Merge tag 'xfs-6.8-merge-3' of git://git.kernel.org/pub/scm/fs/xfs/xfs-linux

Pull xfs updates from Chandan Babu:
 "New features/functionality:
   - Online repair:
       - Reserve disk space for online repairs
       - Fix misinteraction between the AIL and btree bulkloader because
         of which the bulk load fails to queue a buffer for writeback if
         it happens to be on the AIL list
       - Prevent transaction reservation overflows when reaping blocks
         during online repair
       - Whenever possible, bulkloader now copies multiple records into
         a block
       - Support repairing of
           1. Per-AG free space, inode and refcount btrees
           2. Ondisk inodes
           3. File data and attribute fork mappings
       - Verify the contents of
           1. Inode and data fork of realtime bitmap file
           2. Quota files
   - Introduce MF_MEM_PRE_REMOVE. This will be used to notify tasks
     about a pmem device being removed

  Bug fixes:
   - Fix memory leak of recovered attri intent items
   - Fix UAF during log intent recovery
   - Fix realtime geometry integer overflows
   - Prevent scrub from live locking in xchk_iget
   - Prevent fs shutdown when removing files during low free disk space
   - Prevent transaction reservation overflow when extending an RT
     device
   - Prevent incorrect warning from being printed when extending a
     filesystem
   - Fix an off-by-one error in xreap_agextent_binval
   - Serialize access to perag radix tree during deletion operation
   - Fix perag memory leak during growfs
   - Allow allocation of minlen realtime extent when the maximum sized
     realtime free extent is minlen in size

  Cleanups:
   - Remove duplicate boilerplate code spread across functionality
     associated with different log items
   - Cleanup resblks interfaces
   - Pass defer ops pointer to defer helpers instead of an enum
   - Initialize di_crc in xfs_log_dinode to prevent KMSAN warnings
   - Use static_assert() instead of BUILD_BUG_ON_MSG() to validate size
     of structures and structure member offsets. This is done in order
     to be able to share the code with userspace
   - Move XFS documentation under a new directory specific to XFS
   - Do not invoke deferred ops' ->create_done callback if the deferred
     operation does not have an intent item associated with it
   - Remove duplicate inclusion of header files from scrub/health.c
   - Refactor Realtime code
   - Cleanup attr code"

* tag 'xfs-6.8-merge-3' of git://git.kernel.org/pub/scm/fs/xfs/xfs-linux: (123 commits)
  xfs: use the op name in trace_xlog_intent_recovery_failed
  xfs: fix a use after free in xfs_defer_finish_recovery
  xfs: turn the XFS_DA_OP_REPLACE checks in xfs_attr_shortform_addname into asserts
  xfs: remove xfs_attr_sf_hdr_t
  xfs: remove struct xfs_attr_shortform
  xfs: use xfs_attr_sf_findname in xfs_attr_shortform_getvalue
  xfs: remove xfs_attr_shortform_lookup
  xfs: simplify xfs_attr_sf_findname
  xfs: move the xfs_attr_sf_lookup tracepoint
  xfs: return if_data from xfs_idata_realloc
  xfs: make if_data a void pointer
  xfs: fold xfs_rtallocate_extent into xfs_bmap_rtalloc
  xfs: simplify and optimize the RT allocation fallback cascade
  xfs: reorder the minlen and prod calculations in xfs_bmap_rtalloc
  xfs: remove XFS_RTMIN/XFS_RTMAX
  xfs: remove rt-wrappers from xfs_format.h
  xfs: factor out a xfs_rtalloc_sumlevel helper
  xfs: tidy up xfs_rtallocate_extent_exact
  xfs: merge the calls to xfs_rtallocate_range in xfs_rtallocate_block
  xfs: reflow the tail end of xfs_rtallocate_extent_block
  ...

5 files changed, 6963 insertions, 0 deletions

diff --git a/Documentation/filesystems/xfs/index.rst b/Documentation/filesystems/xfs/index.rst
new file mode 100644
index 000000000000..ab66c57a5d18
--- /dev/null
+++ b/Documentation/filesystems/xfs/index.rst
@@ -0,0 +1,14 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+============================
+XFS Filesystem Documentation
+============================
+
+.. toctree::
+   :maxdepth: 2
+   :numbered:
+
+   xfs-delayed-logging-design
+   xfs-maintainer-entry-profile
+   xfs-self-describing-metadata
+   xfs-online-fsck-design
diff --git a/Documentation/filesystems/xfs/xfs-delayed-logging-design.rst b/Documentation/filesystems/xfs/xfs-delayed-logging-design.rst
new file mode 100644
index 000000000000..6402ab8e370c
--- /dev/null
+++ b/Documentation/filesystems/xfs/xfs-delayed-logging-design.rst
@@ -0,0 +1,1087 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+==================
+XFS Logging Design
+==================
+
+Preamble
+========
+
+This document describes the design and algorithms that the XFS journalling
+subsystem is based on. This document describes the design and algorithms that
+the XFS journalling subsystem is based on so that readers may familiarize
+themselves with the general concepts of how transaction processing in XFS works.
+
+We begin with an overview of transactions in XFS, followed by describing how
+transaction reservations are structured and accounted, and then move into how we
+guarantee forwards progress for long running transactions with finite initial
+reservations bounds. At this point we need to explain how relogging works. With
+the basic concepts covered, the design of the delayed logging mechanism is
+documented.
+
+
+Introduction
+============
+
+XFS uses Write Ahead Logging for ensuring changes to the filesystem metadata
+are atomic and recoverable. For reasons of space and time efficiency, the
+logging mechanisms are varied and complex, combining intents, logical and
+physical logging mechanisms to provide the necessary recovery guarantees the
+filesystem requires.
+
+Some objects, such as inodes and dquots, are logged in logical format where the
+details logged are made up of the changes to in-core structures rather than
+on-disk structures. Other objects - typically buffers - have their physical
+changes logged. Long running atomic modifications have individual changes
+chained together by intents, ensuring that journal recovery can restart and
+finish an operation that was only partially done when the system stopped
+functioning.
+
+The reason for these differences is to keep the amount of log space and CPU time
+required to process objects being modified as small as possible and hence the
+logging overhead as low as possible. Some items are very frequently modified,
+and some parts of objects are more frequently modified than others, so keeping
+the overhead of metadata logging low is of prime importance.
+
+The method used to log an item or chain modifications together isn't
+particularly important in the scope of this document. It suffices to know that
+the method used for logging a particular object or chaining modifications
+together are different and are dependent on the object and/or modification being
+performed. The logging subsystem only cares that certain specific rules are
+followed to guarantee forwards progress and prevent deadlocks.
+
+
+Transactions in XFS
+===================
+
+XFS has two types of high level transactions, defined by the type of log space
+reservation they take. These are known as "one shot" and "permanent"
+transactions. Permanent transaction reservations can take reservations that span
+commit boundaries, whilst "one shot" transactions are for a single atomic
+modification.
+
+The type and size of reservation must be matched to the modification taking
+place.  This means that permanent transactions can be used for one-shot
+modifications, but one-shot reservations cannot be used for permanent
+transactions.
+
+In the code, a one-shot transaction pattern looks somewhat like this::
+
+	tp = xfs_trans_alloc(<reservation>)
+	<lock items>
+	<join item to transaction>
+	<do modification>
+	xfs_trans_commit(tp);
+
+As items are modified in the transaction, the dirty regions in those items are
+tracked via the transaction handle.  Once the transaction is committed, all
+resources joined to it are released, along with the remaining unused reservation
+space that was taken at the transaction allocation time.
+
+In contrast, a permanent transaction is made up of multiple linked individual
+transactions, and the pattern looks like this::
+
+	tp = xfs_trans_alloc(<reservation>)
+	xfs_ilock(ip, XFS_ILOCK_EXCL)
+
+	loop {
+		xfs_trans_ijoin(tp, 0);
+		<do modification>
+		xfs_trans_log_inode(tp, ip);
+		xfs_trans_roll(&tp);
+	}
+
+	xfs_trans_commit(tp);
+	xfs_iunlock(ip, XFS_ILOCK_EXCL);
+
+While this might look similar to a one-shot transaction, there is an important
+difference: xfs_trans_roll() performs a specific operation that links two
+transactions together::
+
+	ntp = xfs_trans_dup(tp);
+	xfs_trans_commit(tp);
+	xfs_trans_reserve(ntp);
+
+This results in a series of "rolling transactions" where the inode is locked
+across the entire chain of transactions.  Hence while this series of rolling
+transactions is running, nothing else can read from or write to the inode and
+this provides a mechanism for complex changes to appear atomic from an external
+observer's point of view.
+
+It is important to note that a series of rolling transactions in a permanent
+transaction does not form an atomic change in the journal. While each
+individual modification is atomic, the chain is *not atomic*. If we crash half
+way through, then recovery will only replay up to the last transactional
+modification the loop made that was committed to the journal.
+
+This affects long running permanent transactions in that it is not possible to
+predict how much of a long running operation will actually be recovered because
+there is no guarantee of how much of the operation reached stale storage. Hence
+if a long running operation requires multiple transactions to fully complete,
+the high level operation must use intents and deferred operations to guarantee
+recovery can complete the operation once the first transactions is persisted in
+the on-disk journal.
+
+
+Transactions are Asynchronous
+=============================
+
+In XFS, all high level transactions are asynchronous by default. This means that
+xfs_trans_commit() does not guarantee that the modification has been committed
+to stable storage when it returns. Hence when a system crashes, not all the
+completed transactions will be replayed during recovery.
+
+However, the logging subsystem does provide global ordering guarantees, such
+that if a specific change is seen after recovery, all metadata modifications
+that were committed prior to that change will also be seen.
+
+For single shot operations that need to reach stable storage immediately, or
+ensuring that a long running permanent transaction is fully committed once it is
+complete, we can explicitly tag a transaction as synchronous. This will trigger
+a "log force" to flush the outstanding committed transactions to stable storage
+in the journal and wait for that to complete.
+
+Synchronous transactions are rarely used, however, because they limit logging
+throughput to the IO latency limitations of the underlying storage. Instead, we
+tend to use log forces to ensure modifications are on stable storage only when
+a user operation requires a synchronisation point to occur (e.g. fsync).
+
+
+Transaction Reservations
+========================
+
+It has been mentioned a number of times now that the logging subsystem needs to
+provide a forwards progress guarantee so that no modification ever stalls
+because it can't be written to the journal due to a lack of space in the
+journal. This is achieved by the transaction reservations that are made when
+a transaction is first allocated. For permanent transactions, these reservations
+are maintained as part of the transaction rolling mechanism.
+
+A transaction reservation provides a guarantee that there is physical log space
+available to write the modification into the journal before we start making
+modifications to objects and items. As such, the reservation needs to be large
+enough to take into account the amount of metadata that the change might need to
+log in the worst case. This means that if we are modifying a btree in the
+transaction, we have to reserve enough space to record a full leaf-to-root split
+of the btree. As such, the reservations are quite complex because we have to
+take into account all the hidden changes that might occur.
+
+For example, a user data extent allocation involves allocating an extent from
+free space, which modifies the free space trees. That's two btrees.  Inserting
+the extent into the inode's extent map might require a split of the extent map
+btree, which requires another allocation that can modify the free space trees
+again.  Then we might have to update reverse mappings, which modifies yet
+another btree which might require more space. And so on.  Hence the amount of
+metadata that a "simple" operation can modify can be quite large.
+
+This "worst case" calculation provides us with the static "unit reservation"
+for the transaction that is calculated at mount time. We must guarantee that the
+log has this much space available before the transaction is allowed to proceed
+so that when we come to write the dirty metadata into the log we don't run out
+of log space half way through the write.
+
+For one-shot transactions, a single unit space reservation is all that is
+required for the transaction to proceed. For permanent transactions, however, we
+also have a "log count" that affects the size of the reservation that is to be
+made.
