devicetree/bindings/arm/pmu.txt @@ -9,8 +9,9 @@ Required properties: - compatible : should be one of "apm,potenza-pmu" "arm,armv8-pmuv3" - "arm.cortex-a57-pmu" - "arm.cortex-a53-pmu" + "arm,cortex-a72-pmu" + "arm,cortex-a57-pmu" + "arm,cortex-a53-pmu" "arm,cortex-a17-pmu" "arm,cortex-a15-pmu" "arm,cortex-a12-pmu" diff --git a/Documentation/devicetree/bindings/ata/brcm,sata-brcmstb.txt b/Documentation/devicetree/bindings/ata/brcm,sata-brcmstb.txt index 20ac9bbfa1fd..60872838f1ad 100644 --- a/Documentation/devicetree/bindings/ata/brcm,sata-brcmstb.txt +++ b/Documentation/devicetree/bindings/ata/brcm,sata-brcmstb.txt @@ -4,7 +4,9 @@ SATA nodes are defined to describe on-chip Serial ATA controllers. Each SATA controller should have its own node. Required properties: -- compatible : compatible list, may contain "brcm,bcm7445-ahci" and/or +- compatible : should be one or more of + "brcm,bcm7425-ahci" + "brcm,bcm7445-ahci" "brcm,sata3-ahci" - reg : register mappings for AHCI and SATA_TOP_CTRL - reg-names : "ahci" and "top-ctrl" diff --git a/Documentation/devicetree/bindings/ata/sata_rcar.txt b/Documentation/devicetree/bindings/ata/sata_rcar.txt index 2493a5a31655..0764f9ab63dc 100644 --- a/Documentation/devicetree/bindings/ata/sata_rcar.txt +++ b/Documentation/devicetree/bindings/ata/sata_rcar.txt @@ -8,6 +8,7 @@ Required properties: - "renesas,sata-r8a7790" for R-Car H2 other than ES1 - "renesas,sata-r8a7791" for R-Car M2-W - "renesas,sata-r8a7793" for R-Car M2-N + - "renesas,sata-r8a7795" for R-Car H3 - reg : address and length of the SATA registers; - interrupts : must consist of one interrupt specifier. - clocks : must contain a reference to the functional clock. diff --git a/Documentation/devicetree/bindings/cpufreq/cpufreq-st.txt b/Documentation/devicetree/bindings/cpufreq/cpufreq-st.txt new file mode 100644 index 000000000000..d91a02a3b6b0 --- /dev/null +++ b/Documentation/devicetree/bindings/cpufreq/cpufreq-st.txt @@ -0,0 +1,91 @@ +Binding for ST's CPUFreq driver +=============================== + +ST's CPUFreq driver attempts to read 'process' and 'version' attributes +from the SoC, then supplies the OPP framework with 'prop' and 'supported +hardware' information respectively. The framework is then able to read +the DT and operate in the usual way. + +For more information about the expected DT format [See: ../opp/opp.txt]. + +Frequency Scaling only +---------------------- + +No vendor specific driver required for this. + +Located in CPU's node: + +- operating-points : [See: ../power/opp.txt] + +Example [safe] +-------------- + +cpus { + cpu@0 { + /* kHz uV */ + operating-points = <1500000 0 + 1200000 0 + 800000 0 + 500000 0>; + }; +}; + +Dynamic Voltage and Frequency Scaling (DVFS) +-------------------------------------------- + +This requires the ST CPUFreq driver to supply 'process' and 'version' info. + +Located in CPU's node: + +- operating-points-v2 : [See ../power/opp.txt] + +Example [unsafe] +---------------- + +cpus { + cpu@0 { + operating-points-v2 = <&cpu0_opp_table>; + }; +}; + +cpu0_opp_table: opp_table { + compatible = "operating-points-v2"; + + /* ############################################################### */ + /* # WARNING: Do not attempt to copy/replicate these nodes, # */ + /* # they are only to be supplied by the bootloader !!! # */ + /* ############################################################### */ + opp0 { + /* Major Minor Substrate */ + /* 2 all all */ + opp-supported-hw = <0x00000004 0xffffffff 0xffffffff>; + opp-hz = /bits/ 64 <1500000000>; + clock-latency-ns = <10000000>; + + opp-microvolt-pcode0 = <1200000>; + opp-microvolt-pcode1 = <1200000>; + opp-microvolt-pcode2 = <1200000>; + opp-microvolt-pcode3 = <1200000>; + opp-microvolt-pcode4 = <1170000>; + opp-microvolt-pcode5 = <1140000>; + opp-microvolt-pcode6 = <1100000>; + opp-microvolt-pcode7 = <1070000>; + }; + + opp1 { + /* Major Minor Substrate */ + /* all all all */ + opp-supported-hw = <0xffffffff 0xffffffff 0xffffffff>; + opp-hz = /bits/ 64 <1200000000>; + clock-latency-ns = <10000000>; + + opp-microvolt-pcode0 = <1110000>; + opp-microvolt-pcode1 = <1150000>; + opp-microvolt-pcode2 = <1100000>; + opp-microvolt-pcode3 = <1080000>; + opp-microvolt-pcode4 = <1040000>; + opp-microvolt-pcode5 = <1020000>; + opp-microvolt-pcode6 = <980000>; + opp-microvolt-pcode7 = <930000>; + }; +}; diff --git a/Documentation/devicetree/bindings/crypto/rockchip-crypto.txt b/Documentation/devicetree/bindings/crypto/rockchip-crypto.txt new file mode 100644 index 000000000000..096df34b11c1 --- /dev/null +++ b/Documentation/devicetree/bindings/crypto/rockchip-crypto.txt @@ -0,0 +1,29 @@ +Rockchip Electronics And Security Accelerator + +Required properties: +- compatible: Should be "rockchip,rk3288-crypto" +- reg: Base physical address of the engine and length of memory mapped + region +- interrupts: Interrupt number +- clocks: Reference to the clocks about crypto +- clock-names: "aclk" used to clock data + "hclk" used to clock data + "sclk" used to clock crypto accelerator + "apb_pclk" used to clock dma +- resets: Must contain an entry for each entry in reset-names. + See ../reset/reset.txt for details. +- reset-names: Must include the name "crypto-rst". + +Examples: + + crypto: cypto-controller@ff8a0000 { + compatible = "rockchip,rk3288-crypto"; + reg = <0xff8a0000 0x4000>; + interrupts = <GIC_SPI 48 IRQ_TYPE_LEVEL_HIGH>; + clocks = <&cru ACLK_CRYPTO>, <&cru HCLK_CRYPTO>, + <&cru SCLK_CRYPTO>, <&cru ACLK_DMAC1>; + clock-names = "aclk", "hclk", "sclk", "apb_pclk"; + resets = <&cru SRST_CRYPTO>; + reset-names = "crypto-rst"; + status = "okay"; + }; diff --git a/Documentation/devicetree/bindings/interrupt-controller/allwinner,sun67i-sc-nmi.txt b/Documentation/devicetree/bindings/interrupt-controller/allwinner,sunxi-nmi.txt index d1c5cdabc3e0..81cd3692405e 100644 --- a/Documentation/devicetree/bindings/interrupt-controller/allwinner,sun67i-sc-nmi.txt +++ b/Documentation/devicetree/bindings/interrupt-controller/allwinner,sunxi-nmi.txt @@ -4,7 +4,7 @@ Allwinner Sunxi NMI Controller Required properties: - compatible : should be "allwinner,sun7i-a20-sc-nmi" or - "allwinner,sun6i-a31-sc-nmi" + "allwinner,sun6i-a31-sc-nmi" or "allwinner,sun9i-a80-nmi" - reg : Specifies base physical address and size of the registers. - interrupt-controller : Identifies the node as an interrupt controller - #interrupt-cells : Specifies the number of cells needed to encode an diff --git a/Documentation/devicetree/bindings/interrupt-controller/arm,gic.txt b/Documentation/devicetree/bindings/interrupt-controller/arm,gic.txt index cc56021eb60b..5a1cb4bc3dfe 100644 --- a/Documentation/devicetree/bindings/interrupt-controller/arm,gic.txt +++ b/Documentation/devicetree/bindings/interrupt-controller/arm,gic.txt @@ -18,6 +18,7 @@ Main node required properties: "arm,cortex-a9-gic" "arm,gic-400" "arm,pl390" + "arm,tc11mp-gic" "brcm,brahma-b15-gic" "qcom,msm-8660-qgic" "qcom,msm-qgic2" diff --git a/Documentation/devicetree/bindings/interrupt-controller/hisilicon,mbigen-v2.txt b/Documentation/devicetree/bindings/interrupt-controller/hisilicon,mbigen-v2.txt new file mode 100644 index 000000000000..720f7c92e9a1 --- /dev/null +++ b/Documentation/devicetree/bindings/interrupt-controller/hisilicon,mbigen-v2.txt @@ -0,0 +1,74 @@ +Hisilicon mbigen device tree bindings. +======================================= + +Mbigen means: message based interrupt generator. + +MBI is kind of msi interrupt only used on Non-PCI devices. + +To reduce the wired interrupt number connected to GIC, +Hisilicon designed mbigen to collect and generate interrupt. + + +Non-pci devices can connect to mbigen and generate the +interrupt by writing ITS register. + +The mbigen chip and devices connect to mbigen have the following properties: + +Mbigen main node required properties: +------------------------------------------- +- compatible: Should be "hisilicon,mbigen-v2" + +- reg: Specifies the base physical address and size of the Mbigen + registers. + +- interrupt controller: Identifies the node as an interrupt controller + +- msi-parent: Specifies the MSI controller this mbigen use. + For more detail information,please refer to the generic msi-parent binding in + Documentation/devicetree/bindings/interrupt-controller/msi.txt. + +- num-pins: the total number of pins implemented in this Mbigen + instance. + +- #interrupt-cells : Specifies the number of cells needed to encode an + interrupt source. The value must be 2. + + The 1st cell is hardware pin number of the interrupt.This number is local to + each mbigen chip and in the range from 0 to the maximum interrupts number + of the mbigen. + + The 2nd cell is the interrupt trigger type. + The value of this cell should be: + 1: rising edge triggered + or + 4: high level triggered + +Examples: + + mbigen_device_gmac:intc { + compatible = "hisilicon,mbigen-v2"; + reg = <0x0 0xc0080000 0x0 0x10000>; + interrupt-controller; + msi-parent = <&its_dsa 0x40b1c>; + num-pins = <9>; + #interrupt-cells = <2>; + }; + +Devices connect to mbigen required properties: +---------------------------------------------------- +-interrupt-parent: Specifies the mbigen device node which device connected. + +-interrupts:Specifies the interrupt source. + For the specific information of each cell in this property,please refer to + the "interrupt-cells" description mentioned above. + +Examples: + gmac0: ethernet@c2080000 { + #address-cells = <1>; + #size-cells = <0>; + reg = <0 0xc2080000 0 0x20000>, + <0 0xc0000000 0 0x1000>; + interrupt-parent = <&mbigen_device_gmac>; + interrupts = <656 1>, + <657 1>; + }; diff --git a/Documentation/devicetree/bindings/interrupt-controller/technologic,ts4800.txt b/Documentation/devicetree/bindings/interrupt-controller/technologic,ts4800.txt new file mode 100644 index 000000000000..7f15f1b0325b --- /dev/null +++ b/Documentation/devicetree/bindings/interrupt-controller/technologic,ts4800.txt @@ -0,0 +1,16 @@ +TS-4800 FPGA interrupt controller + +TS-4800 FPGA has an internal interrupt controller. When one of the +interrupts is triggered, the SoC is notified, usually using a GPIO as +parent interrupt source. + +Required properties: +- compatible: should be "technologic,ts4800-irqc" +- interrupt-controller: identifies the node as an interrupt controller +- reg: physical base address of the controller and length of memory mapped + region +- #interrupt-cells: specifies the number of cells needed to encode an interrupt + source, should be 1. +- interrupt-parent: phandle to the parent interrupt controller this one is + cascaded from +- interrupts: specifies the interrupt line in the interrupt-parent controller diff --git a/Documentation/devicetree/bindings/media/stih407-c8sectpfe.txt b/Documentation/devicetree/bindings/media/stih407-c8sectpfe.txt index d4def767bdfe..cc51b1fd6e0c 100644 --- a/Documentation/devicetree/bindings/media/stih407-c8sectpfe.txt +++ b/Documentation/devicetree/bindings/media/stih407-c8sectpfe.txt @@ -35,7 +35,7 @@ Required properties (tsin (child) node): - tsin-num : tsin id of the InputBlock (must be between 0 to 6) - i2c-bus : phandle to the I2C bus DT node which the demodulators & tuners on this tsin channel are connected. -- rst-gpio : reset gpio for this tsin channel. +- reset-gpios : reset gpio for this tsin channel. Optional properties (tsin (child) node): @@ -55,27 +55,27 @@ Example: status = "okay"; reg = <0x08a20000 0x10000>, <0x08a00000 0x4000>; reg-names = "stfe", "stfe-ram"; - interrupts = <0 34 0>, <0 35 0>; + interrupts = <GIC_SPI 34 IRQ_TYPE_NONE>, <GIC_SPI 35 IRQ_TYPE_NONE>; interrupt-names = "stfe-error-irq", "stfe-idle-irq"; - - pinctrl-names = "tsin0-serial", "tsin0-parallel", "tsin3-serial", - "tsin4-serial", "tsin5-serial"; - pinctrl-0 = <&pinctrl_tsin0_serial>; pinctrl-1 = <&pinctrl_tsin0_parallel>; pinctrl-2 = <&pinctrl_tsin3_serial>; pinctrl-3 = <&pinctrl_tsin4_serial_alt3>; pinctrl-4 = <&pinctrl_tsin5_serial_alt1>; - + pinctrl-names = "tsin0-serial", + "tsin0-parallel", + "tsin3-serial", + "tsin4-serial", + "tsin5-serial"; clocks = <&clk_s_c0_flexgen CLK_PROC_STFE>; - clock-names = "stfe"; + clock-names = "c8sectpfe"; /* tsin0 is TSA on NIMA */ tsin0: port@0 { tsin-num = <0>; serial-not-parallel; i2c-bus = <&ssc2>; - rst-gpio = <&pio15 4 0>; + reset-gpios = <&pio15 4 GPIO_ACTIVE_HIGH>; dvb-card = <STV0367_TDA18212_NIMA_1>; }; @@ -83,7 +83,7 @@ Example: tsin-num = <3>; serial-not-parallel; i2c-bus = <&ssc3>; - rst-gpio = <&pio15 