+
+While a permanent transaction can get by with a single unit of space
+reservation, it is somewhat inefficient to do this as it requires the
+transaction rolling mechanism to re-reserve space on every transaction roll. We
+know from the implementation of the permanent transactions how many transaction
+rolls are likely for the common modifications that need to be made.
+
+For example, an inode allocation is typically two transactions - one to
+physically allocate a free inode chunk on disk, and another to allocate an inode
+from an inode chunk that has free inodes in it.  Hence for an inode allocation
+transaction, we might set the reservation log count to a value of 2 to indicate
+that the common/fast path transaction will commit two linked transactions in a
+chain. Each time a permanent transaction rolls, it consumes an entire unit
+reservation.
+
+Hence when the permanent transaction is first allocated, the log space
+reservation is increased from a single unit reservation to multiple unit
+reservations. That multiple is defined by the reservation log count, and this
+means we can roll the transaction multiple times before we have to re-reserve
+log space when we roll the transaction. This ensures that the common
+modifications we make only need to reserve log space once.
+
+If the log count for a permanent transaction reaches zero, then it needs to
+re-reserve physical space in the log. This is somewhat complex, and requires
+an understanding of how the log accounts for space that has been reserved.
+
+
+Log Space Accounting
+====================
+
+The position in the log is typically referred to as a Log Sequence Number (LSN).
+The log is circular, so the positions in the log are defined by the combination
+of a cycle number - the number of times the log has been overwritten - and the
+offset into the log.  A LSN carries the cycle in the upper 32 bits and the
+offset in the lower 32 bits. The offset is in units of "basic blocks" (512
+bytes). Hence we can do realtively simple LSN based math to keep track of
+available space in the log.
+
+Log space accounting is done via a pair of constructs called "grant heads".  The
+position of the grant heads is an absolute value, so the amount of space
+available in the log is defined by the distance between the position of the
+grant head and the current log tail. That is, how much space can be
+reserved/consumed before the grant heads would fully wrap the log and overtake
+the tail position.
+
+The first grant head is the "reserve" head. This tracks the byte count of the
+reservations currently held by active transactions. It is a purely in-memory
+accounting of the space reservation and, as such, actually tracks byte offsets
+into the log rather than basic blocks. Hence it technically isn't using LSNs to
+represent the log position, but it is still treated like a split {cycle,offset}
+tuple for the purposes of tracking reservation space.
+
+The reserve grant head is used to accurately account for exact transaction
+reservations amounts and the exact byte count that modifications actually make
+and need to write into the log. The reserve head is used to prevent new
+transactions from taking new reservations when the head reaches the current
+tail. It will block new reservations in a FIFO queue and as the log tail moves
+forward it will wake them in order once sufficient space is available. This FIFO
+mechanism ensures no transaction is starved of resources when log space
+shortages occur.
+
+The other grant head is the "write" head. Unlike the reserve head, this grant
+head contains an LSN and it tracks the physical space usage in the log. While
+this might sound like it is accounting the same state as the reserve grant head
+- and it mostly does track exactly the same location as the reserve grant head -
+there are critical differences in behaviour between them that provides the
+forwards progress guarantees that rolling permanent transactions require.
+
+These differences when a permanent transaction is rolled and the internal "log
+count" reaches zero and the initial set of unit reservations have been
+exhausted. At this point, we still require a log space reservation to continue
+the next transaction in the sequeunce, but we have none remaining. We cannot
+sleep during the transaction commit process waiting for new log space to become
+available, as we may end up on the end of the FIFO queue and the items we have
+locked while we sleep could end up pinning the tail of the log before there is
+enough free space in the log to fulfill all of the pending reservations and
+then wake up transaction commit in progress.
+
+To take a new reservation without sleeping requires us to be able to take a
+reservation even if there is no reservation space currently available. That is,
+we need to be able to *overcommit* the log reservation space. As has already
+been detailed, we cannot overcommit physical log space. However, the reserve
+grant head does not track physical space - it only accounts for the amount of
+reservations we currently have outstanding. Hence if the reserve head passes
+over the tail of the log all it means is that new reservations will be throttled
+immediately and remain throttled until the log tail is moved forward far enough
+to remove the overcommit and start taking new reservations. In other words, we
+can overcommit the reserve head without violating the physical log head and tail
+rules.
+
+As a result, permanent transactions only "regrant" reservation space during
+xfs_trans_commit() calls, while the physical log space reservation - tracked by
+the write head - is then reserved separately by a call to xfs_log_reserve()
+after the commit completes. Once the commit completes, we can sleep waiting for
+physical log space to be reserved from the write grant head, but only if one
+critical rule has been observed::
+
+	Code using permanent reservations must always log the items they hold
+	locked across each transaction they roll in the chain.
+
+"Re-logging" the locked items on every transaction roll ensures that the items
+attached to the transaction chain being rolled are always relocated to the
+physical head of the log and so do not pin the tail of the log. If a locked item
+pins the tail of the log when we sleep on the write reservation, then we will
+deadlock the log as we cannot take the locks needed to write back that item and
+move the tail of the log forwards to free up write grant space. Re-logging the
+locked items avoids this deadlock and guarantees that the log reservation we are
+making cannot self-deadlock.
+
+If all rolling transactions obey this rule, then they can all make forwards
+progress independently because nothing will block the progress of the log
+tail moving forwards and hence ensuring that write grant space is always
+(eventually) made available to permanent transactions no matter how many times
+they roll.
+
+
+Re-logging Explained
+====================
+
+XFS allows multiple separate modifications to a single object to be carried in
+the log at any given time.  This allows the log to avoid needing to flush each
+change to disk before recording a new change to the object. XFS does this via a
+method called "re-logging". Conceptually, this is quite simple - all it requires
+is that any new change to the object is recorded with a *new copy* of all the
+existing changes in the new transaction that is written to the log.
+
+That is, if we have a sequence of changes A through to F, and the object was
+written to disk after change D, we would see in the log the following series
+of transactions, their contents and the log sequence number (LSN) of the
+transaction::
+
+	Transaction		Contents	LSN
+	   A			   A		   X
+	   B			  A+B		  X+n
+	   C			 A+B+C		 X+n+m
+	   D			A+B+C+D		X+n+m+o
+	    <object written to disk>
+	   E			   E		   Y (> X+n+m+o)
+	   F			  E+F		  Y+p
+
+In other words, each time an object is relogged, the new transaction contains
+the aggregation of all the previous changes currently held only in the log.
+
+This relogging technique allows objects to be moved forward in the log so that
+an object being relogged does not prevent the tail of the log from ever moving
+forward.  This can be seen in the table above by the changing (increasing) LSN
+of each subsequent transaction, and it's the technique that allows us to
+implement long-running, multiple-commit permanent transactions. 
+
+A typical example of a rolling transaction is the removal of extents from an
+inode which can only be done at a rate of two extents per transaction because
+of reservation size limitations. Hence a rolling extent removal transaction
+keeps relogging the inode and btree buffers as they get modified in each
+removal operation. This keeps them moving forward in the log as the operation
+progresses, ensuring that current operation never gets blocked by itself if the
+log wraps around.
+
+Hence it can be seen that the relogging operation is fundamental to the correct
+working of the XFS journalling subsystem. From the above description, most
+people should be able to see why the XFS metadata operations writes so much to
+the log - repeated operations to the same objects write the same changes to
+the log over and over again. Worse is the fact that objects tend to get
+dirtier as they get relogged, so each subsequent transaction is writing more
+metadata into the log.
+
+It should now also be obvious how relogging and asynchronous transactions go
+hand in hand. That is, transactions don't get written to the physical journal
+until either a log buffer is filled (a log buffer can hold multiple
+transactions) or a synchronous operation forces the log buffers holding the
+transactions to disk. This means that XFS is doing aggregation of transactions
+in memory - batching them, if you like - to minimise the impact of the log IO on
+transaction throughput.
+
+The limitation on asynchronous transaction throughput is the number and size of
+log buffers made available by the log manager. By default there are 8 log
+buffers available and the size of each is 32kB - the size can be increased up
+to 256kB by use of a mount option.
+
+Effectively, this gives us the maximum bound of outstanding metadata changes
+that can be made to the filesystem at any point in time - if all the log
+buffers are full and under IO, then no more transactions can be committed until
+the current batch completes. It is now common for a single current CPU core to
+be to able to issue enough transactions to keep the log buffers full and under
+IO permanently. Hence the XFS journalling subsystem can be considered to be IO
+bound.
+
+Delayed Logging: Concepts
+=========================
+
+The key thing to note about the asynchronous logging combined with the
+relogging technique XFS uses is that we can be relogging changed objects
+multiple times before they are committed to disk in the log buffers. If we
+return to the previous relogging example, it is entirely possible that
+transactions A through D are committed to disk in the same log buffer.
+
+That is, a single log buffer may contain multiple copies of the same object,
+but only one of those copies needs to be there - the last one "D", as it
+contains all the changes from the previous changes. In other words, we have one
+necessary copy in the log buffer, and three stale copies that are simply
+wasting space. When we are doing repeated operations on the same set of
+objects, these "stale objects" can be over 90% of the space used in the log
+buffers. It is clear that reducing the number of stale objects written to the
+log would greatly reduce the amount of metadata we write to the log, and this
+is the fundamental goal of delayed logging.
+
+From a conceptual point of view, XFS is already doing relogging in memory (where
+memory == log buffer), only it is doing it extremely inefficiently. It is using
+logical to physical formatting to do the relogging because there is no
+infrastructure to keep track of logical changes in memory prior to physically
+formatting the changes in a transaction to the log buffer. Hence we cannot avoid
+accumulating stale objects in the log buffers.
+
+Delayed logging is the name we've given to keeping and tracking transactional
+changes to objects in memory outside the log buffer infrastructure. Because of
+the relogging concept fundamental to the XFS journalling subsystem, this is
+actually relatively easy to do - all the changes to logged items are already
+tracked in the current infrastructure. The big problem is how to accumulate
+them and get them to the log in a consistent, recoverable manner.
+Describing the problems and how they have been solved is the focus of this
+document.
+
+One of the key changes that delayed logging makes to the operation of the
+journalling subsystem is that it disassociates the amount of outstanding
+metadata changes from the size and number of log buffers available. In other
+words, instead of there only being a maximum of 2MB of transaction changes not
+written to the log at any point in time, there may be a much greater amount
+being accumulated in memory. Hence the potential for loss of metadata on a
+crash is much greater than for the existing logging mechanism.
+
+It should be noted that this does not change the guarantee that log recovery
+will result in a consistent filesystem. What it does mean is that as far as the
+recovered filesystem is concerned, there may be many thousands of transactions
+that simply did not occur as a result of the crash. This makes it even more
+important that applications that care about their data use fsync() where they
+need to ensure application level data integrity is maintained.
+
+It should be noted that delayed logging is not an innovative new concept that
+warrants rigorous proofs to determine whether it is correct or not. The method
+of accumulating changes in memory for some period before writing them to the
+log is used effectively in many filesystems including ext3 and ext4. Hence
+no time is spent in this document trying to convince the reader that the
+concept is sound. Instead it is simply considered a "solved problem" and as
+such implementing it in XFS is purely an exercise in software engineering.
+
+The fundamental requirements for delayed logging in XFS are simple:
+
+	1. Reduce the amount of metadata written to the log by at least
+	   an order of magnitude.
+	2. Supply sufficient statistics to validate Requirement #1.
+	3. Supply sufficient new tracing infrastructure to be able to debug
+	   problems with the new code.
+	4. No on-disk format change (metadata or log format).
+	5. Enable and disable with a mount option.
+	6. No performance regressions for synchronous transaction workloads.