7 0>; + reset-gpios = <&pio15 7 GPIO_ACTIVE_HIGH>; dvb-card = <STV0367_TDA18212_NIMB_1>; }; }; diff --git a/Documentation/devicetree/bindings/mmc/renesas,mmcif.txt b/Documentation/devicetree/bindings/mmc/renesas,mmcif.txt index cae29eb5733d..ff611fa66871 100644 --- a/Documentation/devicetree/bindings/mmc/renesas,mmcif.txt +++ b/Documentation/devicetree/bindings/mmc/renesas,mmcif.txt @@ -11,6 +11,7 @@ Required properties: - "renesas,mmcif-r8a7740" for the MMCIF found in r8a7740 SoCs - "renesas,mmcif-r8a7790" for the MMCIF found in r8a7790 SoCs - "renesas,mmcif-r8a7791" for the MMCIF found in r8a7791 SoCs + - "renesas,mmcif-r8a7793" for the MMCIF found in r8a7793 SoCs - "renesas,mmcif-r8a7794" for the MMCIF found in r8a7794 SoCs - clocks: reference to the functional clock diff --git a/Documentation/devicetree/bindings/net/cpsw.txt b/Documentation/devicetree/bindings/net/cpsw.txt index 9853f8e70966..28a4781ab6d7 100644 --- a/Documentation/devicetree/bindings/net/cpsw.txt +++ b/Documentation/devicetree/bindings/net/cpsw.txt @@ -40,18 +40,18 @@ Optional properties: Slave Properties: Required properties: -- phy_id : Specifies slave phy id - phy-mode : See ethernet.txt file in the same directory Optional properties: - dual_emac_res_vlan : Specifies VID to be used to segregate the ports - mac-address : See ethernet.txt file in the same directory +- phy_id : Specifies slave phy id - phy-handle : See ethernet.txt file in the same directory Slave sub-nodes: - fixed-link : See fixed-link.txt file in the same directory - Either the properties phy_id and phy-mode, - or the sub-node fixed-link can be specified + Either the property phy_id, or the sub-node + fixed-link can be specified Note: "ti,hwmods" field is used to fetch the base address and irq resources from TI, omap hwmod data base during device registration. diff --git a/Documentation/devicetree/bindings/net/dsa/dsa.txt b/Documentation/devicetree/bindings/net/dsa/dsa.txt index 04e6bef3ac3f..5fdbbcdf8c4b 100644 --- a/Documentation/devicetree/bindings/net/dsa/dsa.txt +++ b/Documentation/devicetree/bindings/net/dsa/dsa.txt @@ -31,6 +31,8 @@ A switch child node has the following optional property: switch. Must be set if the switch can not detect the presence and/or size of a connected EEPROM, otherwise optional. +- reset-gpios : phandle and specifier to a gpio line connected to + reset pin of the switch chip. A switch may have multiple "port" children nodes @@ -114,6 +116,7 @@ Example: #size-cells = <0>; reg = <17 1>; /* MDIO address 17, switch 1 in tree */ mii-bus = <&mii_bus1>; + reset-gpios = <&gpio5 1 GPIO_ACTIVE_LOW>; switch1port0: port@0 { reg = <0>; diff --git a/Documentation/devicetree/bindings/net/hisilicon-hns-mdio.txt b/Documentation/devicetree/bindings/net/hisilicon-hns-mdio.txt index 9c23fdf25018..4a7ede9657b0 100644 --- a/Documentation/devicetree/bindings/net/hisilicon-hns-mdio.txt +++ b/Documentation/devicetree/bindings/net/hisilicon-hns-mdio.txt @@ -1,7 +1,12 @@ Hisilicon MDIO bus controller Properties: -- compatible: "hisilicon,mdio","hisilicon,hns-mdio". +- compatible: can be one of: + "hisilicon,hns-mdio" + "hisilicon,mdio" + "hisilicon,hns-mdio" is recommended to be used for hip05 and later SOCs, + while "hisilicon,mdio" is optional for backwards compatibility only on + hip04 Soc. - reg: The base address of the MDIO bus controller register bank. - #address-cells: Must be <1>. - #size-cells: Must be <0>. MDIO addresses have no size component. diff --git a/Documentation/devicetree/bindings/net/ieee802154/adf7242.txt b/Documentation/devicetree/bindings/net/ieee802154/adf7242.txt new file mode 100644 index 000000000000..dea5124cdc52 --- /dev/null +++ b/Documentation/devicetree/bindings/net/ieee802154/adf7242.txt @@ -0,0 +1,18 @@ +* ADF7242 IEEE 802.15.4 * + +Required properties: + - compatible: should be "adi,adf7242" + - spi-max-frequency: maximal bus speed (12.5 MHz) + - reg: the chipselect index + - interrupts: the interrupt generated by the device via pin IRQ1. + IRQ_TYPE_LEVEL_HIGH (4) or IRQ_TYPE_EDGE_FALLING (1) + +Example: + + adf7242@0 { + compatible = "adi,adf7242"; + spi-max-frequency = <10000000>; + reg = <0>; + interrupts = <98 IRQ_TYPE_LEVEL_HIGH>; + interrupt-parent = <&gpio3>; + }; diff --git a/Documentation/devicetree/bindings/net/macb.txt b/Documentation/devicetree/bindings/net/macb.txt index b5d79761ac97..5c397ca14cfe 100644 --- a/Documentation/devicetree/bindings/net/macb.txt +++ b/Documentation/devicetree/bindings/net/macb.txt @@ -4,6 +4,7 @@ Required properties: - compatible: Should be "cdns,[<chip>-]{macb|gem}" Use "cdns,at91sam9260-macb" for Atmel at91sam9 SoCs or the 10/100Mbit IP available on sama5d3 SoCs. + Use "cdns,np4-macb" for NP4 SoC devices. Use "cdns,at32ap7000-macb" for other 10/100 usage or use the generic form: "cdns,macb". Use "cdns,pc302-gem" for Picochip picoXcell pc302 and later devices based on the Cadence GEM, or the generic form: "cdns,gem". @@ -19,6 +20,9 @@ Required properties: Optional elements: 'tx_clk' - clocks: Phandles to input clocks. +Optional properties for PHY child node: +- reset-gpios : Should specify the gpio for phy reset + Examples: macb0: ethernet@fffc4000 { @@ -29,4 +33,8 @@ Examples: local-mac-address = [3a 0e 03 04 05 06]; clock-names = "pclk", "hclk", "tx_clk"; clocks = <&clkc 30>, <&clkc 30>, <&clkc 13>; + ethernet-phy@1 { + reg = <0x1>; + reset-gpios = <&pioE 6 1>; + }; }; diff --git a/Documentation/devicetree/bindings/net/micrel-ksz90x1.txt b/Documentation/devicetree/bindings/net/micrel-ksz90x1.txt index 692076fda0e5..f9c32adab5c6 100644 --- a/Documentation/devicetree/bindings/net/micrel-ksz90x1.txt +++ b/Documentation/devicetree/bindings/net/micrel-ksz90x1.txt @@ -1,8 +1,9 @@ Micrel KSZ9021/KSZ9031 Gigabit Ethernet PHY -Some boards require special tuning values, particularly when it comes to -clock delays. You can specify clock delay values by adding -micrel-specific properties to an Ethernet OF device node. +Some boards require special tuning values, particularly when it comes +to clock delays. You can specify clock delay values in the PHY OF +device node. Deprecated, but still supported, these properties can +also be added to an Ethernet OF device node. Note that these settings are applied after any phy-specific fixup from phy_fixup_list (see phy_init_hw() from drivers/net/phy/phy_device.c), @@ -57,16 +58,6 @@ KSZ9031: Examples: - /* Attach to an Ethernet device with autodetected PHY */ - &enet { - rxc-skew-ps = <3000>; - rxdv-skew-ps = <0>; - txc-skew-ps = <3000>; - txen-skew-ps = <0>; - status = "okay"; - }; - - /* Attach to an explicitly-specified PHY */ mdio { phy0: ethernet-phy@0 { rxc-skew-ps = <3000>; diff --git a/Documentation/devicetree/bindings/net/nfc/st95hf.txt b/Documentation/devicetree/bindings/net/nfc/st95hf.txt new file mode 100644 index 000000000000..ea3178bc9ddd --- /dev/null +++ b/Documentation/devicetree/bindings/net/nfc/st95hf.txt @@ -0,0 +1,50 @@ +* STMicroelectronics : NFC Transceiver ST95HF + +ST NFC Transceiver is required to attach with SPI bus. +ST95HF node should be defined in DT as SPI slave device of SPI +master with which ST95HF transceiver is physically connected. +The properties defined below are required to be the part of DT +to include ST95HF transceiver into the platform. + +Required properties: +=================== +- reg: Address of SPI slave "ST95HF transceiver" on SPI master bus. + +- compatible: should be "st,st95hf" for ST95HF NFC transceiver + +- spi-max-frequency: Max. operating SPI frequency for ST95HF + transceiver. + +- enable-gpio: GPIO line to enable ST95HF transceiver. + +- interrupt-parent : Standard way to specify the controller to which + ST95HF transceiver's interrupt is routed. + +- interrupts : Standard way to define ST95HF transceiver's out + interrupt. + +Optional property: +================= +- st95hfvin-supply : This is an optional property. It contains a + phandle to ST95HF transceiver's regulator supply node in DT. + +Example: +======= +spi@9840000 { + reg = <0x9840000 0x110>; + #address-cells = <1>; + #size-cells = <0>; + cs-gpios = <&pio0 4>; + status = "okay"; + + st95hf@0{ + reg = <0>; + compatible = "st,st95hf"; + status = "okay"; + spi-max-frequency = <1000000>; + enable-gpio = <&pio4 0>; + interrupt-parent = <&pio0>; + interrupts = <7 IRQ_TYPE_EDGE_FALLING>; + }; + +}; diff --git a/Documentation/devicetree/bindings/net/renesas,ravb.txt b/Documentation/devicetree/bindings/net/renesas,ravb.txt index b486f3f5f6a3..81a9f9e6b45f 100644 --- a/Documentation/devicetree/bindings/net/renesas,ravb.txt +++ b/Documentation/devicetree/bindings/net/renesas,ravb.txt @@ -5,8 +5,18 @@ interface contains. Required properties: - compatible: "renesas,etheravb-r8a7790" if the device is a part of R8A7790 SoC. + "renesas,etheravb-r8a7791" if the device is a part of R8A7791 SoC. + "renesas,etheravb-r8a7792" if the device is a part of R8A7792 SoC. + "renesas,etheravb-r8a7793" if the device is a part of R8A7793 SoC. "renesas,etheravb-r8a7794" if the device is a part of R8A7794 SoC. "renesas,etheravb-r8a7795" if the device is a part of R8A7795 SoC. + "renesas,etheravb-rcar-gen2" for generic R-Car Gen 2 compatible interface. + "renesas,etheravb-rcar-gen3" for generic R-Car Gen 3 compatible interface. + + When compatible with the generic version, nodes must list the + SoC-specific version corresponding to the platform first + followed by the generic version. + - reg: offset and length of (1) the register block and (2) the stream buffer. - interrupts: A list of interrupt-specifiers, one for each entry in interrupt-names. @@ -37,7 +47,7 @@ Optional properties: Example: ethernet@e6800000 { - compatible = "renesas,etheravb-r8a7795"; + compatible = "renesas,etheravb-r8a7795", "renesas,etheravb-rcar-gen3"; reg = <0 0xe6800000 0 0x800>, <0 0xe6a00000 0 0x10000>; interrupt-parent = <&gic>; interrupts = <GIC_SPI 39 IRQ_TYPE_LEVEL_HIGH>, diff --git a/Documentation/devicetree/bindings/net/socfpga-dwmac.txt b/Documentation/devicetree/bindings/net/socfpga-dwmac.txt index 3a9d67951606..72d82d684342 100644 --- a/Documentation/devicetree/bindings/net/socfpga-dwmac.txt +++ b/Documentation/devicetree/bindings/net/socfpga-dwmac.txt @@ -11,6 +11,8 @@ Required properties: designware version numbers documented in stmmac.txt - altr,sysmgr-syscon : Should be the phandle to the system manager node that encompasses the glue register, the register offset, and the register shift. + - altr,f2h_ptp_ref_clk use f2h_ptp_ref_clk instead of default eosc1 clock + for ptp ref clk. This affects all emacs as the clock is common. Optional properties: altr,emac-splitter: Should be the phandle to the emac splitter soft IP node if diff --git a/Documentation/devicetree/bindings/net/stmmac.txt b/Documentation/devicetree/bindings/net/stmmac.txt index f34fc3c81a75..e862a922bd3f 100644 --- a/Documentation/devicetree/bindings/net/stmmac.txt +++ b/Documentation/devicetree/bindings/net/stmmac.txt @@ -35,18 +35,18 @@ Optional properties: - reset-names: Should contain the reset signal name "stmmaceth", if a reset phandle is given - max-frame-size: See ethernet.txt file in the same directory -- clocks: If present, the first clock should be the GMAC main clock and - the second clock should be peripheral's register interface clock. Further - clocks may be specified in derived bindings. -- clock-names: One name for each entry in the clocks property, the - first one should be "stmmaceth" and the second one should be "pclk". -- clk_ptp_ref: this is the PTP reference clock; in case of the PTP is - available this clock is used for programming the Timestamp Addend Register. - If not passed then the system clock will be used and this is fine on some - platforms. +- clocks: If present, the first clock should be the GMAC main clock + The optional second clock should be peripheral's register interface clock. + The third optional clock should be the ptp reference clock. + Further clocks may be specified in derived bindings. +- clock-names: One name for each entry in the clocks property. + The first one should be "stmmaceth". + The optional second one should be "pclk". + The optional third one should be "clk_ptp_ref". - snps,burst_len: The AXI burst lenth value of the AXI BUS MODE register. - tx-fifo-depth: See ethernet.txt file in the same directory - rx-fifo-depth: See ethernet.txt file in the same directory +- mdio: with compatible = "snps,dwmac-mdio", create and register mdio bus. Examples: @@ -65,4 +65,11 @@ Examples: tx-fifo-depth = <16384>; clocks = <&clock>; clock-names = "stmmaceth"; + mdio0 { + #address-cells = <1>; + #size-cells = <0>; + compatible = "snps,dwmac-mdio"; + phy1: ethernet-phy@0 { + }; + }; }; diff --git a/Documentation/devicetree/bindings/opp/opp.txt b/Documentation/devicetree/bindings/opp/opp.txt index 0cb44dc21f97..601256fe8c0d 100644 --- a/Documentation/devicetree/bindings/opp/opp.txt +++ b/Documentation/devicetree/bindings/opp/opp.txt @@ -45,21 +45,10 @@ Devices supporting OPPs must set their "operating-points-v2" property with phandle to a OPP table in their DT node. The OPP core will use this phandle to find the operating points for the device. -Devices may want to choose OPP tables at runtime and so can provide a list of -phandles here. But only *one* of them should be chosen at runtime. This must be -accompanied by a corresponding "operating-points-names" property, to uniquely -identify the OPP tables. - If required, this can be extended for SoC vendor specfic bindings. Such bindings should be documented as Documentation/devicetree/bindings/power/<vendor>-opp.txt and should have a compatible description like: "operating-points-v2-<vendor>". -Optional properties: -- operating-points-names: Names of OPP tables (required if multiple OPP - tables are present), to uniquely identify them. The same list must be present - for all the CPUs which are sharing clock/voltage rails and hence the OPP - tables. - * OPP Table Node This describes the OPPs belonging to a device. This node can have following @@ -100,6 +89,14 @@ Optional properties: Entries for multiple regulators must be present in the same order as regulators are specified in device's DT node. +- opp-microvolt-<name>: Named opp-microvolt property. This is exactly similar to + the above opp-microvolt property, but allows multiple voltage ranges to be + provided for the same OPP. At runtime, the platform can pick a <name> and + matching opp-microvolt-<name> property will be enabled for all OPPs. If the + platform doesn't pick a specific <name> or the <name> doesn't match with any + opp-microvolt-<name> properties, then opp-microvolt property shall be used, if + present. + - opp-microamp: The maximum current drawn by the device in microamperes considering system specific parameters (such as transients, process, aging, maximum operating temperature range etc.) as necessary. This may be used to @@ -112,6 +109,9 @@ Optional properties: for few regulators, then this should be marked as zero for them. If it isn't required for any regulator, then this property need not be present. +- opp-microamp-<name>: Named opp-microamp property. Similar to + opp-microvolt-<name> property, but for microamp instead. + - clock-latency-ns: Specifies the maximum possible transition latency (in nanoseconds) for switching to this OPP from any other OPP. @@ -123,6 +123,26 @@ Optional properties: - opp-suspend: Marks the OPP to be used during device suspend. Only one OPP in the table should have this. +- opp-supported-hw: This enables us to select only a subset of OPPs from the + larger OPP table, based on what version of the hardware we are running on. We + still can't have multiple nodes with the same opp-hz value in OPP table. + + It's an user defined array containing a hierarchy of hardware version numbers, + supported by the OPP. For example: a platform with hierarchy of three levels + of versions (A, B and C), this field should be like <X Y Z>, where X + corresponds to Version hierarchy A, Y corresponds to version hierarchy B and Z + corresponds to version hierarchy C. + + Each level of hierarchy is represented by a 32 bit value, and so there can be + only 32 different supported version per hierarchy. i.e. 1 bit per version. A + value of 0xFFFFFFFF will enable the OPP for all versions for that hierarchy + level. And a value of 0x00000000 will disable the OPP completely, and so we + never want that to happen. + + If 32 values aren't sufficient for a version hierarchy, than that version + hierarchy can be contained in multiple 32 bit values. i.e. <X Y Z1 Z2> in the + above example, Z1 & Z2 refer to the version hierarchy Z. + - status: Marks the node enabled/disabled. Example 1: Single cluster Dual-core ARM cortex A9, switch DVFS states together. @@ -157,20 +177,20 @@ Example 1: Single cluster Dual-core ARM cortex A9, switch DVFS states together. compatible = "operating-points-v2"; opp-shared; - opp00 { + opp@1000000000 { opp-hz = /bits/ 64 <1000000000>; opp-microvolt = <970000 975000 985000>; opp-microamp = <70000>; clock-latency-ns = <300000>; opp-suspend; }; - opp01 { + opp@1100000000 { opp-hz = /bits/ 64 <1100000000>; opp-microvolt = <980000 1000000 1010000>; opp-microamp = <80000>; clock-latency-ns = <310000>; }; - opp02 { + opp@1200000000 { opp-hz = /bits/ 64 <1200000000>; opp-microvolt = <1025000>; clock-latency-ns = <290000>; @@ -236,20 +256,20 @@ independently. * independently. */ - opp00 { + opp@1000000000 { opp-hz = /bits/ 64 <1000000000>; opp-microvolt = <970000 975000 985000>; opp-microamp = <70000>; clock-latency-ns = <300000>; opp-suspend; }; - opp01 { + opp@1100000000 { opp-hz = /bits/ 64 <1100000000>; opp-microvolt = <980000 1000000 1010000>; opp-microamp = <80000>; clock-latency-ns = <310000>; }; - opp02 { + opp@1200000000 { opp-hz = /bits/ 64 <1200000000>; opp-microvolt = <1025000>; opp-microamp = <90000; @@ -312,20 +332,20 @@ DVFS state together. compatible = "operating-points-v2"; opp-shared; - opp00 { + opp@1000000000 { opp-hz = /bits/ 64 <1000000000>; opp-microvolt = <970000 975000 985000>; opp-microamp = <70000>; clock-latency-ns = <300000>; opp-suspend; }; - opp01 { + opp@1100000000 { opp-hz = /bits/ 64 <1100000000>; opp-microvolt = <980000 1000000 1010000>; opp-microamp = <80000>; clock-latency-ns = <310000>; }; - opp02 { + opp@1200000000 { opp-hz = /bits/ 64 <1200000000>; opp-microvolt = <1025000>; opp-microamp = <90000>; @@ -338,20 +358,20 @@ DVFS state together. compatible = "operating-points-v2"; opp-shared; - opp10 { + opp@1300000000 { opp-hz = /bits/ 64 <1300000000>; opp-microvolt = <1045000 1050000 1055000>; opp-microamp = <95000>; clock-latency-ns = <400000>; opp-suspend; }; - opp11 { + opp@1400000000 { opp-hz = /bits/ 64 <1400000000>; opp-microvolt = <1075000>; opp-microamp = <100000>; clock-latency-ns = <400000>; }; - opp12 { + opp@1500000000 { opp-hz = /bits/ 64 <1500000000>; opp-microvolt = <1010000 1100000 1110000>; opp-microamp = <95000>; @@ -378,7 +398,7 @@ Example 4: Handling multiple regulators compatible = "operating-points-v2"; opp-shared; - opp00 { + opp@1000000000 { opp-hz = /bits/ 64 <1000000000>; opp-microvolt = <970000>, /* Supply 0 */ <960000>, /* Supply 1 */ @@ -391,7 +411,7 @@ Example 4: Handling multiple regulators /* OR */ - opp00 { + opp@1000000000 { opp-hz = /bits/ 64 <1000000000>; opp-microvolt = <970000 975000 985000>, /* Supply 0 */ <960000 965000 975000>, /* Supply 1 */ @@ -404,7 +424,7 @@ Example 4: Handling multiple regulators /* OR */ - opp00 { + opp@1000000000 { opp-hz = /bits/ 64 <1000000000>; opp-microvolt = <970000 975000 985000>, /* Supply 0 */ <960000 965000 975000>, /* Supply 1 */ @@ -417,7 +437,8 @@ Example 4: Handling multiple regulators }; }; -Example 5: Multiple OPP tables +Example 5: opp-supported-hw +(example: three level hierarchy of versions: cuts, substrate and process) / { cpus { @@ -426,40 +447,73 @@ Example 5: Multiple OPP tables ... cpu-supply = <&cpu_supply> - operating-points-v2 = <&cpu0_opp_table_slow>, <&cpu0_opp_table_fast>; - operating-points-names = "slow", "fast"; + operating-points-v2 = <&cpu0_opp_table_slow>; }; }; - cpu0_opp_table_slow: opp_table_slow { + opp_table { compatible = "operating-points-v2"; status = "okay"; opp-shared; - opp00 { + opp@600000000 { + /* + * Supports all substrate and process versions for 0xF + * cuts, i.e. only first four cuts. + */ + opp-supported-hw = <0xF 0xFFFFFFFF 0xFFFFFFFF> opp-hz = /bits/ 64 <600000000>; + opp-microvolt = <900000 915000 925000>; ... }; - opp01 { + opp@800000000 { + /* + * Supports: + * - cuts: only one, 6th cut (represented by 6th bit). + * - substrate: supports 16 different substrate versions + * - process: supports 9 different process versions + */ + opp-supported-hw = <0x20 0xff0000ff 0x0000f4f0> opp-hz = /bits/ 64 <800000000>; + opp-microvolt = <900000 915000 925000>; ... }; }; +}; + +Example 6: opp-microvolt-<name>, opp-microamp-<name>: +(example: device with two possible microvolt ranges: slow and fast) - cpu0_opp_table_fast: opp_table_fast { +/ { + cpus { + cpu@0 { + compatible = "arm,cortex-a7"; + ... + + operating-points-v2 = <&cpu0_opp_table>; + }; + }; + + cpu0_opp_table: opp_table0 { compatible = "operating-points-v2"; - status = "okay"; opp-shared; - opp10 { + opp@1000000000 { opp-hz = /bits/ 64 <1000000000>; - ... + opp-microvolt-slow = <900000 915000 925000>; + opp-microvolt-fast = <970000 975000 985000>; + opp-microamp-slow = <70000>; + opp-microamp-fast = <71000>; }; - opp11 { - opp-hz = /bits/ 64 <1100000000>; - ... + opp@1200000000 { + opp-hz = /bits/ 64 <1200000000>; + opp-microvolt-slow = <900000 915000 925000>, /* Supply vcc0 */ + <910000 925000 935000>; /* Supply vcc1 */ + opp-microvolt-fast = <970000 975000 985000>, /* Supply vcc0 */ + <960000 965000 975000>; /* Supply vcc1 */ + opp-microamp = <70000>; /* Will be used for both slow/fast */ }; }; }; diff --git a/Documentation/devicetree/bindings/pinctrl/allwinner,sunxi-pinctrl.txt b/Documentation/devicetree/bindings/pinctrl/allwinner,sunxi-pinctrl.txt index b321b26780dc..9213b27e1036 100644 --- a/Documentation/devicetree/bindings/pinctrl/allwinner,sunxi-pinctrl.txt +++ b/Documentation/devicetree/bindings/pinctrl/allwinner,sunxi-pinctrl.txt @@ -17,7 +17,10 @@ Required properties: "allwinner,sun8i-a23-pinctrl" "allwinner,sun8i-a23-r-pinctrl" "allwinner,sun8i-a33-pinctrl" + "allwinner,sun9i-a80-pinctrl" + "allwinner,sun9i-a80-r-pinctrl" "allwinner,sun8i-a83t-pinctrl" + "allwinner,sun8i-h3-pinctrl" - reg: Should contain the register physical address and length for the pin controller. diff --git a/Documentation/devicetree/bindings/pinctrl/brcm,cygnus-gpio.txt b/Documentation/devicetree/bindings/pinctrl/brcm,iproc-gpio.txt index 16589fb6f420..e4277921f3e3 100644 --- a/Documentation/devicetree/bindings/pinctrl/brcm,cygnus-gpio.txt +++ b/Documentation/devicetree/bindings/pinctrl/brcm,iproc-gpio.txt @@ -1,4 +1,4 @@ -Broadcom Cygnus GPIO/PINCONF Controller +Broadcom iProc GPIO/PINCONF Controller Required properties: @@ -7,9 +7,12 @@ Required properties: "brcm,cygnus-crmu-gpio" or "brcm,iproc-gpio" - reg: - Define the base and range of the I/O address space that contains the Cygnus + Define the base and range of the I/O address space that contains SoC GPIO/PINCONF controller registers +- ngpios: + Total number of in-use slots in GPIO controller + - #gpio-cells: Must be two. The first cell is the GPIO pin number (within the controller's pin space) and the second cell is used for the following: @@ -57,6 +60,7 @@ Example: compatible = "brcm,cygnus-ccm-gpio"; reg = <0x1800a000 0x50>, <0x0301d164 0x20>; + ngpios = <24>; #gpio-cells = <2>; gpio-controller; interrupts = <GIC_SPI 84 IRQ_TYPE_LEVEL_HIGH>; @@ -78,6 +82,7 @@ Example: gpio_asiu: gpio@180a5000 { compatible = "brcm,cygnus-asiu-gpio"; reg = <0x180a5000 0x668>; + ngpios = <146>; #gpio-cells = <2>; gpio-controller; interrupts = <GIC_SPI 174 IRQ_TYPE_LEVEL_HIGH>; diff --git a/Documentation/devicetree/bindings/pinctrl/brcm,nsp-gpio.txt b/Documentation/devicetree/bindings/pinctrl/brcm,nsp-gpio.txt new file mode 100644 index 000000000000..0844168a6dd4 --- /dev/null +++ b/Documentation/devicetree/bindings/pinctrl/brcm,nsp-gpio.txt @@ -0,0 +1,80 @@ +Broadcom Northstar plus (NSP) GPIO/PINCONF Controller + +Required properties: +- compatible: + Must be "brcm,nsp-gpio-a" + +- reg: + Should contain the register physical address and length for each of + GPIO base, IO control registers + +- #gpio-cells: + Must be two. The first cell is the GPIO pin number (within the + controller's pin space) and the second cell is used for the following: + bit[0]: polarity (0 for active high and 1 for active low) + +- gpio-controller: + Specifies that the node is a GPIO controller + +- ngpios: + Number of gpios supported (58x25 supports 32 and 58x23 supports 24) + +Optional properties: +- interrupts: + Interrupt ID + +- interrupt-controller: + Specifies that the node is an interrupt controller + +- gpio-ranges: + Specifies the mapping between