+
+Delayed Logging: Design
+=======================
+
+Storing Changes
+---------------
+
+The problem with accumulating changes at a logical level (i.e. just using the
+existing log item dirty region tracking) is that when it comes to writing the
+changes to the log buffers, we need to ensure that the object we are formatting
+is not changing while we do this. This requires locking the object to prevent
+concurrent modification. Hence flushing the logical changes to the log would
+require us to lock every object, format them, and then unlock them again.
+
+This introduces lots of scope for deadlocks with transactions that are already
+running. For example, a transaction has object A locked and modified, but needs
+the delayed logging tracking lock to commit the transaction. However, the
+flushing thread has the delayed logging tracking lock already held, and is
+trying to get the lock on object A to flush it to the log buffer. This appears
+to be an unsolvable deadlock condition, and it was solving this problem that
+was the barrier to implementing delayed logging for so long.
+
+The solution is relatively simple - it just took a long time to recognise it.
+Put simply, the current logging code formats the changes to each item into an
+vector array that points to the changed regions in the item. The log write code
+simply copies the memory these vectors point to into the log buffer during
+transaction commit while the item is locked in the transaction. Instead of
+using the log buffer as the destination of the formatting code, we can use an
+allocated memory buffer big enough to fit the formatted vector.
+
+If we then copy the vector into the memory buffer and rewrite the vector to
+point to the memory buffer rather than the object itself, we now have a copy of
+the changes in a format that is compatible with the log buffer writing code.
+that does not require us to lock the item to access. This formatting and
+rewriting can all be done while the object is locked during transaction commit,
+resulting in a vector that is transactionally consistent and can be accessed
+without needing to lock the owning item.
+
+Hence we avoid the need to lock items when we need to flush outstanding
+asynchronous transactions to the log. The differences between the existing
+formatting method and the delayed logging formatting can be seen in the
+diagram below.
+
+Current format log vector::
+
+    Object    +---------------------------------------------+
+    Vector 1      +----+
+    Vector 2                    +----+
+    Vector 3                                   +----------+
+
+After formatting::
+
+    Log Buffer    +-V1-+-V2-+----V3----+
+
+Delayed logging vector::
+
+    Object    +---------------------------------------------+
+    Vector 1      +----+
+    Vector 2                    +----+
+    Vector 3                                   +----------+
+
+After formatting::
+
+    Memory Buffer +-V1-+-V2-+----V3----+
+    Vector 1      +----+
+    Vector 2           +----+
+    Vector 3                +----------+
+
+The memory buffer and associated vector need to be passed as a single object,
+but still need to be associated with the parent object so if the object is
+relogged we can replace the current memory buffer with a new memory buffer that
+contains the latest changes.
+
+The reason for keeping the vector around after we've formatted the memory
+buffer is to support splitting vectors across log buffer boundaries correctly.
+If we don't keep the vector around, we do not know where the region boundaries
+are in the item, so we'd need a new encapsulation method for regions in the log
+buffer writing (i.e. double encapsulation). This would be an on-disk format
+change and as such is not desirable.  It also means we'd have to write the log
+region headers in the formatting stage, which is problematic as there is per
+region state that needs to be placed into the headers during the log write.
+
+Hence we need to keep the vector, but by attaching the memory buffer to it and
+rewriting the vector addresses to point at the memory buffer we end up with a
+self-describing object that can be passed to the log buffer write code to be
+handled in exactly the same manner as the existing log vectors are handled.
+Hence we avoid needing a new on-disk format to handle items that have been
+relogged in memory.
+
+
+Tracking Changes
+----------------
+
+Now that we can record transactional changes in memory in a form that allows
+them to be used without limitations, we need to be able to track and accumulate
+them so that they can be written to the log at some later point in time.  The
+log item is the natural place to store this vector and buffer, and also makes sense
+to be the object that is used to track committed objects as it will always
+exist once the object has been included in a transaction.
+
+The log item is already used to track the log items that have been written to
+the log but not yet written to disk. Such log items are considered "active"
+and as such are stored in the Active Item List (AIL) which is a LSN-ordered
+double linked list. Items are inserted into this list during log buffer IO
+completion, after which they are unpinned and can be written to disk. An object
+that is in the AIL can be relogged, which causes the object to be pinned again
+and then moved forward in the AIL when the log buffer IO completes for that
+transaction.
+
+Essentially, this shows that an item that is in the AIL can still be modified
+and relogged, so any tracking must be separate to the AIL infrastructure. As
+such, we cannot reuse the AIL list pointers for tracking committed items, nor
+can we store state in any field that is protected by the AIL lock. Hence the
+committed item tracking needs its own locks, lists and state fields in the log
+item.
+
+Similar to the AIL, tracking of committed items is done through a new list
+called the Committed Item List (CIL).  The list tracks log items that have been
+committed and have formatted memory buffers attached to them. It tracks objects
+in transaction commit order, so when an object is relogged it is removed from
+its place in the list and re-inserted at the tail. This is entirely arbitrary
+and done to make it easy for debugging - the last items in the list are the
+ones that are most recently modified. Ordering of the CIL is not necessary for
+transactional integrity (as discussed in the next section) so the ordering is
+done for convenience/sanity of the developers.
+
+
+Delayed Logging: Checkpoints
+----------------------------
+
+When we have a log synchronisation event, commonly known as a "log force",
+all the items in the CIL must be written into the log via the log buffers.
+We need to write these items in the order that they exist in the CIL, and they
+need to be written as an atomic transaction. The need for all the objects to be
+written as an atomic transaction comes from the requirements of relogging and
+log replay - all the changes in all the objects in a given transaction must
+either be completely replayed during log recovery, or not replayed at all. If
+a transaction is not replayed because it is not complete in the log, then
+no later transactions should be replayed, either.
+
+To fulfill this requirement, we need to write the entire CIL in a single log
+transaction. Fortunately, the XFS log code has no fixed limit on the size of a
+transaction, nor does the log replay code. The only fundamental limit is that
+the transaction cannot be larger than just under half the size of the log.  The
+reason for this limit is that to find the head and tail of the log, there must
+be at least one complete transaction in the log at any given time. If a
+transaction is larger than half the log, then there is the possibility that a
+crash during the write of a such a transaction could partially overwrite the
+only complete previous transaction in the log. This will result in a recovery
+failure and an inconsistent filesystem and hence we must enforce the maximum
+size of a checkpoint to be slightly less than a half the log.
+
+Apart from this size requirement, a checkpoint transaction looks no different
+to any other transaction - it contains a transaction header, a series of
+formatted log items and a commit record at the tail. From a recovery
+perspective, the checkpoint transaction is also no different - just a lot
+bigger with a lot more items in it. The worst case effect of this is that we
+might need to tune the recovery transaction object hash size.
+
+Because the checkpoint is just another transaction and all the changes to log
+items are stored as log vectors, we can use the existing log buffer writing
+code to write the changes into the log. To do this efficiently, we need to
+minimise the time we hold the CIL locked while writing the checkpoint
+transaction. The current log write code enables us to do this easily with the
+way it separates the writing of the transaction contents (the log vectors) from
+the transaction commit record, but tracking this requires us to have a
+per-checkpoint context that travels through the log write process through to
+checkpoint completion.
+
+Hence a checkpoint has a context that tracks the state of the current
+checkpoint from initiation to checkpoint completion. A new context is initiated
+at the same time a checkpoint transaction is started. That is, when we remove
+all the current items from the CIL during a checkpoint operation, we move all
+those changes into the current checkpoint context. We then initialise a new
+context and attach that to the CIL for aggregation of new transactions.
+
+This allows us to unlock the CIL immediately after transfer of all the
+committed items and effectively allows new transactions to be issued while we
+are formatting the checkpoint into the log. It also allows concurrent
+checkpoints to be written into the log buffers in the case of log force heavy
+workloads, just like the existing transaction commit code does. This, however,
+requires that we strictly order the commit records in the log so that
+checkpoint sequence order is maintained during log replay.
+
+To ensure that we can be writing an item into a checkpoint transaction at
+the same time another transaction modifies the item and inserts the log item
+into the new CIL, then checkpoint transaction commit code cannot use log items
+to store the list of log vectors that need to be written into the transaction.
+Hence log vectors need to be able to be chained together to allow them to be
+detached from the log items. That is, when the CIL is flushed the memory
+buffer and log vector attached to each log item needs to be attached to the
+checkpoint context so that the log item can be released. In diagrammatic form,
+the CIL would look like this before the flush::
+
+	CIL Head
+	   |
+	   V
+	Log Item <-> log vector 1	-> memory buffer
+	   |				-> vector array
+	   V
+	Log Item <-> log vector 2	-> memory buffer
+	   |				-> vector array
+	   V
+	......
+	   |
+	   V
+	Log Item <-> log vector N-1	-> memory buffer
+	   |				-> vector array
+	   V
+	Log Item <-> log vector N	-> memory buffer
+					-> vector array
+
+And after the flush the CIL head is empty, and the checkpoint context log
+vector list would look like::
+
+	Checkpoint Context
+	   |
+	   V
+	log vector 1	-> memory buffer
+	   |		-> vector array
+	   |		-> Log Item
+	   V
+	log vector 2	-> memory buffer
+	   |		-> vector array
+	   |		-> Log Item
+	   V
+	......
+	   |
+	   V
+	log vector N-1	-> memory buffer
+	   |		-> vector array
+	   |		-> Log Item
+	   V
+	log vector N	-> memory buffer
+			-> vector array
+			-> Log Item
+
+Once this transfer is done, the CIL can be unlocked and new transactions can
+start, while the checkpoint flush code works over the log vector chain to
+commit the checkpoint.
+
+Once the checkpoint is written into the log buffers, the checkpoint context is
+attached to the log buffer that the commit record was written to along with a
+completion callback. Log IO completion will call that callback, which can then
+run transaction committed processing for the log items (i.e. insert into AIL
+and unpin) in the log vector chain and then free the log vector chain and
+checkpoint context.
+
+Discussion Point: I am uncertain as to whether the log item is the most
+efficient way to track vectors, even though it seems like the natural way to do
+it. The fact that we walk the log items (in the CIL) just to chain the log
+vectors and break the link between the log item and the log vector means that
+we take a cache line hit for the log item list modification, then another for
+the log vector chaining. If we track by the log vectors, then we only need to
+break the link between the log item and the log vector, which means we should
+dirty only the log item cachelines. Normally I wouldn't be concerned about one
+vs two dirty cachelines except for the fact I've seen upwards of 80,000 log
+vectors in one checkpoint transaction. I'd guess this is a "measure and
+compare" situation that can be done after a working and reviewed implementation
+is in the dev tree....
+
+Delayed Logging: Checkpoint Sequencing
+--------------------------------------
+
+One of the key aspects of the XFS transaction subsystem is that it tags
+committed transactions with the log sequence number of the transaction commit.
+This allows transactions to be issued asynchronously even though there may be
+future operations that cannot be completed until that transaction is fully
+committed to the log. In the rare case that a dependent operation occurs (e.g.
+re-using a freed metadata extent for a data extent), a special, optimised log
+force can be issued to force the dependent transaction to disk immediately.
+
+To do this, transactions need to record the LSN of the commit record of the
+transaction. This LSN comes directly from the log buffer the transaction is
+written into. While this works just fine for the existing transaction
+mechanism, it does not work for delayed logging because transactions are not
+written directly into the log buffers. Hence some other method of sequencing
+transactions is required.
+
+As discussed in the checkpoint section, delayed logging uses per-checkpoint
+contexts, and as such it is simple to assign a sequence number to each
+checkpoint. Because the switching of checkpoint contexts must be done
+atomically, it is simple to ensure that each new context has a monotonically
+increasing sequence number assigned to it without the need for an external
+atomic counter - we can just take the current context sequence number and add
+one to it for the new context.
+
+Then, instead of assigning a log buffer LSN to the transaction commit LSN
+during the commit, we can assign the current checkpoint sequence. This allows
+operations that track transactions that have not yet completed know what
+checkpoint sequence needs to be committed before they can continue. As a
+result, the code that forces the log to a specific LSN now needs to ensure that
+the log forces to a specific checkpoint.
+
+To ensure that we can do this, we need to track all the checkpoint contexts
+that are currently committing to the log. When we flush a checkpoint, the
+context gets added to a "committing" list which can be searched. When a
+checkpoint commit completes, it is removed from the committing list. Because
+the checkpoint context records the LSN of the commit record for the checkpoint,
+we can also wait on the log buffer that contains the commit record, thereby
+using the existing log force mechanisms to execute synchronous forces.
+
+It should be noted that the synchronous forces may need to be extended with
+mitigation algorithms similar to the current log buffer code to allow
+aggregation of multiple synchronous transactions if there are already
+synchronous transactions being flushed. Investigation of the performance of the
+current design is needed before making any decisions here.
+
+The main concern with log forces is to ensure that all the previous checkpoints
+are also committed to disk before the one we need to wait for. Therefore we
+need to check that all the prior contexts in the committing list are also
+complete before waiting on the one we need to complete. We do this
+synchronisation in the log force code so that we don't need to wait anywhere
+else for such serialisation - it only matters when we do a log force.
+
+The only remaining complexity is that a log force now also has to handle the
+case where the forcing sequence number is the same as the current context. That
+is, we need to flush the CIL and potentially wait for it to complete. This is a
+simple addition to the existing log forcing code to check the sequence numbers
+and push if required. Indeed, placing the current sequence checkpoint flush in
+the log force code enables the current mechanism for issuing synchronous
+transactions to remain untouched (i.e. commit an asynchronous transaction, then
+force the log at the LSN of that transaction) and so the higher level code
+behaves the same regardless of whether delayed logging is being used or not.
+
+Delayed Logging: Checkpoint Log Space Accounting
+------------------------------------------------
+
+The big issue for a checkpoint transaction is the log space reservation for the
+transaction. We don't know how big a checkpoint transaction is going to be
+ahead of time, nor how many log buffers it will take to write out, nor the
+number of split log vector regions are going to be used. We can track the
+amount of log space required as we add items to the commit item list, but we
+still need to reserve the space in the log for the checkpoint.
+
+A typical transaction reserves enough space in the log for the worst case space
+usage of the transaction. The reservation accounts for log record headers,
+transaction and region headers, headers for split regions, buffer tail padding,
+etc. as well as the actual space for all the changed metadata in the
+transaction. While some of this is fixed overhead, much of it is dependent on
+the size of the transaction and the number of regions being logged (the number
+of log vectors in the transaction).
+
+An example of the differences would be logging directory changes versus logging
+inode changes. If you modify lots of inode cores (e.g. ``chmod -R g+w *``), then
+there are lots of transactions that only contain an inode core and an inode log
+format structure. That is, two vectors totaling roughly 150 bytes. If we modify
+10,000 inodes, we have about 1.5MB of metadata to write in 20,000 vectors. Each
+vector is 12 bytes, so the total to be logged is approximately 1.75MB. In
+comparison, if we are logging full directory buffers, they are typically 4KB
+each, so we in 1.5MB of directory buffers we'd have roughly 400 buffers and a
+buffer format structure for each buffer - roughly 800 vectors or 1.51MB total
+space.  From this, it should be obvious that a static log space reservation is
+not particularly flexible and is difficult to select the "optimal value" for
+all workloads.
+
+Further, if we are going to use a static reservation, which bit of the entire
+reservation does it cover? We account for space used by the transaction
+reservation by tracking the space currently used by the object in the CIL and
+then calculating the increase or decrease in space used as the object is
+relogged. This allows for a checkpoint reservation to only have to account for
+log buffer metadata used such as log header records.
+
+However, even using a static reservation for just the log metadata is
+problematic. Typically log record headers use at least 16KB of log space per
+1MB of log space consumed (512 bytes per 32k) and the reservation needs to be
+large enough to handle arbitrary sized checkpoint transactions. This
+reservation needs to be made before the checkpoint is started, and we need to
+be able to reserve the space without sleeping.  For a 8MB checkpoint, we need a
+reservation of around 150KB, which is a non-trivial amount of space.
+
+A static reservation needs to manipulate the log grant counters - we can take a
+permanent reservation on the space, but we still need to make sure we refresh
+the write reservation (the actual space available to the transaction) after
+every checkpoint transaction completion. Unfortunately, if this space is not
+available when required, then the regrant code will sleep waiting for it.
+
+The problem with this is that it can lead to deadlocks as we may need to commit
+checkpoints to be able to free up log space (refer back to the description of
+rolling transactions for an example of this).  Hence we *must* always have
+space available in the log if we are to use static reservations, and that is
+very difficult and complex to arrange. It is possible to do, but there is a
+simpler way.
+
+The simpler way of doing this is tracking the entire log space used by the
+items in the CIL and using this to dynamically calculate the amount of log
+space required by the log metadata. If this log metadata space changes as a
+result of a transaction commit inserting a new memory buffer into the CIL, then
+the difference in space required is removed from the transaction that causes
+the change. Transactions at this level will *always* have enough space
+available in their reservation for this as they have already reserved the
+maximal amount of log metadata space they require, and such a delta reservation
+will always be less than or equal to the maximal amount in the reservation.
+
+Hence we can grow the checkpoint transaction reservation dynamically as items
+are added to the CIL and avoid the need for reserving and regranting log space
+up front. This avoids deadlocks and removes a blocking point from the
+checkpoint flush code.
+
+As mentioned early, transactions can't grow to more than half the size of the
+log. Hence as part of the reservation growing, we need to also check the size
+of the reservation against the maximum allowed transaction size. If we reach
+the maximum threshold, we need to push the CIL to the log. This is effectively
+a "background flush" and is done on demand. This is identical to
+a CIL push triggered by a log force, only that there is no waiting for the
+checkpoint commit to complete. This background push is checked and executed by
+transaction commit code.
+
+If the transaction subsystem goes idle while we still have items in the CIL,
+they will be flushed by the periodic log force issued by the xfssyncd. This log
+force will push the CIL to disk, and if the transaction subsystem stays idle,
+allow the idle log to be covered (effectively marked clean) in exactly the same
+manner that is done for the existing logging method. A discussion point is
+whether this log force needs to be done more frequently than the current rate
+which is once every 30s.
+
+
+Delayed Logging: Log Item Pinning
+---------------------------------
+
+Currently log items are pinned during transaction commit while the items are
+still locked. This happens just after the items are formatted, though it could
+be done any time before the items are unlocked. The result of this mechanism is
+that items get pinned once for every transaction that is committed to the log
+buffers. Hence items that are relogged in the log buffers will have a pin count
+for every outstanding transaction they were dirtied in. When each of these
+transactions is completed, they will unpin the item once. As a result, the item
+only becomes unpinned when all the transactions complete and there are no
+pending transactions. Thus the pinning and unpinning of a log item is symmetric
+as there is a 1:1 relationship with transaction commit and log item completion.
+
+For delayed logging, however, we have an asymmetric transaction commit to
+completion relationship. Every time an object is relogged in the CIL it goes
+through the commit process without a corresponding completion being registered.
+That is, we now have a many-to-one relationship between transaction commit and
+log item completion. The result of this is that pinning and unpinning of the
+log items becomes unbalanced if we retain the "pin on transaction commit, unpin
+on transaction completion" model.
+
+To keep pin/unpin symmetry, the algorithm needs to change to a "pin on
+insertion into the CIL, unpin on checkpoint completion". In other words, the
+pinning and unpinning becomes symmetric around a checkpoint context. We have to
+pin the object the first time it is inserted into the CIL - if it is already in
+the CIL during a transaction commit, then we do not pin it again. Because there
+can be multiple outstanding checkpoint contexts, we can still see elevated pin
+counts, but as each checkpoint completes the pin count will retain the correct
+value according to its context.
+
+Just to make matters slightly more complex, this checkpoint level context
+for the pin count means that the pinning of an item must take place under the
+CIL commit/flush lock. If we pin the object outside this lock, we cannot
+guarantee which context the pin count is associated with. This is because of
+the fact pinning the item is dependent on whether the item is present in the
+current CIL or not. If we don't pin the CIL first before we check and pin the
+object, we have a race with CIL being flushed between the check and the pin
+(or not pinning, as the case may be). Hence we must hold the CIL flush/commit
+lock to guarantee that we pin the items correctly.
+
+Delayed Logging: Concurrent Scalability
+---------------------------------------
+
+A fundamental requirement for the CIL is that accesses through transaction
+commits must scale to many concurrent commits. The current transaction commit
+code does not break down even when there are transactions coming from 2048
+processors at once. The current transaction code does not go any faster than if
+there was only one CPU using it, but it does not slow down either.
+
+As a result, the delayed logging transaction commit code needs to be designed
+for concurrency from the ground up. It is obvious that there are serialisation
+points in the design - the three important ones are:
+
+	1. Locking out new transaction commits while flushing the CIL
+	2. Adding items to the CIL and updating item space accounting
+	3. Checkpoint commit ordering
+
+Looking at the transaction commit and CIL flushing interactions, it is clear
+that we have a many-to-one interaction here. That is, the only restriction on
+the number of concurrent transactions that can be trying to commit at once is
+the amount of space available in the log for their reservations. The practical
+limit here is in the order of several hundred concurrent transactions for a
+128MB log, which means that it is generally one per CPU in a machine.
+
+The amount of time a transaction commit needs to hold out a flush is a
+relatively long period of time - the pinning of log items needs to be done
+while we are holding out a CIL flush, so at the moment that means it is held
+across the formatting of the objects into memory buffers (i.e. while memcpy()s
+are in progress). Ultimately a two pass algorithm where the formatting is done
+separately to the pinning of objects could be used to reduce the hold time of
+the transaction commit side.
+
+Because of the number of potential transaction commit side holders, the lock
+really needs to be a sleeping lock - if the CIL flush takes the lock, we do not
+want every other CPU in the machine spinning on the CIL lock. Given that
+flushing the CIL could involve walking a list of tens of thousands of log
+items, it will get held for a significant time and so spin contention is a
+significant concern. Preventing lots of CPUs spinning doing nothing is the
+main reason for choosing a sleeping lock even though nothing in either the
+transaction commit or CIL flush side sleeps with the lock held.
+
+It should also be noted that CIL flushing is also a relatively rare operation
+compared to transaction commit for asynchronous transaction workloads - only
+time will tell if using a read-write semaphore for exclusion will limit
+transaction commit concurrency due to cache line bouncing of the lock on the
+read side.
+
+The second serialisation point is on the transaction commit side where items
+are inserted into the CIL. Because transactions can enter this code
+concurrently, the CIL needs to be protected separately from the above
+commit/flush exclusion. It also needs to be an exclusive lock but it is only
+held for a very short time and so a spin lock is appropriate here. It is
+possible that this lock will become a contention point, but given the short
+hold time once per transaction I think that contention is unlikely.
+
+The final serialisation point is the checkpoint commit record ordering code
+that is run as part of the checkpoint commit and log force sequencing. The code
+path that triggers a CIL flush (i.e. whatever triggers the log force) will enter
+an ordering loop after writing all the log vectors into the log buffers but
+before writing the commit record. This loop walks the list of committing
+checkpoints and needs to block waiting for checkpoints to complete their commit
+record write. As a result it needs a lock and a wait variable. Log force
+sequencing also requires the same lock, list walk, and blocking mechanism to
+ensure completion of checkpoints.
+
+These two sequencing operations can use the mechanism even though the
+events they are waiting for are different. The checkpoint commit record
+sequencing needs to wait until checkpoint contexts contain a commit LSN
+(obtained through completion of a commit record write) while log force
+sequencing needs to wait until previous checkpoint contexts are removed from
+the committing list (i.e. they've completed). A simple wait variable and
+broadcast wakeups (thundering herds) has been used to implement these two
+serialisation queues. They use the same lock as the CIL, too. If we see too
+much contention on the CIL lock, or too many context switches as a result of
+the broadcast wakeups these operations can be put under a new spinlock and
+given separate wait lists to reduce lock contention and the number of processes
+woken by the wrong event.