gpio controller and pin-controllers pins. + This requires 4 fields in cells defined as - + 1. Phandle of pin-controller. + 2. GPIO base pin offset. + 3 Pin-control base pin offset. + 4. number of gpio pins which are linearly mapped from pin base. + +Supported generic PINCONF properties in child nodes: +- pins: + The list of pins (within the controller's own pin space) that properties + in the node apply to. Pin names are "gpio-<pin>" + +- bias-disable: + Disable pin bias + +- bias-pull-up: + Enable internal pull up resistor + +- bias-pull-down: + Enable internal pull down resistor + +- drive-strength: + Valid drive strength values include 2, 4, 6, 8, 10, 12, 14, 16 (mA) + +Example: + + gpioa: gpio@18000020 { + compatible = "brcm,nsp-gpio-a"; + reg = <0x18000020 0x100>, + <0x1803f1c4 0x1c>; + #gpio-cells = <2>; + gpio-controller; + ngpios = <32>; + gpio-ranges = <&pinctrl 0 0 31>; + interrupt-controller; + interrupts = <GIC_SPI 85 IRQ_TYPE_LEVEL_HIGH>; + + /* Hog a few default settings */ + pinctrl-names = "default"; + pinctrl-0 = <&led>; + led: led { + pins = "gpio-1"; + bias-pull-up; + }; + + pwr: pwr { + gpio-hog; + gpios = <3 1>; + output-high; + }; + }; diff --git a/Documentation/devicetree/bindings/pinctrl/lantiq,pinctrl-xway.txt b/Documentation/devicetree/bindings/pinctrl/lantiq,pinctrl-xway.txt index e89b4677567d..8e5216bcd748 100644 --- a/Documentation/devicetree/bindings/pinctrl/lantiq,pinctrl-xway.txt +++ b/Documentation/devicetree/bindings/pinctrl/lantiq,pinctrl-xway.txt @@ -1,7 +1,16 @@ Lantiq XWAY pinmux controller Required properties: -- compatible: "lantiq,pinctrl-xway" or "lantiq,pinctrl-xr9" +- compatible: "lantiq,pinctrl-xway", (DEPRECATED: Use "lantiq,pinctrl-danube") + "lantiq,pinctrl-xr9", (DEPRECATED: Use "lantiq,xrx100-pinctrl" or + "lantiq,xrx200-pinctrl") + "lantiq,pinctrl-ase", (DEPRECATED: Use "lantiq,ase-pinctrl") + "lantiq,<chip>-pinctrl", where <chip> is: + "ase" (XWAY AMAZON Family) + "danube" (XWAY DANUBE Family) + "xrx100" (XWAY xRX100 Family) + "xrx200" (XWAY xRX200 Family) + "xrx300" (XWAY xRX300 Family) - reg: Should contain the physical address and length of the gpio/pinmux register range @@ -36,19 +45,87 @@ Required subnode-properties: Valid values for group and function names: +XWAY: (DEPRECATED: Use DANUBE) mux groups: exin0, exin1, exin2, jtag, ebu a23, ebu a24, ebu a25, ebu clk, ebu cs1, ebu wait, nand ale, nand cs1, nand cle, spi, spi_cs1, spi_cs2, spi_cs3, - spi_cs4, spi_cs5, spi_cs6, asc0, asc0 cts rts, stp, nmi , gpt1, gpt2, + spi_cs4, spi_cs5, spi_cs6, asc0, asc0 cts rts, stp, nmi, gpt1, gpt2, gpt3, clkout0, clkout1, clkout2, clkout3, gnt1, gnt2, gnt3, req1, req2, req3 - additional mux groups (XR9 only): - mdio, nand rdy, nand rd, exin3, exin4, gnt4, req4 + functions: + spi, asc, cgu, jtag, exin, stp, gpt, nmi, pci, ebu + +XR9: ( DEPRECATED: Use xRX100/xRX200) + mux groups: + exin0, exin1, exin2, exin3, exin4, jtag, ebu a23, ebu a24, ebu a25, + ebu clk, ebu cs1, ebu wait, nand ale, nand cs1, nand cle, nand rdy, + nand rd, spi, spi_cs1, spi_cs2, spi_cs3, spi_cs4, spi_cs5, spi_cs6, + asc0, asc0 cts rts, stp, nmi, gpt1, gpt2, gpt3, clkout0, clkout1, + clkout2, clkout3, gnt1, gnt2, gnt3, gnt4, req1, req2, req3, req4, mdio, + gphy0 led0, gphy0 led1, gphy0 led2, gphy1 led0, gphy1 led1, gphy1 led2 + + functions: + spi, asc, cgu, jtag, exin, stp, gpt, nmi, pci, ebu, mdio, gphy + +AMAZON: + mux groups: + exin0, exin1, exin2, jtag, spi_di, spi_do, spi_clk, spi_cs1, spi_cs2, + spi_cs3, spi_cs4, spi_cs5, spi_cs6, asc, stp, gpt1, gpt2, gpt3, clkout0, + clkout1, clkout2, mdio, dfe led0, dfe led1, ephy led0, ephy led1, ephy led2 + + functions: + spi, asc, cgu, jtag, exin, stp, gpt, mdio, ephy, dfe + +DANUBE: + mux groups: + exin0, exin1, exin2, jtag, ebu a23, ebu a24, ebu a25, ebu clk, ebu cs1, + ebu wait, nand ale, nand cs1, nand cle, spi_di, spi_do, spi_clk, spi_cs1, + spi_cs2, spi_cs3, spi_cs4, spi_cs5, spi_cs6, asc0, asc0 cts rts, stp, nmi, + gpt1, gpt2, gpt3, clkout0, clkout1, clkout2, clkout3, gnt1, gnt2, gnt3, + req1, req2, req3, dfe led0, dfe led1 functions: - spi, asc, cgu, jtag, exin, stp, gpt, nmi, pci, ebu, mdio + spi, asc, cgu, jtag, exin, stp, gpt, nmi, pci, ebu, dfe +xRX100: + mux groups: + exin0, exin1, exin2, exin3, exin4, ebu a23, ebu a24, ebu a25, ebu clk, + ebu cs1, ebu wait, nand ale, nand cs1, nand cle, nand rdy, nand rd, + spi_di, spi_do, spi_clk, spi_cs1, spi_cs2, spi_cs3, spi_cs4, spi_cs5, + spi_cs6, asc0, asc0 cts rts, stp, nmi, gpt1, gpt2, gpt3, clkout0, clkout1, + clkout2, clkout3, gnt1, gnt2, gnt3, gnt4, req1, req2, req3, req4, mdio, + dfe led0, dfe led1 + + functions: + spi, asc, cgu, exin, stp, gpt, nmi, pci, ebu, mdio, dfe + +xRX200: + mux groups: + exin0, exin1, exin2, exin3, exin4, ebu a23, ebu a24, ebu a25, ebu clk, + ebu cs1, ebu wait, nand ale, nand cs1, nand cle, nand rdy, nand rd, + spi_di, spi_do, spi_clk, spi_cs1, spi_cs2, spi_cs3, spi_cs4, spi_cs5, + spi_cs6, usif uart_rx, usif uart_tx, usif uart_rts, usif uart_cts, + usif uart_dtr, usif uart_dsr, usif uart_dcd, usif uart_ri, usif spi_di, + usif spi_do, usif spi_clk, usif spi_cs0, usif spi_cs1, usif spi_cs2, + stp, nmi, gpt1, gpt2, gpt3, clkout0, clkout1, clkout2, clkout3, gnt1, + gnt2, gnt3, gnt4, req1, req2, req3, req4, mdio, dfe led0, dfe led1, + gphy0 led0, gphy0 led1, gphy0 led2, gphy1 led0, gphy1 led1, gphy1 led2 + + functions: + spi, usif, cgu, exin, stp, gpt, nmi, pci, ebu, mdio, dfe, gphy + +xRX300: + mux groups: + exin0, exin1, exin2, exin4, nand ale, nand cs0, nand cs1, nand cle, + nand rdy, nand rd, nand_d0, nand_d1, nand_d2, nand_d3, nand_d4, nand_d5, + nand_d6, nand_d7, nand_d1, nand wr, nand wp, nand se, spi_di, spi_do, + spi_clk, spi_cs1, spi_cs4, spi_cs6, usif uart_rx, usif uart_tx, + usif spi_di, usif spi_do, usif spi_clk, usif spi_cs0, stp, clkout2, + mdio, dfe led0, dfe led1, ephy0 led0, ephy0 led1, ephy1 led0, ephy1 led1 + + functions: + spi, usif, cgu, exin, stp, ebu, mdio, dfe, ephy Definition of pin configurations: @@ -62,15 +139,32 @@ Optional subnode-properties: 0: none, 1: down, 2: up. - lantiq,open-drain: Boolean, enables open-drain on the defined pin. -Valid values for XWAY pin names: +Valid values for XWAY pin names: (DEPRECATED: Use DANUBE) Pinconf pins can be referenced via the names io0-io31. -Valid values for XR9 pin names: +Valid values for XR9 pin names: (DEPRECATED: Use xrX100/xRX200) Pinconf pins can be referenced via the names io0-io55. +Valid values for AMAZON pin names: + Pinconf pins can be referenced via the names io0-io31. + +Valid values for DANUBE pin names: + Pinconf pins can be referenced via the names io0-io31. + +Valid values for xRX100 pin names: + Pinconf pins can be referenced via the names io0-io55. + +Valid values for xRX200 pin names: + Pinconf pins can be referenced via the names io0-io49. + +Valid values for xRX300 pin names: + Pinconf pins can be referenced via the names io0-io1,io3-io6,io8-io11, + io13-io19,io23-io27,io34-io36, + io42-io43,io48-io61. + Example: gpio: pinmux@E100B10 { - compatible = "lantiq,pinctrl-xway"; + compatible = "lantiq,danube-pinctrl"; pinctrl-names = "default"; pinctrl-0 = <&state_default>; diff --git a/Documentation/devicetree/bindings/pinctrl/pinctrl-mt65xx.txt b/Documentation/devicetree/bindings/pinctrl/pinctrl-mt65xx.txt index 0480bc31bfd7..9ffb0b276bb4 100644 --- a/Documentation/devicetree/bindings/pinctrl/pinctrl-mt65xx.txt +++ b/Documentation/devicetree/bindings/pinctrl/pinctrl-mt65xx.txt @@ -4,10 +4,11 @@ The Mediatek's Pin controller is used to control SoC pins. Required properties: - compatible: value should be one of the following. - (a) "mediatek,mt8135-pinctrl", compatible with mt8135 pinctrl. - (b) "mediatek,mt8173-pinctrl", compatible with mt8173 pinctrl. - (c) "mediatek,mt6397-pinctrl", compatible with mt6397 pinctrl. - (d) "mediatek,mt8127-pinctrl", compatible with mt8127 pinctrl. + "mediatek,mt2701-pinctrl", compatible with mt2701 pinctrl. + "mediatek,mt6397-pinctrl", compatible with mt6397 pinctrl. + "mediatek,mt8127-pinctrl", compatible with mt8127 pinctrl. + "mediatek,mt8135-pinctrl", compatible with mt8135 pinctrl. + "mediatek,mt8173-pinctrl", compatible with mt8173 pinctrl. - pins-are-numbered: Specify the subnodes are using numbered pinmux to specify pins. - gpio-controller : Marks the device node as a gpio controller. diff --git a/Documentation/devicetree/bindings/pinctrl/qcom,msm8996-pinctrl.txt b/Documentation/devicetree/bindings/pinctrl/qcom,msm8996-pinctrl.txt new file mode 100644 index 000000000000..e312a71b2f94 --- /dev/null +++ b/Documentation/devicetree/bindings/pinctrl/qcom,msm8996-pinctrl.txt @@ -0,0 +1,199 @@ +Qualcomm MSM8996 TLMM block + +This binding describes the Top Level Mode Multiplexer block found in the +MSM8996 platform. + +- compatible: + Usage: required + Value type: <string> + Definition: must be "qcom,msm8996-pinctrl" + +- reg: + Usage: required + Value type: <prop-encoded-array> + Definition: the base address and size of the TLMM register space. + +- interrupts: + Usage: required + Value type: <prop-encoded-array> + Definition: should specify the TLMM summary IRQ. + +- interrupt-controller: + Usage: required + Value type: <none> + Definition: identifies this node as an interrupt controller + +- #interrupt-cells: + Usage: required + Value type: <u32> + Definition: must be 2. Specifying the pin number and flags, as defined + in <dt-bindings/interrupt-controller/irq.h> + +- gpio-controller: + Usage: required + Value type: <none> + Definition: identifies this node as a gpio controller + +- #gpio-cells: + Usage: required + Value type: <u32> + Definition: must be 2. Specifying the pin number and flags, as defined + in <dt-bindings/gpio/gpio.h> + +Please refer to ../gpio/gpio.txt and ../interrupt-controller/interrupts.txt for +a general description of GPIO and interrupt bindings. + +Please refer to pinctrl-bindings.txt in this directory for details of the +common pinctrl bindings used by client devices, including the meaning of the +phrase "pin configuration node". + +The pin configuration nodes act as a container for an arbitrary number of +subnodes. Each of these subnodes represents some desired configuration for a +pin, a group, or a list of pins or groups. This configuration can include the +mux function to select on those pin(s)/group(s), and various pin configuration +parameters, such as pull-up, drive strength, etc. + + +PIN CONFIGURATION NODES: + +The name of each subnode is not important; all subnodes should be enumerated +and processed purely based on their content. + +Each subnode only affects those parameters that are explicitly listed. In +other words, a subnode that lists a mux function but no pin configuration +parameters implies no information about any pin configuration parameters. +Similarly, a pin subnode that describes a pullup parameter implies no +information about e.g. the mux function. + + +The following generic properties as defined in pinctrl-bindings.txt are valid +to specify in a pin configuration subnode: + +- pins: + Usage: required + Value type: <string-array> + Definition: List of gpio pins affected by the properties specified in + this subnode. + + Valid pins are: + gpio0-gpio149 + Supports mux, bias and drive-strength + + sdc1_clk, sdc1_cmd, sdc1_data sdc2_clk, sdc2_cmd, + sdc2_data sdc1_rclk + Supports bias and drive-strength + +- function: + Usage: required + Value type: <string> + Definition: Specify the alternative function to be configured for the + specified pins. Functions are only valid for gpio pins. + Valid values are: + + blsp_uart1, blsp_spi1, blsp_i2c1, blsp_uim1, atest_tsens, + bimc_dte1, dac_calib0, blsp_spi8, blsp_uart8, blsp_uim8, + qdss_cti_trig_out_b, bimc_dte0, dac_calib1, qdss_cti_trig_in_b, + dac_calib2, atest_tsens2, atest_usb1, blsp_spi10, blsp_uart10, + blsp_uim10, atest_bbrx1, atest_usb13, atest_bbrx0, atest_usb12, + mdp_vsync, edp_lcd, blsp_i2c10, atest_gpsadc1, atest_usb11, + atest_gpsadc0, edp_hot, atest_usb10, m_voc, dac_gpio, atest_char, + cam_mclk, pll_bypassnl, qdss_stm7, blsp_i2c8, qdss_tracedata_b, + pll_reset, qdss_stm6, qdss_stm5, qdss_stm4, atest_usb2, cci_i2c, + qdss_stm3, dac_calib3, atest_usb23, atest_char3, dac_calib4, + qdss_stm2, atest_usb22, atest_char2, qdss_stm1, dac_calib5, + atest_usb21, atest_char1, dbg_out, qdss_stm0, dac_calib6, + atest_usb20, atest_char0, dac_calib10, qdss_stm10, + qdss_cti_trig_in_a, cci_timer4, blsp_spi6, blsp_uart6, blsp_uim6, + blsp2_spi, qdss_stm9, qdss_cti_trig_out_a, dac_calib11, + qdss_stm8, cci_timer0, qdss_stm13, dac_calib7, cci_timer1, + qdss_stm12, dac_calib8, cci_timer2, blsp1_spi, qdss_stm11, + dac_calib9, cci_timer3, cci_async, dac_calib12, blsp_i2c6, + qdss_tracectl_a, dac_calib13, qdss_traceclk_a, dac_calib14, + dac_calib15, hdmi_rcv, dac_calib16, hdmi_cec, pwr_modem, + dac_calib17, hdmi_ddc, pwr_nav, dac_calib18, pwr_crypto, + dac_calib19, hdmi_hot, dac_calib20, dac_calib21, pci_e0, + dac_calib22, dac_calib23, dac_calib24, tsif1_sync, dac_calib25, + sd_write, tsif1_error, blsp_spi2, blsp_uart2, blsp_uim2, + qdss_cti, blsp_i2c2, blsp_spi3, blsp_uart3, blsp_uim3, blsp_i2c3, + uim3, blsp_spi9, blsp_uart9, blsp_uim9, blsp10_spi, blsp_i2c9, + blsp_spi7, blsp_uart7, blsp_uim7, qdss_tracedata_a, blsp_i2c7, + qua_mi2s, gcc_gp1_clk_a, ssc_irq, uim4, blsp_spi11, blsp_uart11, + blsp_uim11, gcc_gp2_clk_a, gcc_gp3_clk_a, blsp_i2c11, cri_trng0, + cri_trng1, cri_trng, qdss_stm18, pri_mi2s, qdss_stm17, blsp_spi4, + blsp_uart4, blsp_uim4, qdss_stm16, qdss_stm15, blsp_i2c4, + qdss_stm14, dac_calib26, spkr_i2s, audio_ref, lpass_slimbus, + isense_dbg, tsense_pwm1, tsense_pwm2, btfm_slimbus, ter_mi2s, + qdss_stm22, qdss_stm21, qdss_stm20, qdss_stm19, gcc_gp1_clk_b, + sec_mi2s, blsp_spi5, blsp_uart5, blsp_uim5, gcc_gp2_clk_b, + gcc_gp3_clk_b, blsp_i2c5, blsp_spi12, blsp_uart12, blsp_uim12, + qdss_stm25, qdss_stm31, blsp_i2c12, qdss_stm30, qdss_stm29, + tsif1_clk, qdss_stm28, tsif1_en, tsif1_data, sdc4_cmd, qdss_stm27, + qdss_traceclk_b, tsif2_error, sdc43, vfr_1, qdss_stm26, tsif2_clk, + sdc4_clk, qdss_stm24, tsif2_en, sdc42, qdss_stm23, qdss_tracectl_b, + sd_card, tsif2_data, sdc41, tsif2_sync, sdc40, mdp_vsync_p_b, + ldo_en, mdp_vsync_s_b, ldo_update, blsp11_uart_tx_b, blsp11_uart_rx_b, + blsp11_i2c_sda_b, prng_rosc, blsp11_i2c_scl_b, uim2, uim1, uim_batt, + pci_e2, pa_indicator, adsp_ext, ddr_bist, qdss_tracedata_11, + qdss_tracedata_12, modem_tsync, nav_dr, nav_pps, pci_e1, gsm_tx, + qspi_cs, ssbi2, ssbi1, mss_lte, qspi_clk, qspi0, qspi1, qspi2, qspi3, + gpio + +- bias-disable: + Usage: optional + Value type: <none> + Definition: The specified pins should be configued as no pull. + +- bias-pull-down: + Usage: optional + Value type: <none> + Definition: The specified pins should be configued as pull down. + +- bias-pull-up: + Usage: optional + Value type: <none> + Definition: The specified pins should be configued as pull up. + +- output-high: + Usage: optional + Value type: <none> + Definition: The specified pins are configured in output mode, driven + high. + Not valid for sdc pins. + +- output-low: + Usage: optional + Value type: <none> + Definition: The specified pins are configured in output mode, driven + low. + Not valid for sdc pins. + +- drive-strength: + Usage: optional + Value type: <u32> + Definition: Selects the drive strength for the specified pins, in mA. + Valid values are: 2, 4, 6, 8, 10, 12, 14 and 16 + +Example: + + tlmm: pinctrl@01010000 { + compatible = "qcom,msm8996-pinctrl"; + reg = <0x01010000 0x300000>; + interrupts = <0 208 0>; + gpio-controller; + #gpio-cells = <2>; + interrupt-controller; + #interrupt-cells = <2>; + + uart_console_active: uart_console_active { + mux { + pins = "gpio4", "gpio5"; + function = "blsp_uart8"; + }; + + config { + pins = "gpio4", "gpio5"; + drive-strength = <2>; + bias-disable; + }; + }; + }; diff --git a/Documentation/devicetree/bindings/pinctrl/qcom,pmic-gpio.txt b/Documentation/devicetree/bindings/pinctrl/qcom,pmic-gpio.txt index 1ae63c0acd40..a90c812ad642 100644 --- a/Documentation/devicetree/bindings/pinctrl/qcom,pmic-gpio.txt +++ b/Documentation/devicetree/bindings/pinctrl/qcom,pmic-gpio.txt @@ -14,6 +14,7 @@ PMIC's from Qualcomm. "qcom,pm8917-gpio" "qcom,pm8921-gpio" "qcom,pm8941-gpio" + "qcom,pm8994-gpio" "qcom,pma8084-gpio" - reg: @@ -79,6 +80,7 @@ to specify in a pin configuration subnode: gpio1-gpio38 for pm8917 gpio1-gpio44 for pm8921 gpio1-gpio36 for pm8941 + gpio1-gpio22 for pm8994 gpio1-gpio22 for pma8084 - function: diff --git a/Documentation/devicetree/bindings/pinctrl/qcom,pmic-mpp.txt b/Documentation/devicetree/bindings/pinctrl/qcom,pmic-mpp.txt index d7803a2a94e9..d74e631e10da 100644 --- a/Documentation/devicetree/bindings/pinctrl/qcom,pmic-mpp.txt +++ b/Documentation/devicetree/bindings/pinctrl/qcom,pmic-mpp.txt @@ -15,6 +15,7 @@ of PMIC's from Qualcomm. "qcom,pm8917-mpp", "qcom,pm8921-mpp", "qcom,pm8941-mpp", + "qcom,pm8994-mpp", "qcom,pma8084-mpp", - reg: diff --git a/Documentation/devicetree/bindings/pinctrl/rockchip,pinctrl.txt b/Documentation/devicetree/bindings/pinctrl/rockchip,pinctrl.txt index 391ef4be8d50..0cd701b1947f 100644 --- a/Documentation/devicetree/bindings/pinctrl/rockchip,pinctrl.txt +++ b/Documentation/devicetree/bindings/pinctrl/rockchip,pinctrl.txt @@ -21,7 +21,8 @@ defined as gpio sub-nodes of the pinmux controller. Required properties for iomux controller: - compatible: one of "rockchip,rk2928-pinctrl", "rockchip,rk3066a-pinctrl" "rockchip,rk3066b-pinctrl", "rockchip,rk3188-pinctrl" - "rockchip,rk3288-pinctrl", "rockchip,rk3368-pinctrl" + "rockchip,rk3228-pinctrl", "rockchip,rk3288-pinctrl" + "rockchip,rk3368-pinctrl" - rockchip,grf: phandle referencing a syscon providing the "general register files" diff --git a/Documentation/devicetree/bindings/pinctrl/samsung-pinctrl.txt b/Documentation/devicetree/bindings/pinctrl/samsung-pinctrl.txt index 9d2a995293e6..6db16b90873a 100644 --- a/Documentation/devicetree/bindings/pinctrl/samsung-pinctrl.txt +++ b/Documentation/devicetree/bindings/pinctrl/samsung-pinctrl.txt @@ -17,6 +17,7 @@ Required Properties: - "samsung,exynos4x12-pinctrl": for Exynos4x12 compatible pin-controller. - "samsung,exynos5250-pinctrl": for Exynos5250 compatible pin-controller. - "samsung,exynos5260-pinctrl": for Exynos5260 compatible pin-controller. + - "samsung,exynos5410-pinctrl": for Exynos5410 compatible pin-controller. - "samsung,exynos5420-pinctrl": for Exynos5420 compatible pin-controller. - "samsung,exynos7-pinctrl": for Exynos7 compatible pin-controller. diff --git a/Documentation/dvb/README.dvb-usb b/Documentation/dvb/README.dvb-usb index 8eb92264ee04..669dc6ce4330 100644 --- a/Documentation/dvb/README.dvb-usb +++ b/Documentation/dvb/README.dvb-usb @@ -45,7 +45,7 @@ Supported devices See the LinuxTV DVB Wiki at www.linuxtv.org for a complete list of cards/drivers/firmwares: -http://www.linuxtv.org/wiki/index.php/DVB_USB +https://linuxtv.org/wiki/index.php/DVB_USB 0. History & News: 2005-06-30 - added support for WideView WT-220U (Thanks to Steve Chang) @@ -121,7 +121,7 @@ working. Have a look at the Wikipage for the DVB-USB-drivers to find out, which firmware you need for your device: -http://www.linuxtv.org/wiki/index.php/DVB_USB +https://linuxtv.org/wiki/index.php/DVB_USB 1.2. Compiling diff --git a/Documentation/dvb/faq.txt b/Documentation/dvb/faq.txt index 97b1373f2428..a0be92012877 100644 --- a/Documentation/dvb/faq.txt +++ b/Documentation/dvb/faq.txt @@ -76,7 +76,7 @@ Some very frequently asked questions about linuxtv-dvb the TuxBox CVS many interesting DVB applications and the dBox2 DVB source - http://www.linuxtv.org/downloads/ + https://linuxtv.org/downloads DVB Swiss Army Knife library and utilities http://www.nenie.org/misc/mpsys/ diff --git a/Documentation/dvb/get_dvb_firmware b/Documentation/dvb/get_dvb_firmware index 91b43d2738c7..1a0a04125f71 100755 --- a/Documentation/dvb/get_dvb_firmware +++ b/Documentation/dvb/get_dvb_firmware @@ -152,7 +152,7 @@ sub tda10046lifeview { sub av7110 { my $sourcefile = "dvb-ttpci-01.fw-261d"; - my $url = "http://www.linuxtv.org/downloads/firmware/$sourcefile"; + my $url = "https://linuxtv.org/downloads/firmware/$sourcefile"; my $hash = "603431b6259715a8e88f376a53b64e2f"; my $outfile = "dvb-ttpci-01.fw"; @@ -303,7 +303,7 @@ sub vp7049 { } sub dibusb { - my $url = "http://www.linuxtv.org/downloads/firmware/dvb-usb-dibusb-5.0.0.11.fw"; + my $url = "https://linuxtv.org/downloads/firmware/dvb-usb-dibusb-5.0.0.11.fw"; my $outfile = "dvb-dibusb-5.0.0.11.fw"; my $hash = "fa490295a527360ca16dcdf3224ca243"; @@ -351,7 +351,7 @@ sub nxt2004 { sub or51211 { my $fwfile = "dvb-fe-or51211.fw"; - my $url = "http://linuxtv.org/downloads/firmware/$fwfile"; + my $url = "https://linuxtv.org/downloads/firmware/$fwfile"; my $hash = "d830949c771a289505bf9eafc225d491"; checkstandard(); @@ -364,7 +364,7 @@ sub or51211 { sub cx231xx { my $fwfile = "v4l-cx231xx-avcore-01.fw"; - my $url = "http://linuxtv.org/downloads/firmware/$fwfile"; + my $url = "https://linuxtv.org/downloads/firmware/$fwfile"; my $hash = "7d3bb956dc9df0eafded2b56ba57cc42"; checkstandard(); @@ -376,7 +376,7 @@ sub cx231xx { } sub cx18 { - my $url = "http://linuxtv.org/downloads/firmware/"; + my $url = "https://linuxtv.org/downloads/firmware/"; my %files = ( 'v4l-cx23418-apu.fw' => '588f081b562f5c653a3db1ad8f65939a', @@ -450,7 +450,7 @@ sub mpc718 { } sub cx23885 { - my $url = "http://linuxtv.org/downloads/firmware/"; + my $url = "https://linuxtv.org/downloads/firmware/"; my %files = ( 'v4l-cx23885-avcore-01.fw' => 'a9f8f5d901a7fb42f552e1ee6384f3bb', @@ -472,7 +472,7 @@ sub cx23885 { } sub pvrusb2 { - my $url = "http://linuxtv.org/downloads/firmware/"; + my $url = "https://linuxtv.org/downloads/firmware/"; my %files = ( 'v4l-cx25840.fw' => 'dadb79e9904fc8af96e8111d9cb59320', @@ -494,7 +494,7 @@ sub pvrusb2 { sub or51132_qam { my $fwfile = "dvb-fe-or51132-qam.fw"; - my $url = "http://linuxtv.org/downloads/firmware/$fwfile"; + my $url = "https://linuxtv.org/downloads/firmware/$fwfile"; my $hash = "7702e8938612de46ccadfe9b413cb3b5"; checkstandard(); @@ -507,7 +507,7 @@ sub or51132_qam { sub or51132_vsb { my $fwfile = "dvb-fe-or51132-vsb.fw"; - my $url = "http://linuxtv.org/downloads/firmware/$fwfile"; + my $url = "https://linuxtv.org/downloads/firmware/$fwfile"; my $hash = "c16208e02f36fc439a557ad4c613364a"; checkstandard(); @@ -519,7 +519,7 @@ sub or51132_vsb { } sub bluebird { - my $url = "http://www.linuxtv.org/download/dvb/firmware/dvb-usb-bluebird-01.fw"; + my $url = "https://linuxtv.org/download/dvb/firmware/dvb-usb-bluebird-01.fw"; my $outfile = "dvb-usb-bluebird-01.fw"; my $hash = "658397cb9eba9101af9031302671f49d"; @@ -677,7 +677,7 @@ sub drxk_hauppauge_hvr930c { } sub drxk_terratec_h5 { - my $url = "http://www.linuxtv.org/downloads/firmware/"; + my $url = "https://linuxtv.org/downloads/firmware/"; my $hash = "19000dada8e2741162ccc50cc91fa7f1"; my $fwfile = "dvb-usb-terratec-h5-drxk.fw"; diff --git a/Documentation/dvb/readme.txt b/Documentation/dvb/readme.txt index 0b0380c91990..89965041a266 100644 --- a/Documentation/dvb/readme.txt +++ b/Documentation/dvb/readme.txt @@ -2,12 +2,12 @@ Linux Digital Video Broadcast (DVB) subsystem ============================================= The main development site and CVS repository for these -drivers is http://linuxtv.org/. +drivers is https://linuxtv.org. The developer mailing list linux-dvb is also hosted there, -see http://linuxtv.org/lists.php. Please check -the archive http://linuxtv.org/pipermail/linux-dvb/ -and the Wiki http://linuxtv.org/wiki/ +see https://linuxtv.org/lists.php. Please check +the archive https://linuxtv.org/pipermail/linux-dvb/ +and the Wiki https://linuxtv.org/wiki/ before asking newbie questions on the list. API documentation, utilities and test/example programs @@ -16,7 +16,7 @@ are available as part of the old driver package for Linux 2.4 We plan to split this into separate packages, but it's not been done yet. -http://linuxtv.org/downloads/ +https://linuxtv.org/downloads/ What's inside this directory: diff --git a/Documentation/edac.txt b/Documentation/edac.txt index 80841a2d640c..f89cfd85ae13 100644 --- a/Documentation/edac.txt +++ b/Documentation/edac.txt @@ -1,9 +1,13 @@ EDAC - Error Detection And Correction ===================================== -"bluesmoke" was the name for this device driver when it was "out-of-tree" -and maintained at sourceforge.net. When it was pushed into 2.6.16 for the -first time, it was renamed to 'EDAC'. +"bluesmoke" was the name for this device driver when it +was "out-of-tree" and maintained at sourceforge.net - +bluesmoke.sourceforge.net. That site is mostly archaic now and can be +used only for historical purposes. + +When the subsystem was pushed into 2.6.16 for the first time, it was +renamed to 'EDAC'. PURPOSE ------- diff --git a/Documentation/fault-injection/notifier-error-inject.txt b/Documentation/fault-injection/notifier-error-inject.txt