+
+
+Lifecycle Changes
+-----------------
+
+The existing log item life cycle is as follows::
+
+	1. Transaction allocate
+	2. Transaction reserve
+	3. Lock item
+	4. Join item to transaction
+		If not already attached,
+			Allocate log item
+			Attach log item to owner item
+		Attach log item to transaction
+	5. Modify item
+		Record modifications in log item
+	6. Transaction commit
+		Pin item in memory
+		Format item into log buffer
+		Write commit LSN into transaction
+		Unlock item
+		Attach transaction to log buffer
+
+	<log buffer IO dispatched>
+	<log buffer IO completes>
+
+	7. Transaction completion
+		Mark log item committed
+		Insert log item into AIL
+			Write commit LSN into log item
+		Unpin log item
+	8. AIL traversal
+		Lock item
+		Mark log item clean
+		Flush item to disk
+
+	<item IO completion>
+
+	9. Log item removed from AIL
+		Moves log tail
+		Item unlocked
+
+Essentially, steps 1-6 operate independently from step 7, which is also
+independent of steps 8-9. An item can be locked in steps 1-6 or steps 8-9
+at the same time step 7 is occurring, but only steps 1-6 or 8-9 can occur
+at the same time. If the log item is in the AIL or between steps 6 and 7
+and steps 1-6 are re-entered, then the item is relogged. Only when steps 8-9
+are entered and completed is the object considered clean.
+
+With delayed logging, there are new steps inserted into the life cycle::
+
+	1. Transaction allocate
+	2. Transaction reserve
+	3. Lock item
+	4. Join item to transaction
+		If not already attached,
+			Allocate log item
+			Attach log item to owner item
+		Attach log item to transaction
+	5. Modify item
+		Record modifications in log item
+	6. Transaction commit
+		Pin item in memory if not pinned in CIL
+		Format item into log vector + buffer
+		Attach log vector and buffer to log item
+		Insert log item into CIL
+		Write CIL context sequence into transaction
+		Unlock item
+
+	<next log force>
+
+	7. CIL push
+		lock CIL flush
+		Chain log vectors and buffers together
+		Remove items from CIL
+		unlock CIL flush
+		write log vectors into log
+		sequence commit records
+		attach checkpoint context to log buffer
+
+	<log buffer IO dispatched>
+	<log buffer IO completes>
+
+	8. Checkpoint completion
+		Mark log item committed
+		Insert item into AIL
+			Write commit LSN into log item
+		Unpin log item
+	9. AIL traversal
+		Lock item
+		Mark log item clean
+		Flush item to disk
+	<item IO completion>
+	10. Log item removed from AIL
+		Moves log tail
+		Item unlocked
+
+From this, it can be seen that the only life cycle differences between the two
+logging methods are in the middle of the life cycle - they still have the same
+beginning and end and execution constraints. The only differences are in the
+committing of the log items to the log itself and the completion processing.
+Hence delayed logging should not introduce any constraints on log item
+behaviour, allocation or freeing that don't already exist.
+
+As a result of this zero-impact "insertion" of delayed logging infrastructure
+and the design of the internal structures to avoid on disk format changes, we
+can basically switch between delayed logging and the existing mechanism with a
+mount option. Fundamentally, there is no reason why the log manager would not
+be able to swap methods automatically and transparently depending on load
+characteristics, but this should not be necessary if delayed logging works as
+designed.
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+XFS Maintainer Entry Profile
+============================
+
+Overview
+--------
+XFS is a well known high-performance filesystem in the Linux kernel.
+The aim of this project is to provide and maintain a robust and
+performant filesystem.
+
+Patches are generally merged to the for-next branch of the appropriate
+git repository.
+After a testing period, the for-next branch is merged to the master
+branch.
+
+Kernel code are merged to the xfs-linux tree[0].
+Userspace code are merged to the xfsprogs tree[1].
+Test cases are merged to the xfstests tree[2].
+Ondisk format documentation are merged to the xfs-documentation tree[3].
+
+All patchsets involving XFS *must* be cc'd in their entirety to the mailing
+list linux-xfs@vger.kernel.org.
+
+Roles
+-----
+There are eight key roles in the XFS project.
+A person can take on multiple roles, and a role can be filled by
+multiple people.
+Anyone taking on a role is advised to check in with themselves and
+others on a regular basis about burnout.
+
+- **Outside Contributor**: Anyone who sends a patch but is not involved
+  in the XFS project on a regular basis.
+  These folks are usually people who work on other filesystems or
+  elsewhere in the kernel community.
+
+- **Developer**: Someone who is familiar with the XFS codebase enough to
+  write new code, documentation, and tests.
+
+  Developers can often be found in the IRC channel mentioned by the ``C:``
+  entry in the kernel MAINTAINERS file.
+
+- **Senior Developer**: A developer who is very familiar with at least
+  some part of the XFS codebase and/or other subsystems in the kernel.
+  These people collectively decide the long term goals of the project
+  and nudge the community in that direction.
+  They should help prioritize development and review work for each release
+  cycle.
+
+  Senior developers tend to be more active participants in the IRC channel.
+
+- **Reviewer**: Someone (most likely also a developer) who reads code
+  submissions to decide:
+
+  0. Is the idea behind the contribution sound?
+  1. Does the idea fit the goals of the project?
+  2. Is the contribution designed correctly?
+  3. Is the contribution polished?
+  4. Can the contribution be tested effectively?
+
+  Reviewers should identify themselves with an ``R:`` entry in the kernel
+  and fstests MAINTAINERS files.
+
+- **Testing Lead**: This person is responsible for setting the test
+  coverage goals of the project, negotiating with developers to decide
+  on new tests for new features, and making sure that developers and
+  release managers execute on the testing.
+
+  The testing lead should identify themselves with an ``M:`` entry in
+  the XFS section of the fstests MAINTAINERS file.
+
+- **Bug Triager**: Someone who examines incoming bug reports in just
+  enough detail to identify the person to whom the report should be
+  forwarded.
+
+  The bug triagers should identify themselves with a ``B:`` entry in
+  the kernel MAINTAINERS file.
+
+- **Release Manager**: This person merges reviewed patchsets into an
+  integration branch, tests the result locally, pushes the branch to a
+  public git repository, and sends pull requests further upstream.
+  The release manager is not expected to work on new feature patchsets.
+  If a developer and a reviewer fail to reach a resolution on some point,
+  the release manager must have the ability to intervene to try to drive a
+  resolution.
+
+  The release manager should identify themselves with an ``M:`` entry in
+  the kernel MAINTAINERS file.
+
+- **Community Manager**: This person calls and moderates meetings of as many
+  XFS participants as they can get when mailing list discussions prove
+  insufficient for collective decisionmaking.
+  They may also serve as liaison between managers of the organizations
+  sponsoring work on any part of XFS.
+
+- **LTS Maintainer**: Someone who backports and tests bug fixes from
+  uptream to the LTS kernels.
+  There tend to be six separate LTS trees at any given time.
+
+  The maintainer for a given LTS release should identify themselves with an
+  ``M:`` entry in the MAINTAINERS file for that LTS tree.
+  Unmaintained LTS kernels should be marked with status ``S: Orphan`` in that
+  same file.
+
+Submission Checklist Addendum
+-----------------------------
+Please follow these additional rules when submitting to XFS:
+
+- Patches affecting only the filesystem itself should be based against
+  the latest -rc or the for-next branch.
+  These patches will be merged back to the for-next branch.
+
+- Authors of patches touching other subsystems need to coordinate with
+  the maintainers of XFS and the relevant subsystems to decide how to
+  proceed with a merge.
+
+- Any patchset changing XFS should be cc'd in its entirety to linux-xfs.
+  Do not send partial patchsets; that makes analysis of the broader
+  context of the changes unnecessarily difficult.
+
+- Anyone making kernel changes that have corresponding changes to the
+  userspace utilities should send the userspace changes as separate
+  patchsets immediately after the kernel patchsets.
+
+- Authors of bug fix patches are expected to use fstests[2] to perform
+  an A/B test of the patch to determine that there are no regressions.
+  When possible, a new regression test case should be written for
+  fstests.
+
+- Authors of new feature patchsets must ensure that fstests will have
+  appropriate functional and input corner-case test cases for the new
+  feature.
+
+- When implementing a new feature, it is strongly suggested that the
+  developers write a design document to answer the following questions:
+
+  * **What** problem is this trying to solve?
+
+  * **Who** will benefit from this solution, and **where** will they
+    access it?
+
+  * **How** will this new feature work?  This should touch on major data
+    structures and algorithms supporting the solution at a higher level
+    than code comments.
+
+  * **What** userspace interfaces are necessary to build off of the new
+    features?
+
+  * **How** will this work be tested to ensure that it solves the
+    problems laid out in the design document without causing new
+    problems?
+
+  The design document should be committed in the kernel documentation
+  directory.
+  It may be omitted if the feature is already well known to the
+  community.
+
+- Patchsets for the new tests should be submitted as separate patchsets
+  immediately after the kernel and userspace code patchsets.
+
+- Changes to the on-disk format of XFS must be described in the ondisk
+  format document[3] and submitted as a patchset after the fstests
+  patchsets.
+
+- Patchsets implementing bug fixes and further code cleanups should put
+  the bug fixes at the beginning of the series to ease backporting.
+
+Key Release Cycle Dates
+-----------------------
+Bug fixes may be sent at any time, though the release manager may decide to
+defer a patch when the next merge window is close.
+
+Code submissions targeting the next merge window should be sent between
+-rc1 and -rc6.
+This gives the community time to review the changes, to suggest other changes,
+and for the author to retest those changes.
+
+Code submissions also requiring changes to fs/iomap and targeting the
+next merge window should be sent between -rc1 and -rc4.
+This allows the broader kernel community adequate time to test the
+infrastructure changes.
+
+Review Cadence
+--------------
+In general, please wait at least one week before pinging for feedback.
+To find reviewers, either consult the MAINTAINERS file, or ask
+developers that have Reviewed-by tags for XFS changes to take a look and
+offer their opinion.
+
+References
+----------
+| [0] https://git.kernel.org/pub/scm/fs/xfs/xfs-linux.git/
+| [1] https://git.kernel.org/pub/scm/fs/xfs/xfsprogs-dev.git/
+| [2] https://git.kernel.org/pub/scm/fs/xfs/xfstests-dev.git/
+| [3] https://git.kernel.org/pub/scm/fs/xfs/xfs-documentation.git/
diff --git a/Documentation/filesystems/xfs/xfs-online-fsck-design.rst b/Documentation/filesystems/xfs/xfs-online-fsck-design.rst
new file mode 100644
index 000000000000..352516feef6f
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@@ -0,0 +1,5315 @@
+.. SPDX-License-Identifier: GPL-2.0
+.. _xfs_online_fsck_design:
+
+..
+        Mapping of heading styles within this document:
+        Heading 1 uses "====" above and below
+        Heading 2 uses "===="
+        Heading 3 uses "----"
+        Heading 4 uses "````"
+        Heading 5 uses "^^^^"
+        Heading 6 uses "~~~~"
+        Heading 7 uses "...."
+
+        Sections are manually numbered because apparently that's what everyone
+        does in the kernel.
+
+======================
+XFS Online Fsck Design
+======================
+
+This document captures the design of the online filesystem check feature for
+XFS.
+The purpose of this document is threefold:
+
+- To help kernel distributors understand exactly what the XFS online fsck
+  feature is, and issues about which they should be aware.
+
+- To help people reading the code to familiarize themselves with the relevant
+  concepts and design points before they start digging into the code.
+
+- To help developers maintaining the system by capturing the reasons
+  supporting higher level decision making.
+
+As the online fsck code is merged, the links in this document to topic branches
+will be replaced with links to code.