index 09adabef513f..83d3f4e43e91 100644 --- a/Documentation/fault-injection/notifier-error-inject.txt +++ b/Documentation/fault-injection/notifier-error-inject.txt @@ -10,6 +10,7 @@ modules that can be used to test the following notifiers. * PM notifier * Memory hotplug notifier * powerpc pSeries reconfig notifier + * Netdevice notifier CPU notifier error injection module ----------------------------------- @@ -87,6 +88,30 @@ Possible pSeries reconfig notifier events to be failed are: * PSERIES_DRCONF_MEM_ADD * PSERIES_DRCONF_MEM_REMOVE +Netdevice notifier error injection module +---------------------------------------------- +This feature is controlled through debugfs interface +/sys/kernel/debug/notifier-error-inject/netdev/actions/<notifier event>/error + +Netdevice notifier events which can be failed are: + + * NETDEV_REGISTER + * NETDEV_CHANGEMTU + * NETDEV_CHANGENAME + * NETDEV_PRE_UP + * NETDEV_PRE_TYPE_CHANGE + * NETDEV_POST_INIT + * NETDEV_PRECHANGEMTU + * NETDEV_PRECHANGEUPPER + * NETDEV_CHANGEUPPER + +Example: Inject netdevice mtu change error (-22 == -EINVAL) + + # cd /sys/kernel/debug/notifier-error-inject/netdev + # echo -22 > actions/NETDEV_CHANGEMTU/error + # ip link set eth0 mtu 1024 + RTNETLINK answers: Invalid argument + For more usage examples ----------------------- There are tools/testing/selftests using the notifier error injection features diff --git a/Documentation/features/seccomp/seccomp-filter/arch-support.txt b/Documentation/features/seccomp/seccomp-filter/arch-support.txt index 76d39d66a5d7..4f66ec133951 100644 --- a/Documentation/features/seccomp/seccomp-filter/arch-support.txt +++ b/Documentation/features/seccomp/seccomp-filter/arch-support.txt @@ -33,7 +33,7 @@ | sh: | TODO | | sparc: | TODO | | tile: | ok | - | um: | TODO | + | um: | ok | | unicore32: | TODO | | x86: | ok | | xtensa: | TODO | diff --git a/Documentation/features/time/irq-time-acct/arch-support.txt b/Documentation/features/time/irq-time-acct/arch-support.txt index e63316239938..4199ffecc0ff 100644 --- a/Documentation/features/time/irq-time-acct/arch-support.txt +++ b/Documentation/features/time/irq-time-acct/arch-support.txt @@ -9,7 +9,7 @@ | alpha: | .. | | arc: | TODO | | arm: | ok | - | arm64: | .. | + | arm64: | ok | | avr32: | TODO | | blackfin: | TODO | | c6x: | TODO | diff --git a/Documentation/filesystems/Locking b/Documentation/filesystems/Locking index 06d443450f21..619af9bfdcb3 100644 --- a/Documentation/filesystems/Locking +++ b/Documentation/filesystems/Locking @@ -50,8 +50,7 @@ prototypes: int (*rename2) (struct inode *, struct dentry *, struct inode *, struct dentry *, unsigned int); int (*readlink) (struct dentry *, char __user *,int); - const char *(*follow_link) (struct dentry *, void **); - void (*put_link) (struct inode *, void *); + const char *(*get_link) (struct dentry *, struct inode *, void **); void (*truncate) (struct inode *); int (*permission) (struct inode *, int, unsigned int); int (*get_acl)(struct inode *, int); @@ -83,8 +82,7 @@ rmdir: yes (both) (see below) rename: yes (all) (see below) rename2: yes (all) (see below) readlink: no -follow_link: no -put_link: no +get_link: no setattr: yes permission: no (may not block if called in rcu-walk mode) get_acl: no diff --git a/Documentation/filesystems/configfs/configfs.txt b/Documentation/filesystems/configfs/configfs.txt index af68efdbbfad..e5fe521eea1d 100644 --- a/Documentation/filesystems/configfs/configfs.txt +++ b/Documentation/filesystems/configfs/configfs.txt @@ -51,15 +51,27 @@ configfs tree is always there, whether mounted on /config or not. An item is created via mkdir(2). The item's attributes will also appear at this time. readdir(3) can determine what the attributes are, read(2) can query their default values, and write(2) can store new -values. Like sysfs, attributes should be ASCII text files, preferably -with only one value per file. The same efficiency caveats from sysfs -apply. Don't mix more than one attribute in one attribute file. - -Like sysfs, configfs expects write(2) to store the entire buffer at -once. When writing to configfs attributes, userspace processes should -first read the entire file, modify the portions they wish to change, and -then write the entire buffer back. Attribute files have a maximum size -of one page (PAGE_SIZE, 4096 on i386). +values. Don't mix more than one attribute in one attribute file. + +There are two types of configfs attributes: + +* Normal attributes, which similar to sysfs attributes, are small ASCII text +files, with a maximum size of one page (PAGE_SIZE, 4096 on i386). Preferably +only one value per file should be used, and the same caveats from sysfs apply. +Configfs expects write(2) to store the entire buffer at once. When writing to +normal configfs attributes, userspace processes should first read the entire +file, modify the portions they wish to change, and then write the entire +buffer back. + +* Binary attributes, which are somewhat similar to sysfs binary attributes, +but with a few slight changes to semantics. The PAGE_SIZE limitation does not +apply, but the whole binary item must fit in single kernel vmalloc'ed buffer. +The write(2) calls from user space are buffered, and the attributes' +write_bin_attribute method will be invoked on the final close, therefore it is +imperative for user-space to check the return code of close(2) in order to +verify that the operation finished successfully. +To avoid a malicious user OOMing the kernel, there's a per-binary attribute +maximum buffer value. When an item needs to be destroyed, remove it with rmdir(2). An item cannot be destroyed if any other item has a link to it (via @@ -171,6 +183,7 @@ among other things. For that, it needs a type. struct configfs_item_operations *ct_item_ops; struct configfs_group_operations *ct_group_ops; struct configfs_attribute **ct_attrs; + struct configfs_bin_attribute **ct_bin_attrs; }; The most basic function of a config_item_type is to define what @@ -201,6 +214,32 @@ be called whenever userspace asks for a read(2) on the attribute. If an attribute is writable and provides a ->store method, that method will be be called whenever userspace asks for a write(2) on the attribute. +[struct configfs_bin_attribute] + + struct configfs_attribute { + struct configfs_attribute cb_attr; + void *cb_private; + size_t cb_max_size; + }; + +The binary attribute is used when the one needs to use binary blob to +appear as the contents of a file in the item's configfs directory. +To do so add the binary attribute to the NULL-terminated array +config_item_type->ct_bin_attrs, and the item appears in configfs, the +attribute file will appear with the configfs_bin_attribute->cb_attr.ca_name +filename. configfs_bin_attribute->cb_attr.ca_mode specifies the file +permissions. +The cb_private member is provided for use by the driver, while the +cb_max_size member specifies the maximum amount of vmalloc buffer +to be used. + +If binary attribute is readable and the config_item provides a +ct_item_ops->read_bin_attribute() method, that method will be called +whenever userspace asks for a read(2) on the attribute. The converse +will happen for write(2). The reads/writes are bufferred so only a +single read/write will occur; the attributes' need not concern itself +with it. + [struct config_group] A config_item cannot live in a vacuum. The only way one can be created diff --git a/Documentation/filesystems/porting b/Documentation/filesystems/porting index f24d1b833957..0f88e6020487 100644 --- a/Documentation/filesystems/porting +++ b/Documentation/filesystems/porting @@ -504,3 +504,20 @@ in your dentry operations instead. [mandatory] __fd_install() & fd_install() can now sleep. Callers should not hold a spinlock or other resources that do not allow a schedule. +-- +[mandatory] + any symlink that might use page_follow_link_light/page_put_link() must + have inode_nohighmem(inode) called before anything might start playing with + its pagecache. +-- +[mandatory] + ->follow_link() is replaced with ->get_link(); same API, except that + * ->get_link() gets inode as a separate argument + * ->get_link() may be called in RCU mode - in that case NULL + dentry is passed +-- +[mandatory] + ->get_link() gets struct delayed_call *done now, and should do + set_delayed_call() where it used to set *cookie. + ->put_link() is gone - just give the destructor to set_delayed_call() + in ->get_link(). diff --git a/Documentation/filesystems/vfs.txt b/Documentation/filesystems/vfs.txt index 8c6f07ad373a..b02a7d598258 100644 --- a/Documentation/filesystems/vfs.txt +++ b/Documentation/filesystems/vfs.txt @@ -350,8 +350,8 @@ struct inode_operations { int (*rename2) (struct inode *, struct dentry *, struct inode *, struct dentry *, unsigned int); int (*readlink) (struct dentry *, char __user *,int); - const char *(*follow_link) (struct dentry *, void **); - void (*put_link) (struct inode *, void *); + const char *(*get_link) (struct dentry *, struct inode *, + struct delayed_call *); int (*permission) (struct inode *, int); int (*get_acl)(struct inode *, int); int (*setattr) (struct dentry *, struct iattr *); @@ -434,20 +434,19 @@ otherwise noted. readlink: called by the readlink(2) system call. Only required if you want to support reading symbolic links - follow_link: called by the VFS to follow a symbolic link to the + get_link: called by the VFS to follow a symbolic link to the inode it points to. Only required if you want to support symbolic links. This method returns the symlink body to traverse (and possibly resets the current position with nd_jump_link()). If the body won't go away until the inode is gone, nothing else is needed; if it needs to be otherwise - pinned, the data needed to release whatever we'd grabbed - is to be stored in void * variable passed by address to - follow_link() instance. - - put_link: called by the VFS to release resources allocated by - follow_link(). The cookie stored by follow_link() is passed - to this method as the last parameter; only called when - cookie isn't NULL. + pinned, arrange for its release by having get_link(..., ..., done) + do set_delayed_call(done, destructor, argument). + In that case destructor(argument) will be called once VFS is + done with the body you've returned. + May be called in RCU mode; that is indicated by NULL dentry + argument. If request can't be handled without leaving RCU mode, + have it return ERR_PTR(-ECHILD). permission: called by the VFS to check for access rights on a POSIX-like filesystem. diff --git a/Documentation/hwmon/htu21 b/Documentation/hwmon/htu21 deleted file mode 100644 index f39a215fb6ae..000000000000 --- a/Documentation/hwmon/htu21 +++ /dev/null @@ -1,46 +0,0 @@ -Kernel driver htu21 -=================== - -Supported chips: - * Measurement Specialties HTU21D - Prefix: 'htu21' - Addresses scanned: none - Datasheet: Publicly available at the Measurement Specialties website - http://www.meas-spec.com/downloads/HTU21D.pdf - - -Author: - William Markezana <william.markezana@meas-spec.com> - -Description ------------ - -The HTU21D is a humidity and temperature sensor in a DFN package of -only 3 x 3 mm footprint and 0.9 mm height. - -The devices communicate with the I2C protocol. All sensors are set to the -same I2C address 0x40, so an entry with I2C_BOARD_INFO("htu21", 0x40) can -be used in the board setup code. - -This driver does not auto-detect devices. You will have to instantiate the -devices explicitly. Please see Documentation/i2c/instantiating-devices -for details. - -sysfs-Interface ---------------- - -temp1_input - temperature input -humidity1_input - humidity input - -Notes ------ - -The driver uses the default resolution settings of 12 bit for humidity and 14 -bit for temperature, which results in typical measurement times of 11 ms for -humidity and 44 ms for temperature. To keep self heating below 0.1 degree -Celsius, the device should not be active for more than 10% of the time. For -this reason, the driver performs no more than two measurements per second and -reports cached information if polled more frequently. - -Different resolutions, the on-chip heater, using the CRC checksum and reading -the serial number are not supported yet. diff --git a/Documentation/hwmon/ltc3815 b/Documentation/hwmon/ltc3815 new file mode 100644 index 000000000000..eb7db2d13587 --- /dev/null +++ b/Documentation/hwmon/ltc3815 @@ -0,0 +1,61 @@ +Kernel driver ltc3815 +===================== + +Supported chips: + * Linear Technology LTC3815 + Prefix: 'ltc3815' + Addresses scanned: - + Datasheet: http://www.linear.com/product/ltc3815 + +Author: Guenter Roeck <linux@roeck-us.net> + + +Description +----------- + +LTC3815 is a Monolithic Synchronous DC/DC Step-Down Converter. + + +Usage Notes +----------- + +This driver does not probe for PMBus devices. You will have to instantiate +devices explicitly. + +Example: the following commands will load the driver for an LTC3815 +at address 0x20 on I2C bus #1: + +# modprobe ltc3815 +# echo ltc3815 0x20 > /sys/bus/i2c/devices/i2c-1/new_device + + +Sysfs attributes +---------------- + +in1_label "vin" +in1_input Measured input voltage. +in1_alarm Input voltage alarm. +in1_highest Highest input voltage. +in1_reset_history Reset input voltage history. + +in2_label "vout1". +in2_input Measured output voltage. +in2_alarm Output voltage alarm. +in2_highest Highest output voltage. +in2_reset_history Reset output voltage history. + +temp1_input Measured chip temperature. +temp1_alarm Temperature alarm. +temp1_highest Highest measured temperature. +temp1_reset_history Reset temperature history. + +curr1_label "iin". +curr1_input Measured input current. +curr1_highest Highest input current. +curr1_reset_history Reset input current history. + +curr2_label "iout1". +curr2_input Measured output current. +curr2_alarm Output current alarm. +curr2_highest Highest output current. +curr2_reset_history Reset output current history. diff --git a/Documentation/kernel-parameters.txt b/Documentation/kernel-parameters.txt index fd333fbfd2ab..8b9fc655ddc7 100644 --- a/Documentation/kernel-parameters.txt +++ b/Documentation/kernel-parameters.txt @@ -472,6 +472,15 @@ bytes respectively. Such letter suffixes can also be entirely omitted. Change the amount of debugging information output when initialising the APIC and IO-APIC components. + apic_extnmi= [APIC,X86] External NMI delivery setting + Format: { bsp (default) | all | none } + bsp: External NMI is delivered only to CPU 0 + all: External NMIs are broadcast to all CPUs as a + backup of CPU 0 + none: External NMI is masked for all CPUs. This is + useful so that a dump capture kernel won't be + shot down by NMI + autoconf= [IPV6] See Documentation/networking/ipv6.txt. @@ -3294,18 +3303,35 @@ bytes respectively. Such letter suffixes can also be entirely omitted. rcutorture.verbose= [KNL] Enable additional printk() statements. + rcupdate.rcu_cpu_stall_suppress= [KNL] + Suppress RCU CPU stall warning messages. + + rcupdate.rcu_cpu_stall_timeout= [KNL] + Set timeout for RCU CPU stall warning messages. + rcupdate.rcu_expedited= [KNL] Use expedited grace-period primitives, for example, synchronize_rcu_expedited() instead of synchronize_rcu(). This reduces latency, but can increase CPU utilization, degrade real-time latency, and degrade energy efficiency. - - rcupdate.rcu_cpu_stall_suppress= [KNL] - Suppress RCU CPU stall warning messages. - - rcupdate.rcu_cpu_stall_timeout= [KNL] - Set timeout for RCU CPU stall warning messages. + No effect on CONFIG_TINY_RCU kernels. + + rcupdate.rcu_normal= [KNL] + Use only normal grace-period primitives, + for example, synchronize_rcu() instead of + synchronize_rcu_expedited(). This improves + real-time latency, CPU utilization, and + energy efficiency, but can expose users to + increased grace-period latency. This parameter + overrides rcupdate.rcu_expedited. No effect on + CONFIG_TINY_RCU kernels. + + rcupdate.rcu_normal_after_boot= [KNL] + Once boot has completed (that is, after + rcu_end_inkernel_boot() has been invoked), use + only normal grace-period primitives. No effect + on CONFIG_TINY_RCU kernels. rcupdate.rcu_task_stall_timeout= [KNL] Set timeout in jiffies for RCU task stall warning @@ -4118,6 +4144,15 @@ bytes respectively. Such letter suffixes can also be entirely omitted. or other driver-specific files in the Documentation/watchdog/ directory. + workqueue.watchdog_thresh= + If CONFIG_WQ_WATCHDOG is configured, workqueue can + warn stall conditions and dump internal state to + help debugging. 0 disables workqueue stall + detection; otherwise, it's the stall threshold + duration in seconds. The default value is 30 and + it can be updated at runtime by writing to the + corresponding sysfs file. + workqueue.disable_numa By default, all work items queued to unbound workqueues are affine to the NUMA nodes they're diff --git a/Documentation/leds/leds-class.txt b/Documentation/leds/leds-class.txt index 62261c04060a..d406d98339b2 100644 --- a/Documentation/leds/leds-class.txt +++ b/Documentation/leds/leds-class.txt @@ -52,6 +52,19 @@ above leaves scope for further attributes should they be needed. If sections of the name don't apply, just leave that section blank. +Brightness setting API +====================== + +LED subsystem core exposes following API for setting brightness: + + - led_set_brightness : it is guaranteed not to sleep, passing LED_OFF stops + blinking, + - led_set_brightness_sync : for use cases when immediate effect is desired - + it can block the caller for the time required for accessing + device registers and can sleep, passing LED_OFF stops hardware + blinking, returns -EBUSY if software blink fallback is enabled. + + Hardware accelerated blink of LEDs ================================== diff --git a/Documentation/memory-barriers.txt b/Documentation/memory-barriers.txt index aef9487303d0..a61be39c7b51 100644 --- a/Documentation/memory-barriers.txt +++ b/Documentation/memory-barriers.txt @@ -194,7 +194,7 @@ There are some minimal guarantees that may be expected of a CPU: (*) On any given CPU, dependent memory accesses will be issued in order, with respect to itself. This means that for: - WRITE_ONCE(Q, P); smp_read_barrier_depends(); D = READ_ONCE(*Q); + Q = READ_ONCE(P); smp_read_barrier_depends(); D = READ_ONCE(*Q); the CPU will issue the following memory operations: @@ -202,9 +202,9 @@ There are some minimal guarantees that may be expected of a CPU: and always in that order. On most systems, smp_read_barrier_depends() does nothing, but it is required for DEC Alpha. The READ_ONCE() - and WRITE_ONCE() are required to prevent compiler mischief. Please - note that you should normally use something like rcu_dereference() - instead of open-coding smp_read_barrier_depends(). + is required to prevent compiler mischief. Please note that you + should normally use something like rcu_dereference() instead of + open-coding smp_read_barrier_depends(). (*) Overlapping loads and stores within a particular CPU will appear to be ordered within that CPU. This means that for: @@ -1673,8 +1673,8 @@ There are some more advanced barrier functions: (*) smp_store_mb(var, value) This assigns the value to the variable and then inserts a full memory - barrier after it, depending on the function. It isn't guaranteed to - insert anything more than a compiler barrier in a UP compilation. + barrier after it. It isn't guaranteed to insert anything more than a + compiler barrier in a UP compilation. (*) smp_mb__before_atomic(); diff --git a/Documentation/networking/batman-adv.txt b/Documentation/networking/batman-adv.txt index 58e49042fc20..ff23b755f5e4 100644 --- a/Documentation/networking/batman-adv.txt +++ b/Documentation/networking/batman-adv.txt @@ -115,14 +115,17 @@ The "bat0" interface can be used like any other regular inter- face. It needs an IP address which can be either statically con- figured or dynamically (by using DHCP or similar services): -# NodeA: ifconfig bat0 192.168.0.1 -# NodeB: ifconfig bat0 192.168.0.2 +# NodeA: ip link set up dev bat0 +# NodeA: ip addr add 192.168.0.1/24 dev bat0 + +# NodeB: ip link set up dev bat0 +# NodeB: ip addr add 192.168.0.2/24 dev bat0 # NodeB: ping 192.168.0.1 Note: In order to avoid problems remove all IP addresses previ- ously assigned to interfaces now used by batman advanced, e.g. -# ifconfig eth0 0.0.0.0 +# ip addr flush dev eth0 LOGGING/DEBUGGING diff --git a/Documentation/networking/ip-sysctl.txt b/Documentation/networking/ip-sysctl.txt index 2ea4c45cf1c8..ceb44a095a27 100644 --- a/Documentation/networking/ip-sysctl.txt +++ b/Documentation/networking/ip-sysctl.txt @@ -335,6 +335,14 @@ tcp_keepalive_intvl - INTEGER after probes started. Default value: 75sec i.e. connection will be aborted after ~11 minutes of retries. +tcp_l3mdev_accept - BOOLEAN + Enables child sockets to inherit the L3 master device index. + Enabling this option allows a "global" listen socket to work + across L3 master domains (e.g., VRFs) with connected sockets + derived from the listen socket to be bound to the L3 domain in + which the packets originated. Only valid when the kernel was + compiled with CONFIG_NET_L3_MASTER_DEV. + tcp_low_latency - BOOLEAN If set, the TCP stack makes decisions that prefer lower latency as opposed to higher throughput. By default, this @@ -1723,6 +1731,25 @@ addip_enable - BOOLEAN Default: 0 +pf_enable - INTEGER + Enable or disable pf (pf is short for potentially failed) state. A value + of pf_retrans > path_max_retrans also disables pf state. That is, one of + both pf_enable and pf_retrans > path_max_retrans can disable pf state. + Since pf_retrans and path_max_retrans can be changed by userspace + application, sometimes user expects to disable pf state by the value of + pf_retrans > path_max_retrans, but occasionally the value of pf_retrans + or path_max_retrans is changed by the user application, this pf state is + enabled. As such, it is necessary to add this to dynamically enable + and disable pf state. See: + https://datatracker.ietf.org/doc/draft-ietf-tsvwg-sctp-failover for + details. + + 1: Enable pf. + + 0: Disable pf. + + Default: 1 + addip_noauth_enable - BOOLEAN Dynamic Address Reconfiguration (ADD-IP) requires the use of authentication to protect the operations of adding or removing new @@ -1799,7 +1826,9 @@ pf_retrans - INTEGER having to reduce path_max_retrans to a very low value. See: http://www.ietf.org/id/draft-nishida-tsvwg-sctp-failover-05.txt for details. Note also that a value of pf_retrans > path_max_retrans - disables this feature + disables this feature. Since both pf_retrans and path_max_retrans can + be changed by userspace application, a variable pf_enable is used to + disable pf state. Default: 0 diff --git a/Documentation/networking/switchdev.txt b/Documentation/networking/switchdev.txt index 91994134efca..fad63136ee3e 100644 --- a/Documentation/networking/switchdev.txt +++ b/Documentation/networking/switchdev.txt @@ -304,8 +304,12 @@ certain netdevs from flooding unicast traffic for which there is no FDB entry. IGMP Snooping ^^^^^^^^^^^^^ -XXX: complete this section - +In order to support IGMP snooping, the port netdevs should trap to the bridge +driver all IGMP join and leave messages. +The bridge multicast module will notify port netdevs on every multicast group +changed whether it is static configured or dynamically joined/leave. +The hardware implementation should be forwarding all registered multicast +traffic groups only to the configured ports. L3 Routing Offload ------------------ diff --git a/Documentation/power/pci.txt b/Documentation/power/pci.txt index b0e911e0e8f5..44558882aa60 100644 --- a/Documentation/power/pci.txt +++ b/Documentation/power/pci.txt @@ -999,7 +999,7 @@ from its probe routine to make runtime PM work for the device. It is important to remember that the driver's runtime_suspend() callback may be executed right after the usage counter has been decremented, because -user space may already have cuased the pm_runtime_allow() helper function +user space may already have caused the pm_runtime_allow() helper function unblocking the runtime PM of the device to run via sysfs, so the driver must be prepared to cope with that. diff --git a/Documentation/power/runtime_pm.txt b/Documentation/power/runtime_pm.txt index 0784bc3a2ab5..7328cf85236c 100644 --- a/Documentation/power/runtime_pm.txt +++ b/Documentation/power/runtime_pm.txt @@ -371,6 +371,12 @@ drivers/base/power/runtime.c and include/linux/pm_runtime.h: - increment the device's usage counter, run pm_runtime_resume(dev) and return its result + int pm_runtime_get_if_in_use(struct device *dev); + - return -EINVAL if 'power.disable_depth' is nonzero; otherwise, if the + runtime PM status is RPM_ACTIVE and the runtime PM usage counter is + nonzero, increment the counter and return 1; otherwise return 0 without + changing the counter + void