+
+This document is licensed under the terms of the GNU Public License, v2.
+The primary author is Darrick J. Wong.
+
+This design document is split into seven parts.
+Part 1 defines what fsck tools are and the motivations for writing a new one.
+Parts 2 and 3 present a high level overview of how online fsck process works
+and how it is tested to ensure correct functionality.
+Part 4 discusses the user interface and the intended usage modes of the new
+program.
+Parts 5 and 6 show off the high level components and how they fit together, and
+then present case studies of how each repair function actually works.
+Part 7 sums up what has been discussed so far and speculates about what else
+might be built atop online fsck.
+
+.. contents:: Table of Contents
+   :local:
+
+1. What is a Filesystem Check?
+==============================
+
+A Unix filesystem has four main responsibilities:
+
+- Provide a hierarchy of names through which application programs can associate
+  arbitrary blobs of data for any length of time,
+
+- Virtualize physical storage media across those names, and
+
+- Retrieve the named data blobs at any time.
+
+- Examine resource usage.
+
+Metadata directly supporting these functions (e.g. files, directories, space
+mappings) are sometimes called primary metadata.
+Secondary metadata (e.g. reverse mapping and directory parent pointers) support
+operations internal to the filesystem, such as internal consistency checking
+and reorganization.
+Summary metadata, as the name implies, condense information contained in
+primary metadata for performance reasons.
+
+The filesystem check (fsck) tool examines all the metadata in a filesystem
+to look for errors.
+In addition to looking for obvious metadata corruptions, fsck also
+cross-references different types of metadata records with each other to look
+for inconsistencies.
+People do not like losing data, so most fsck tools also contains some ability
+to correct any problems found.
+As a word of caution -- the primary goal of most Linux fsck tools is to restore
+the filesystem metadata to a consistent state, not to maximize the data
+recovered.
+That precedent will not be challenged here.
+
+Filesystems of the 20th century generally lacked any redundancy in the ondisk
+format, which means that fsck can only respond to errors by erasing files until
+errors are no longer detected.
+More recent filesystem designs contain enough redundancy in their metadata that
+it is now possible to regenerate data structures when non-catastrophic errors
+occur; this capability aids both strategies.
+
++--------------------------------------------------------------------------+
+| **Note**:                                                                |
++--------------------------------------------------------------------------+
+| System administrators avoid data loss by increasing the number of        |
+| separate storage systems through the creation of backups; and they avoid |
+| downtime by increasing the redundancy of each storage system through the |
+| creation of RAID arrays.                                                 |
+| fsck tools address only the first problem.                               |
++--------------------------------------------------------------------------+
+
+TLDR; Show Me the Code!
+-----------------------
+
+Code is posted to the kernel.org git trees as follows:
+`kernel changes <https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-symlink>`_,
+`userspace changes <https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfsprogs-dev.git/log/?h=scrub-media-scan-service>`_, and
+`QA test changes <https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfstests-dev.git/log/?h=repair-dirs>`_.
+Each kernel patchset adding an online repair function will use the same branch
+name across the kernel, xfsprogs, and fstests git repos.
+
+Existing Tools
+--------------
+
+The online fsck tool described here will be the third tool in the history of
+XFS (on Linux) to check and repair filesystems.
+Two programs precede it:
+
+The first program, ``xfs_check``, was created as part of the XFS debugger
+(``xfs_db``) and can only be used with unmounted filesystems.
+It walks all metadata in the filesystem looking for inconsistencies in the
+metadata, though it lacks any ability to repair what it finds.
+Due to its high memory requirements and inability to repair things, this
+program is now deprecated and will not be discussed further.
+
+The second program, ``xfs_repair``, was created to be faster and more robust
+than the first program.
+Like its predecessor, it can only be used with unmounted filesystems.
+It uses extent-based in-memory data structures to reduce memory consumption,
+and tries to schedule readahead IO appropriately to reduce I/O waiting time
+while it scans the metadata of the entire filesystem.
+The most important feature of this tool is its ability to respond to
+inconsistencies in file metadata and directory tree by erasing things as needed
+to eliminate problems.
+Space usage metadata are rebuilt from the observed file metadata.
+
+Problem Statement
+-----------------
+
+The current XFS tools leave several problems unsolved:
+
+1. **User programs** suddenly **lose access** to the filesystem when unexpected
+   shutdowns occur as a result of silent corruptions in the metadata.
+   These occur **unpredictably** and often without warning.
+
+2. **Users** experience a **total loss of service** during the recovery period
+   after an **unexpected shutdown** occurs.
+
+3. **Users** experience a **total loss of service** if the filesystem is taken
+   offline to **look for problems** proactively.
+
+4. **Data owners** cannot **check the integrity** of their stored data without
+   reading all of it.
+   This may expose them to substantial billing costs when a linear media scan
+   performed by the storage system administrator might suffice.
+
+5. **System administrators** cannot **schedule** a maintenance window to deal
+   with corruptions if they **lack the means** to assess filesystem health
+   while the filesystem is online.
+
+6. **Fleet monitoring tools** cannot **automate periodic checks** of filesystem
+   health when doing so requires **manual intervention** and downtime.
+
+7. **Users** can be tricked into **doing things they do not desire** when
+   malicious actors **exploit quirks of Unicode** to place misleading names
+   in directories.
+
+Given this definition of the problems to be solved and the actors who would
+benefit, the proposed solution is a third fsck tool that acts on a running
+filesystem.
+
+This new third program has three components: an in-kernel facility to check
+metadata, an in-kernel facility to repair metadata, and a userspace driver
+program to drive fsck activity on a live filesystem.
+``xfs_scrub`` is the name of the driver program.
+The rest of this document presents the goals and use cases of the new fsck
+tool, describes its major design points in connection to those goals, and
+discusses the similarities and differences with existing tools.
+
++--------------------------------------------------------------------------+
+| **Note**:                                                                |
++--------------------------------------------------------------------------+
+| Throughout this document, the existing offline fsck tool can also be     |
+| referred to by its current name "``xfs_repair``".                        |
+| The userspace driver program for the new online fsck tool can be         |
+| referred to as "``xfs_scrub``".                                          |
+| The kernel portion of online fsck that validates metadata is called      |
+| "online scrub", and portion of the kernel that fixes metadata is called  |
+| "online repair".                                                         |
++--------------------------------------------------------------------------+
+
+The naming hierarchy is broken up into objects known as directories and files
+and the physical space is split into pieces known as allocation groups.
+Sharding enables better performance on highly parallel systems and helps to
+contain the damage when corruptions occur.
+The division of the filesystem into principal objects (allocation groups and
+inodes) means that there are ample opportunities to perform targeted checks and
+repairs on a subset of the filesystem.
+
+While this is going on, other parts continue processing IO requests.
+Even if a piece of filesystem metadata can only be regenerated by scanning the
+entire system, the scan can still be done in the background while other file
+operations continue.
+
+In summary, online fsck takes advantage of resource sharding and redundant
+metadata to enable targeted checking and repair operations while the system
+is running.
+This capability will be coupled to automatic system management so that
+autonomous self-healing of XFS maximizes service availability.
+
+2. Theory of Operation
+======================
+
+Because it is necessary for online fsck to lock and scan live metadata objects,
+online fsck consists of three separate code components.
+The first is the userspace driver program ``xfs_scrub``, which is responsible
+for identifying individual metadata items, scheduling work items for them,
+reacting to the outcomes appropriately, and reporting results to the system
+administrator.
+The second and third are in the kernel, which implements functions to check
+and repair each type of online fsck work item.
+
++------------------------------------------------------------------+
+| **Note**:                                                        |
++------------------------------------------------------------------+
+| For brevity, this document shortens the phrase "online fsck work |
+| item" to "scrub item".                                           |
++------------------------------------------------------------------+
+
+Scrub item types are delineated in a manner consistent with the Unix design
+philosophy, which is to say that each item should handle one aspect of a
+metadata structure, and handle it well.
+
+Scope
+-----
+
+In principle, online fsck should be able to check and to repair everything that
+the offline fsck program can handle.
+However, online fsck cannot be running 100% of the time, which means that
+latent errors may creep in after a scrub completes.
+If these errors cause the next mount to fail, offline fsck is the only
+solution.
+This limitation means that maintenance of the offline fsck tool will continue.
+A second limitation of online fsck is that it must follow the same resource
+sharing and lock acquisition rules as the regular filesystem.
+This means that scrub cannot take *any* shortcuts to save time, because doing
+so could lead to concurrency problems.
+In other words, online fsck is not a complete replacement for offline fsck, and
+a complete run of online fsck may take longer than online fsck.
+However, both of these limitations are acceptable tradeoffs to satisfy the
+different motivations of online fsck, which are to **minimize system downtime**
+and to **increase predictability of operation**.
+
+.. _scrubphases:
+
+Phases of Work
+--------------
+
+The userspace driver program ``xfs_scrub`` splits the work of checking and
+repairing an entire filesystem into seven phases.
+Each phase concentrates on checking specific types of scrub items and depends
+on the success of all previous phases.
+The seven phases are as follows:
+
+1. Collect geometry information about the mounted filesystem and computer,
+   discover the online fsck capabilities of the kernel, and open the
+   underlying storage devices.
+
+2. Check allocation group metadata, all realtime volume metadata, and all quota
+   files.
+   Each metadata structure is scheduled as a separate scrub item.
+   If corruption is found in the inode header or inode btree and ``xfs_scrub``
+   is permitted to perform repairs, then those scrub items are repaired to
+   prepare for phase 3.
+   Repairs are implemented by using the information in the scrub item to
+   resubmit the kernel scrub call with the repair flag enabled; this is
+   discussed in the next section.
+   Optimizations and all other repairs are deferred to phase 4.
+
+3. Check all metadata of every file in the filesystem.
+   Each metadata structure is also scheduled as a separate scrub item.
+   If repairs are needed and ``xfs_scrub`` is permitted to perform repairs,
+   and there were no problems detected during phase 2, then those scrub items
+   are repaired immediately.
+   Optimizations, deferred repairs, and unsuccessful repairs are deferred to
+   phase 4.
+
+4. All remaining repairs and scheduled optimizations are performed during this
+   phase, if the caller permits them.
+   Before starting repairs, the summary counters are checked and any necessary
+   repairs are performed so that subsequent repairs will not fail the resource
+   reservation step due to wildly incorrect summary counters.
+   Unsuccessful repairs are requeued as long as forward progress on repairs is
+   made somewhere in the filesystem.
+   Free space in the filesystem is trimmed at the end of phase 4 if the
+   filesystem is clean.
+
+5. By the start of this phase, all primary and secondary filesystem metadata
+   must be correct.
+   Summary counters such as the free space counts and quota resource counts
+   are checked and corrected.
+   Directory entry names and extended attribute names are checked for
+   suspicious entries such as control characters or confusing Unicode sequences
+   appearing in names.
+
+6. If the caller asks for a media scan, read all allocated and written data
+   file extents in the filesystem.
+   The ability to use hardware-assisted data file integrity checking is new
+   to online fsck; neither of the previous tools have this capability.
+   If media errors occur, they will be mapped to the owning files and reported.
+
+7. Re-check the summary counters and presents the caller with a summary of
+   space usage and file counts.
+
+This allocation of responsibilities will be :ref:`revisited <scrubcheck>`
+later in this document.
+
+Steps for Each Scrub Item
+-------------------------
+
+The kernel scrub code uses a three-step strategy for checking and repairing
+the one aspect of a metadata object represented by a scrub item:
+
+1. The scrub item of interest is checked for corruptions; opportunities for
+   optimization; and for values that are directly controlled by the system
+   administrator but look suspicious.
+   If the item is not corrupt or does not need optimization, resource are
+   released and the positive scan results are returned to userspace.
+   If the item is corrupt or could be optimized but the caller does not permit
+   this, resources are released and the negative scan results are returned to
+   userspace.
+   Otherwise, the kernel moves on to the second step.
+
+2. The repair function is called to rebuild the data structure.
+   Repair functions generally choose rebuild a structure from other metadata
+   rather than try to salvage the existing structure.
+   If the repair fails, the scan results from the first step are returned to
+   userspace.
+   Otherwise, the kernel moves on to the third step.
+
+3. In the third step, the kernel runs the same checks over the new metadata
+   item to assess the efficacy of the repairs.
+   The results of the reassessment are returned to userspace.
+
+Classification of Metadata
+--------------------------
+
+Each type of metadata object (and therefore each type of scrub item) is
+classified as follows:
+
+Primary Metadata
+````````````````
+
+Metadata structures in this category should be most familiar to filesystem
+users either because they are directly created by the user or they index
+objects created by the user
+Most filesystem objects fall into this class:
+
+- Free space and reference count information
+
+- Inode records and indexes
+
+- Storage mapping information for file data
+
+- Directories
+
+- Extended attributes
+
+- Symbolic links
+
+- Quota limits
+
+Scrub obeys the same rules as regular filesystem accesses for resource and lock
+acquisition.
+
+Primary metadata objects are the simplest for scrub to process.
+The principal filesystem object (either an allocation group or an inode) that
+owns the item being scrubbed is locked to guard against concurrent updates.
+The check function examines every record associated with the type for obvious
+errors and cross-references healthy records against other metadata to look for
+inconsistencies.
+Repairs for this class of scrub item are simple, since the repair function
+starts by holding all the resources acquired in the previous step.
+The repair function scans available metadata as needed to record all the
+observations needed to complete the structure.
+Next, it stages the observations in a new ondisk structure and commits it
+atomically to complete the repair.
+Finally, the storage from the old data structure are carefully reaped.
+
+Because ``xfs_scrub`` locks a primary object for the duration of the repair,
+this is effectively an offline repair operation performed on a subset of the
+filesystem.
+This minimizes the complexity of the repair code because it is not necessary to
+handle concurrent updates from other threads, nor is it necessary to access
+any other part of the filesystem.
+As a result, indexed structures can be rebuilt very quickly, and programs
+trying to access the damaged structure will be blocked until repairs complete.
+The only infrastructure needed by the repair code are the staging area for
+observations and a means to write new structures to disk.
+Despite these limitations, the advantage that online repair holds is clear:
+targeted work on individual shards of the filesystem avoids total loss of
+service.
+
+This mechanism is described in section 2.1 ("Off-Line Algorithm") of
+V. Srinivasan and M. J. Carey, `"Performance of On-Line Index Construction
+Algorithms" <https://minds.wisconsin.edu/bitstream/handle/1793/59524/TR1047.pdf>`_,
+*Extending Database Technology*, pp. 293-309, 1992.
+
+Most primary metadata repair functions stage their intermediate results in an
+in-memory array prior to formatting the new ondisk structure, which is very
+similar to the list-based algorithm discussed in section 2.3 ("List-Based
+Algorithms") of Srinivasan.
+However, any data structure builder that maintains a resource lock for the
+duration of the repair is *always* an offline algorithm.
+
+.. _secondary_metadata:
+
+Secondary Metadata
+``````````````````
+
+Metadata structures in this category reflect records found in primary metadata,
+but are only needed for online fsck or for reorganization of the filesystem.
+
+Secondary metadata include:
+
+- Reverse mapping information
+
+- Directory parent pointers
+
+This class of metadata is difficult for scrub to process because scrub attaches
+to the secondary object but needs to check primary metadata, which runs counter
+to the usual order of resource acquisition.
+Frequently, this means that full filesystems scans are necessary to rebuild the
+metadata.
+Check functions can be limited in scope to reduce runtime.
+Repairs, however, require a full scan of primary metadata, which can take a
+long time to complete.
+Under these conditions, ``xfs_scrub`` cannot lock resources for the entire
+duration of the repair.
+
+Instead, repair functions set up an in-memory staging structure to store
+observations.
+Depending on the requirements of the specific repair function, the staging
+index will either have the same format as the ondisk structure or a design
+specific to that repair function.
+The next step is to release all locks and start the filesystem scan.
+When the repair scanner needs to record an observation, the staging data are
+locked long enough to apply the update.
+While the filesystem scan is in progress, the repair function hooks the
+filesystem so that it can apply pending filesystem updates to the staging
+information.
+Once the scan is done, the owning object is re-locked, the live data is used to
+write a new ondisk structure, and the repairs are committed atomically.
+The hooks are disabled and the staging staging area is freed.
+Finally, the storage from the old data structure are carefully reaped.
+
+Introducing concurrency helps online repair avoid various locking problems, but
+comes at a high cost to code complexity.
+Live filesystem code has to be hooked so that the repair function can observe
+updates in progress.
+The staging area has to become a fully functional parallel structure so that
+updates can be merged from the hooks.
+Finally, the hook, the filesystem scan, and the inode locking model must be
+sufficiently well integrated that a hook event can decide if a given update
+should be applied to the staging structure.
+
+In theory, the scrub implementation could apply these same techniques for
+primary metadata, but doing so would make it massively more complex and less
+performant.
+Programs attempting to access the damaged structures are not blocked from
+operation, which may cause application failure or an unplanned filesystem
+shutdown.
+
+Inspiration for the secondary metadata repair strategy was drawn from section
+2.4 of Srinivasan above, and sections 2 ("NSF: Inded Build Without Side-File")
+and 3.1.1 ("Duplicate Key Insert Problem") in C. Mohan, `"Algorithms for
+Creating Indexes for Very Large Tables Without Quiescing Updates"
+<https://dl.acm.org/doi/10.1145/130283.130337>`_, 1992.
+
+The sidecar index mentioned above bears some resemblance to the side file
+method mentioned in Srinivasan and Mohan.
+Their method consists of an index builder that extracts relevant record data to
+build the new structure as quickly as possible; and an auxiliary structure that
+captures all updates that would be committed to the index by other threads were
+the new index already online.
+After the index building scan finishes, the updates recorded in the side file
+are applied to the new index.
+To avoid conflicts between the index builder and other writer threads, the
+builder maintains a publicly visible cursor that tracks the progress of the
+scan through the record space.
+To avoid duplication of work between the side file and the index builder, side
+file updates are elided when the record ID for the update is greater than the
+cursor position within the record ID space.
+
+To minimize changes to the rest of the codebase, XFS online repair keeps the
+replacement index hidden until it's completely ready to go.
+In other words, there is no attempt to expose the keyspace of the new index
+while repair is running.
+The complexity of such an approach would be very high and perhaps more
+appropriate to building *new* indices.
+
+**Future Work Question**: Can the full scan and live update code used to
+facilitate a repair also be used to implement a comprehensive check?
+
+*Answer*: In theory, yes.  Check would be much stronger if each scrub function
+employed these live scans to build a shadow copy of the metadata and then
+compared the shadow records to the ondisk records.
+However, doing that is a fair amount more work than what the checking functions
+do now.
+The live scans and hooks were developed much later.
+That in turn increases the runtime of those scrub functions.
+
+Summary Information
+```````````````````
+
+Metadata structures in this last category summarize the contents of primary
+metadata records.
+These are often used to speed up resource usage queries, and are many times
+smaller than the primary metadata which they represent.
+
+Examples of summary information include:
+
+- Summary counts of free space and inodes
+
+- File link counts from directories
+
+- Quota resource usage counts
+
+Check and repair require full filesystem scans, but resource and lock
+acquisition follow the same paths as regular filesystem accesses.
+
+The superblock summary counters have special requirements due to the underlying
+implementation of the incore counters, and will be treated separately.
+Check and repair of the other types of summary counters (quota resource counts
+and file link counts) employ the same filesystem scanning and hooking
+techniques as outlined above, but because the underlying data are sets of
+integer counters, the staging data need not be a fully functional mirror of the
+ondisk structure.
+
+Inspiration for quota and file link count repair strategies were drawn from
+sections 2.12 ("Online Index Operations") through 2.14 ("Incremental View
+Maintenance") of G.  Graefe, `"Concurrent Queries and Updates in Summary Views
+and Their Indexes"
+<http://www.odbms.org/wp-content/uploads/2014/06/Increment-locks.pdf>`_, 2011.
+
+Since quotas are non-negative integer counts of resource usage, online
+quotacheck can use the incremental view deltas described in section 2.14 to
+track pending changes to the block and inode usage counts in each transaction,
+and commit those changes to a dquot side file when the transaction commits.
+Delta tracking is necessary for dquots because the index builder scans inodes,
+whereas the data structure being rebuilt is an index of dquots.
+Link count checking combines the view deltas and commit step into one because
+it sets attributes of the objects being scanned instead of writing them to a
+separate data structure.
+Each online fsck function will be discussed as case studies later in this
+document.
+
+Risk Management
+---------------
+
+During the development of online fsck, several risk factors were identified
+that may make the feature unsuitable for certain distributors and users.
+Steps can be taken to mitigate or eliminate those risks, though at a cost to
+functionality.
+
+- **Decreased performance**: Adding metadata indices to the filesystem
+  increases the time cost of persisting changes to disk, and the reverse space
+  mapping and directory parent pointers are no exception.
+  System administrators who require the maximum performance can disable the
+  reverse mapping features at format time, though this choice dramatically
+  reduces the ability of online fsck to find inconsistencies and repair them.
+
+- **Incorrect repairs**: As with all software, there might be defects in the
+  software that result in incorrect repairs being written to the filesystem.
+  Systematic fuzz testing (detailed in the next section) is employed by the
+  authors to find bugs early, but it might not catch everything.
+  The kernel build system provides Kconfig options (``CONFIG_XFS_ONLINE_SCRUB``
+  and ``CONFIG_XFS_ONLINE_REPAIR``) to enable distributors to choose not to
+  accept this risk.
+  The xfsprogs build system has a configure option (``--enable-scrub=no``) that
+  disables building of the ``xfs_scrub`` binary, though this is not a risk
+  mitigation if the kernel functionality remains enabled.
+
+- **Inability to repair**: Sometimes, a filesystem is too badly damaged to be
+  repairable.
+  If the keyspaces of several metadata indices overlap in some manner but a
+  coherent narrative cannot be formed from records collected, then the repair
+  fails.
+  To reduce the chance that a repair will fail with a dirty transaction and
+  render the filesystem unusable, the online repair functions have been
+  designed to stage and validate all new records before committing the new
+  structure.
+
+- **Misbehavior**: Online fsck requires many privileges -- raw IO to block
+  devices, opening files by handle, ignoring Unix discretionary access control,
+  and the ability to perform administrative changes.
+  Running this automatically in the background scares people, so the systemd
+  background service is configured to run with only the privileges required.
+  Obviously, this cannot address certain problems like the kernel crashing or
+  deadlocking, but it should be sufficient to prevent the scrub process from
+  escaping and reconfiguring the system.
+  The cron job does not have this protection.
+
+- **Fuzz Kiddiez**: There are many people now who seem to think that running
+  automated fuzz testing of ondisk artifacts to find mischievous behavior and
+  spraying exploit code onto the public mailing list for instant zero-day
+  disclosure is somehow of some social benefit.
+  In the view of this author, the benefit is realized only when the fuzz
+  operators help to **fix** the flaws, but this opinion apparently is not
+  widely shared among security "researchers".
+  The XFS maintainers' continuing ability to manage these events presents an
+  ongoing risk to the stability of the development process.
+  Automated testing should front-load some of the risk while the feature is
+  considered EXPERIMENTAL.
+
+Many of these risks are inherent to software programming.
+Despite this, it is hoped that this new functionality will prove useful in
+reducing unexpected downtime.
+
+3. Testing Plan
+===============
+
+As stated before, fsck tools have three main goals:
+
+1. Detect inconsistencies in the metadata;
+
+2. Eliminate those inconsistencies; and
+
+3. Minimize further loss of data.