pm_runtime_put_noidle(struct device *dev); - decrement the device's usage counter diff --git a/Documentation/printk-formats.txt b/Documentation/printk-formats.txt index b784c270105f..6389551bbad6 100644 --- a/Documentation/printk-formats.txt +++ b/Documentation/printk-formats.txt @@ -250,6 +250,12 @@ dentry names: Passed by reference. +block_device names: + + %pg sda, sda1 or loop0p1 + + For printing name of block_device pointers. + struct va_format: %pV diff --git a/Documentation/sysctl/kernel.txt b/Documentation/sysctl/kernel.txt index af70d1541d3a..73c6b1ef0e84 100644 --- a/Documentation/sysctl/kernel.txt +++ b/Documentation/sysctl/kernel.txt @@ -551,6 +551,21 @@ the recommended setting is 60. ============================================================== +panic_on_io_nmi: + +Controls the kernel's behavior when a CPU receives an NMI caused by +an IO error. + +0: try to continue operation (default) + +1: panic immediately. The IO error triggered an NMI. This indicates a + serious system condition which could result in IO data corruption. + Rather than continuing, panicking might be a better choice. Some + servers issue this sort of NMI when the dump button is pushed, + and you can use this option to take a crash dump. + +============================================================== + panic_on_oops: Controls the kernel's behaviour when an oops or BUG is encountered. diff --git a/Documentation/trace/events-msr.txt b/Documentation/trace/events-msr.txt new file mode 100644 index 000000000000..78c383bf06aa --- /dev/null +++ b/Documentation/trace/events-msr.txt @@ -0,0 +1,37 @@ + +The x86 kernel supports tracing most MSR (Model Specific Register) accesses. +To see the definition of the MSRs on Intel systems please see the SDM +at http://www.intel.com/sdm (Volume 3) + +Available trace points: + +/sys/kernel/debug/tracing/events/msr/ + +Trace MSR reads + +read_msr + +msr: MSR number +val: Value written +failed: 1 if the access failed, otherwise 0 + + +Trace MSR writes + +write_msr + +msr: MSR number +val: Value written +failed: 1 if the access failed, otherwise 0 + + +Trace RDPMC in kernel + +rdpmc + +The trace data can be post processed with the postprocess/decode_msr.py script + +cat /sys/kernel/debug/tracing/trace | decode_msr.py /usr/src/linux/include/asm/msr-index.h + +to add symbolic MSR names. + diff --git a/Documentation/trace/postprocess/decode_msr.py b/Documentation/trace/postprocess/decode_msr.py new file mode 100644 index 000000000000..0ab40e0db580 --- /dev/null +++ b/Documentation/trace/postprocess/decode_msr.py @@ -0,0 +1,37 @@ +#!/usr/bin/python +# add symbolic names to read_msr / write_msr in trace +# decode_msr msr-index.h < trace +import sys +import re + +msrs = dict() + +with open(sys.argv[1] if len(sys.argv) > 1 else "msr-index.h", "r") as f: + for j in f: + m = re.match(r'#define (MSR_\w+)\s+(0x[0-9a-fA-F]+)', j) + if m: + msrs[int(m.group(2), 16)] = m.group(1) + +extra_ranges = ( + ( "MSR_LASTBRANCH_%d_FROM_IP", 0x680, 0x69F ), + ( "MSR_LASTBRANCH_%d_TO_IP", 0x6C0, 0x6DF ), + ( "LBR_INFO_%d", 0xdc0, 0xddf ), +) + +for j in sys.stdin: + m = re.search(r'(read|write)_msr:\s+([0-9a-f]+)', j) + if m: + r = None + num = int(m.group(2), 16) + if num in msrs: + r = msrs[num] + else: + for er in extra_ranges: + if er[1] <= num <= er[2]: + r = er[0] % (num - er[1],) + break + if r: + j = j.replace(" " + m.group(2), " " + r + "(" + m.group(2) + ")") + print j, + + diff --git a/Documentation/video4linux/API.html b/Documentation/video4linux/API.html index 256f8efa992c..eaf948cf1ae7 100644 --- a/Documentation/video4linux/API.html +++ b/Documentation/video4linux/API.html @@ -9,7 +9,7 @@ <table border="0"> <tr> <td> - <a href="http://linuxtv.org/downloads/legacy/video4linux/API/V4L1_API.html">V4L original API</a> + <a href="https://linuxtv.org/downloads/legacy/video4linux/API/V4L1_API.html">V4L original API</a> </td> <td> Obsoleted by V4L2 API diff --git a/Documentation/video4linux/CARDLIST.em28xx b/Documentation/video4linux/CARDLIST.em28xx index 9e57ce43c4f4..67209998a439 100644 --- a/Documentation/video4linux/CARDLIST.em28xx +++ b/Documentation/video4linux/CARDLIST.em28xx @@ -41,8 +41,8 @@ 40 -> Plextor ConvertX PX-TV100U (em2861) [093b:a005] 41 -> Kworld 350 U DVB-T (em2870) [eb1a:e350] 42 -> Kworld 355 U DVB-T (em2870) [eb1a:e355,eb1a:e357,eb1a:e359] - 43 -> Terratec Cinergy T XS (em2870) [0ccd:0043] - 44 -> Terratec Cinergy T XS (MT2060) (em2870) + 43 -> Terratec Cinergy T XS (em2870) + 44 -> Terratec Cinergy T XS (MT2060) (em2870) [0ccd:0043] 45 -> Pinnacle PCTV DVB-T (em2870) 46 -> Compro, VideoMate U3 (em2870) [185b:2870] 47 -> KWorld DVB-T 305U (em2880) [eb1a:e305] diff --git a/Documentation/video4linux/fimc.txt b/Documentation/video4linux/fimc.txt index e0c6b8bc4743..4fab231be52e 100644 --- a/Documentation/video4linux/fimc.txt +++ b/Documentation/video4linux/fimc.txt @@ -58,7 +58,7 @@ Not currently supported: 4.1. Media device interface The driver supports Media Controller API as defined at -http://linuxtv.org/downloads/v4l-dvb-apis/media_common.html +https://linuxtv.org/downloads/v4l-dvb-apis/media_common.html The media device driver name is "SAMSUNG S5P FIMC". The purpose of this interface is to allow changing assignment of FIMC instances @@ -83,11 +83,11 @@ undefined behaviour. 4.3. Capture video node The driver supports V4L2 Video Capture Interface as defined at: -http://linuxtv.org/downloads/v4l-dvb-apis/devices.html +https://linuxtv.org/downloads/v4l-dvb-apis/devices.html At the capture and mem-to-mem video nodes only the multi-planar API is supported. For more details see: -http://linuxtv.org/downloads/v4l-dvb-apis/planar-apis.html +https://linuxtv.org/downloads/v4l-dvb-apis/planar-apis.html 4.4. Camera capture subdevs diff --git a/Documentation/video4linux/omap4_camera.txt b/Documentation/video4linux/omap4_camera.txt index 25d9b40a4651..a6734aa77242 100644 --- a/Documentation/video4linux/omap4_camera.txt +++ b/Documentation/video4linux/omap4_camera.txt @@ -47,7 +47,7 @@ Tested platforms File list --------- drivers/staging/media/omap4iss/ -include/media/omap4iss.h +include/linux/platform_data/media/omap4iss.h References ---------- diff --git a/Documentation/video4linux/si4713.txt b/Documentation/video4linux/si4713.txt index 2e7392a4fee1..2ddc6b095a76 100644 --- a/Documentation/video4linux/si4713.txt +++ b/Documentation/video4linux/si4713.txt @@ -157,7 +157,7 @@ int main (int argc, char *argv[]) } The struct si4713_rnl and SI4713_IOC_MEASURE_RNL are defined under -include/media/si4713.h. +include/linux/platform_data/media/si4713.h. Stereo/Mono and RDS subchannels =============================== diff --git a/Documentation/video4linux/v4l2-pci-skeleton.c b/Documentation/video4linux/v4l2-pci-skeleton.c index 95ae82860092..79af0c041056 100644 --- a/Documentation/video4linux/v4l2-pci-skeleton.c +++ b/Documentation/video4linux/v4l2-pci-skeleton.c @@ -163,11 +163,10 @@ static irqreturn_t skeleton_irq(int irq, void *dev_id) * minimum number: many DMA engines need a minimum of 2 buffers in the * queue and you need to have another available for userspace processing. */ -static int queue_setup(struct vb2_queue *vq, const void *parg, +static int queue_setup(struct vb2_queue *vq, unsigned int *nbuffers, unsigned int *nplanes, unsigned int sizes[], void *alloc_ctxs[]) { - const struct v4l2_format *fmt = parg; struct skeleton *skel = vb2_get_drv_priv(vq); skel->field = skel->format.field; @@ -183,12 +182,12 @@ static int queue_setup(struct vb2_queue *vq, const void *parg, if (vq->num_buffers + *nbuffers < 3) *nbuffers = 3 - vq->num_buffers; + alloc_ctxs[0] = skel->alloc_ctx; - if (fmt && fmt->fmt.pix.sizeimage < skel->format.sizeimage) - return -EINVAL; + if (*nplanes) + return sizes[0] < skel->format.sizeimage ? -EINVAL : 0; *nplanes = 1; - sizes[0] = fmt ? fmt->fmt.pix.sizeimage : skel->format.sizeimage; - alloc_ctxs[0] = skel->alloc_ctx; + sizes[0] = skel->format.sizeimage; return 0; } @@ -509,7 +508,7 @@ static int skeleton_s_dv_timings(struct file *file, void *_fh, return -EINVAL; /* Return 0 if the new timings are the same as the current timings. */ - if (v4l2_match_dv_timings(timings, &skel->timings, 0)) + if (v4l2_match_dv_timings(timings, &skel->timings, 0, false)) return 0; /* diff --git a/Documentation/virtual/kvm/api.txt b/Documentation/virtual/kvm/api.txt index 092ee9fbaf2b..053f613fc9a9 100644 --- a/Documentation/virtual/kvm/api.txt +++ b/Documentation/virtual/kvm/api.txt @@ -1451,6 +1451,7 @@ struct kvm_irq_routing_entry { struct kvm_irq_routing_irqchip irqchip; struct kvm_irq_routing_msi msi; struct kvm_irq_routing_s390_adapter adapter; + struct kvm_irq_routing_hv_sint hv_sint; __u32 pad[8]; } u; }; @@ -1459,6 +1460,7 @@ struct kvm_irq_routing_entry { #define KVM_IRQ_ROUTING_IRQCHIP 1 #define KVM_IRQ_ROUTING_MSI 2 #define KVM_IRQ_ROUTING_S390_ADAPTER 3 +#define KVM_IRQ_ROUTING_HV_SINT 4 No flags are specified so far, the corresponding field must be set to zero. @@ -1482,6 +1484,10 @@ struct kvm_irq_routing_s390_adapter { __u32 adapter_id; }; +struct kvm_irq_routing_hv_sint { + __u32 vcpu; + __u32 sint; +}; 4.53 KVM_ASSIGN_SET_MSIX_NR (deprecated) @@ -3331,6 +3337,28 @@ the userspace IOAPIC should process the EOI and retrigger the interrupt if it is still asserted. Vector is the LAPIC interrupt vector for which the EOI was received. + struct kvm_hyperv_exit { +#define KVM_EXIT_HYPERV_SYNIC 1 + __u32 type; + union { + struct { + __u32 msr; + __u64 control; + __u64 evt_page; + __u64 msg_page; + } synic; + } u; + }; + /* KVM_EXIT_HYPERV */ + struct kvm_hyperv_exit hyperv; +Indicates that the VCPU exits into userspace to process some tasks +related to Hyper-V emulation. +Valid values for 'type' are: + KVM_EXIT_HYPERV_SYNIC -- synchronously notify user-space about +Hyper-V SynIC state change. Notification is used to remap SynIC +event/message pages and to enable/disable SynIC messages/events processing +in userspace. + /* Fix the size of the union. */ char padding[256]; }; @@ -3685,3 +3713,16 @@ available, means that that the kernel has an implementation of the H_RANDOM hypercall backed by a hardware random-number generator. If present, the kernel H_RANDOM handler can be enabled for guest use with the KVM_CAP_PPC_ENABLE_HCALL capability. + +8.2 KVM_CAP_HYPERV_SYNIC + +Architectures: x86 +This capability, if KVM_CHECK_EXTENSION indicates that it is +available, means that that the kernel has an implementation of the +Hyper-V Synthetic interrupt controller(SynIC). Hyper-V SynIC is +used to support Windows Hyper-V based guest paravirt drivers(VMBus). + +In order to use SynIC, it has to be activated by setting this +capability via KVM_ENABLE_CAP ioctl on the vcpu fd. Note that this +will disable the use of APIC hardware virtualization even if supported +by the CPU, as it's incompatible with SynIC auto-EOI behavior. diff --git a/Documentation/virtual/kvm/devices/vm.txt b/Documentation/virtual/kvm/devices/vm.txt index 2d09d1ed86d0..f083a168eb35 100644 --- a/Documentation/virtual/kvm/devices/vm.txt +++ b/Documentation/virtual/kvm/devices/vm.txt @@ -37,7 +37,8 @@ Returns: -EFAULT if the given address is not accessible Allows userspace to query the actual limit and set a new limit for the maximum guest memory size. The limit will be rounded up to 2048 MB, 4096 GB, 8192 TB respectively, as this limit is governed by -the number of page table levels. +the number of page table levels. In the case that there is no limit we will set +the limit to KVM_S390_NO_MEM_LIMIT (U64_MAX). 2. GROUP: KVM_S390_VM_CPU_MODEL Architectures: s390 diff --git a/Documentation/virtual/kvm/mmu.txt b/Documentation/virtual/kvm/mmu.txt index 3a4d681c3e98..daf9c0f742d2 100644 --- a/Documentation/virtual/kvm/mmu.txt +++ b/Documentation/virtual/kvm/mmu.txt @@ -203,10 +203,10 @@ Shadow pages contain the following information: page cannot be destroyed. See role.invalid. parent_ptes: The reverse mapping for the pte/ptes pointing at this page's spt. If - parent_ptes bit 0 is zero, only one spte points at this pages and + parent_ptes bit 0 is zero, only one spte points at this page and parent_ptes points at this single spte, otherwise, there exists multiple sptes pointing at this page and (parent_ptes & ~0x1) points at a data - structure with a list of parent_ptes. + structure with a list of parent sptes. unsync: If true, then the translations in this page may not match the guest's translation. This is equivalent to the state of the tlb when a pte is |