+
+Demonstrations of correct operation are necessary to build users' confidence
+that the software behaves within expectations.
+Unfortunately, it was not really feasible to perform regular exhaustive testing
+of every aspect of a fsck tool until the introduction of low-cost virtual
+machines with high-IOPS storage.
+With ample hardware availability in mind, the testing strategy for the online
+fsck project involves differential analysis against the existing fsck tools and
+systematic testing of every attribute of every type of metadata object.
+Testing can be split into four major categories, as discussed below.
+
+Integrated Testing with fstests
+-------------------------------
+
+The primary goal of any free software QA effort is to make testing as
+inexpensive and widespread as possible to maximize the scaling advantages of
+community.
+In other words, testing should maximize the breadth of filesystem configuration
+scenarios and hardware setups.
+This improves code quality by enabling the authors of online fsck to find and
+fix bugs early, and helps developers of new features to find integration
+issues earlier in their development effort.
+
+The Linux filesystem community shares a common QA testing suite,
+`fstests <https://git.kernel.org/pub/scm/fs/xfs/xfstests-dev.git/>`_, for
+functional and regression testing.
+Even before development work began on online fsck, fstests (when run on XFS)
+would run both the ``xfs_check`` and ``xfs_repair -n`` commands on the test and
+scratch filesystems between each test.
+This provides a level of assurance that the kernel and the fsck tools stay in
+alignment about what constitutes consistent metadata.
+During development of the online checking code, fstests was modified to run
+``xfs_scrub -n`` between each test to ensure that the new checking code
+produces the same results as the two existing fsck tools.
+
+To start development of online repair, fstests was modified to run
+``xfs_repair`` to rebuild the filesystem's metadata indices between tests.
+This ensures that offline repair does not crash, leave a corrupt filesystem
+after it exists, or trigger complaints from the online check.
+This also established a baseline for what can and cannot be repaired offline.
+To complete the first phase of development of online repair, fstests was
+modified to be able to run ``xfs_scrub`` in a "force rebuild" mode.
+This enables a comparison of the effectiveness of online repair as compared to
+the existing offline repair tools.
+
+General Fuzz Testing of Metadata Blocks
+---------------------------------------
+
+XFS benefits greatly from having a very robust debugging tool, ``xfs_db``.
+
+Before development of online fsck even began, a set of fstests were created
+to test the rather common fault that entire metadata blocks get corrupted.
+This required the creation of fstests library code that can create a filesystem
+containing every possible type of metadata object.
+Next, individual test cases were created to create a test filesystem, identify
+a single block of a specific type of metadata object, trash it with the
+existing ``blocktrash`` command in ``xfs_db``, and test the reaction of a
+particular metadata validation strategy.
+
+This earlier test suite enabled XFS developers to test the ability of the
+in-kernel validation functions and the ability of the offline fsck tool to
+detect and eliminate the inconsistent metadata.
+This part of the test suite was extended to cover online fsck in exactly the
+same manner.
+
+In other words, for a given fstests filesystem configuration:
+
+* For each metadata object existing on the filesystem:
+
+  * Write garbage to it
+
+  * Test the reactions of:
+
+    1. The kernel verifiers to stop obviously bad metadata
+    2. Offline repair (``xfs_repair``) to detect and fix
+    3. Online repair (``xfs_scrub``) to detect and fix
+
+Targeted Fuzz Testing of Metadata Records
+-----------------------------------------
+
+The testing plan for online fsck includes extending the existing fs testing
+infrastructure to provide a much more powerful facility: targeted fuzz testing
+of every metadata field of every metadata object in the filesystem.
+``xfs_db`` can modify every field of every metadata structure in every
+block in the filesystem to simulate the effects of memory corruption and
+software bugs.
+Given that fstests already contains the ability to create a filesystem
+containing every metadata format known to the filesystem, ``xfs_db`` can be
+used to perform exhaustive fuzz testing!
+
+For a given fstests filesystem configuration:
+
+* For each metadata object existing on the filesystem...
+
+  * For each record inside that metadata object...
+
+    * For each field inside that record...
+
+      * For each conceivable type of transformation that can be applied to a bit field...
+
+        1. Clear all bits
+        2. Set all bits
+        3. Toggle the most significant bit
+        4. Toggle the middle bit
+        5. Toggle the least significant bit
+        6. Add a small quantity
+        7. Subtract a small quantity
+        8. Randomize the contents
+
+        * ...test the reactions of:
+
+          1. The kernel verifiers to stop obviously bad metadata
+          2. Offline checking (``xfs_repair -n``)
+          3. Offline repair (``xfs_repair``)
+          4. Online checking (``xfs_scrub -n``)
+          5. Online repair (``xfs_scrub``)
+          6. Both repair tools (``xfs_scrub`` and then ``xfs_repair`` if online repair doesn't succeed)
+
+This is quite the combinatoric explosion!
+
+Fortunately, having this much test coverage makes it easy for XFS developers to
+check the responses of XFS' fsck tools.
+Since the introduction of the fuzz testing framework, these tests have been
+used to discover incorrect repair code and missing functionality for entire
+classes of metadata objects in ``xfs_repair``.
+The enhanced testing was used to finalize the deprecation of ``xfs_check`` by
+confirming that ``xfs_repair`` could detect at least as many corruptions as
+the older tool.
+
+These tests have been very valuable for ``xfs_scrub`` in the same ways -- they
+allow the online fsck developers to compare online fsck against offline fsck,
+and they enable XFS developers to find deficiencies in the code base.
+
+Proposed patchsets include
+`general fuzzer improvements
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfstests-dev.git/log/?h=fuzzer-improvements>`_,
+`fuzzing baselines
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfstests-dev.git/log/?h=fuzz-baseline>`_,
+and `improvements in fuzz testing comprehensiveness
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfstests-dev.git/log/?h=more-fuzz-testing>`_.
+
+Stress Testing
+--------------
+
+A unique requirement to online fsck is the ability to operate on a filesystem
+concurrently with regular workloads.
+Although it is of course impossible to run ``xfs_scrub`` with *zero* observable
+impact on the running system, the online repair code should never introduce
+inconsistencies into the filesystem metadata, and regular workloads should
+never notice resource starvation.
+To verify that these conditions are being met, fstests has been enhanced in
+the following ways:
+
+* For each scrub item type, create a test to exercise checking that item type
+  while running ``fsstress``.
+* For each scrub item type, create a test to exercise repairing that item type
+  while running ``fsstress``.
+* Race ``fsstress`` and ``xfs_scrub -n`` to ensure that checking the whole
+  filesystem doesn't cause problems.
+* Race ``fsstress`` and ``xfs_scrub`` in force-rebuild mode to ensure that
+  force-repairing the whole filesystem doesn't cause problems.
+* Race ``xfs_scrub`` in check and force-repair mode against ``fsstress`` while
+  freezing and thawing the filesystem.
+* Race ``xfs_scrub`` in check and force-repair mode against ``fsstress`` while
+  remounting the filesystem read-only and read-write.
+* The same, but running ``fsx`` instead of ``fsstress``.  (Not done yet?)
+
+Success is defined by the ability to run all of these tests without observing
+any unexpected filesystem shutdowns due to corrupted metadata, kernel hang
+check warnings, or any other sort of mischief.
+
+Proposed patchsets include `general stress testing
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfstests-dev.git/log/?h=race-scrub-and-mount-state-changes>`_
+and the `evolution of existing per-function stress testing
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfstests-dev.git/log/?h=refactor-scrub-stress>`_.
+
+4. User Interface
+=================
+
+The primary user of online fsck is the system administrator, just like offline
+repair.
+Online fsck presents two modes of operation to administrators:
+A foreground CLI process for online fsck on demand, and a background service
+that performs autonomous checking and repair.
+
+Checking on Demand
+------------------
+
+For administrators who want the absolute freshest information about the
+metadata in a filesystem, ``xfs_scrub`` can be run as a foreground process on
+a command line.
+The program checks every piece of metadata in the filesystem while the
+administrator waits for the results to be reported, just like the existing
+``xfs_repair`` tool.
+Both tools share a ``-n`` option to perform a read-only scan, and a ``-v``
+option to increase the verbosity of the information reported.
+
+A new feature of ``xfs_scrub`` is the ``-x`` option, which employs the error
+correction capabilities of the hardware to check data file contents.
+The media scan is not enabled by default because it may dramatically increase
+program runtime and consume a lot of bandwidth on older storage hardware.
+
+The output of a foreground invocation is captured in the system log.
+
+The ``xfs_scrub_all`` program walks the list of mounted filesystems and
+initiates ``xfs_scrub`` for each of them in parallel.
+It serializes scans for any filesystems that resolve to the same top level
+kernel block device to prevent resource overconsumption.
+
+Background Service
+------------------
+
+To reduce the workload of system administrators, the ``xfs_scrub`` package
+provides a suite of `systemd <https://systemd.io/>`_ timers and services that
+run online fsck automatically on weekends by default.
+The background service configures scrub to run with as little privilege as
+possible, the lowest CPU and IO priority, and in a CPU-constrained single
+threaded mode.
+This can be tuned by the systemd administrator at any time to suit the latency
+and throughput requirements of customer workloads.
+
+The output of the background service is also captured in the system log.
+If desired, reports of failures (either due to inconsistencies or mere runtime
+errors) can be emailed automatically by setting the ``EMAIL_ADDR`` environment
+variable in the following service files:
+
+* ``xfs_scrub_fail@.service``
+* ``xfs_scrub_media_fail@.service``
+* ``xfs_scrub_all_fail.service``
+
+The decision to enable the background scan is left to the system administrator.
+This can be done by enabling either of the following services:
+
+* ``xfs_scrub_all.timer`` on systemd systems
+* ``xfs_scrub_all.cron`` on non-systemd systems
+
+This automatic weekly scan is configured out of the box to perform an
+additional media scan of all file data once per month.
+This is less foolproof than, say, storing file data block checksums, but much
+more performant if application software provides its own integrity checking,
+redundancy can be provided elsewhere above the filesystem, or the storage
+device's integrity guarantees are deemed sufficient.
+
+The systemd unit file definitions have been subjected to a security audit
+(as of systemd 249) to ensure that the xfs_scrub processes have as little
+access to the rest of the system as possible.
+This was performed via ``systemd-analyze security``, after which privileges
+were restricted to the minimum required, sandboxing was set up to the maximal
+extent possible with sandboxing and system call filtering; and access to the
+filesystem tree was restricted to the minimum needed to start the program and
+access the filesystem being scanned.
+The service definition files restrict CPU usage to 80% of one CPU core, and
+apply as nice of a priority to IO and CPU scheduling as possible.
+This measure was taken to minimize delays in the rest of the filesystem.
+No such hardening has been performed for the cron job.
+
+Proposed patchset:
+`Enabling the xfs_scrub background service
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfsprogs-dev.git/log/?h=scrub-media-scan-service>`_.
+
+Health Reporting
+----------------
+
+XFS caches a summary of each filesystem's health status in memory.
+The information is updated whenever ``xfs_scrub`` is run, or whenever
+inconsistencies are detected in the filesystem metadata during regular
+operations.
+System administrators should use the ``health`` command of ``xfs_spaceman`` to
+download this information into a human-readable format.
+If problems have been observed, the administrator can schedule a reduced
+service window to run the online repair tool to correct the problem.
+Failing that, the administrator can decide to schedule a maintenance window to
+run the traditional offline repair tool to correct the problem.
+
+**Future Work Question**: Should the health reporting integrate with the new
+inotify fs error notification system?
+Would it be helpful for sysadmins to have a daemon to listen for corruption
+notifications and initiate a repair?
+
+*Answer*: These questions remain unanswered, but should be a part of the
+conversation with early adopters and potential downstream users of XFS.
+
+Proposed patchsets include
+`wiring up health reports to correction returns
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=corruption-health-reports>`_
+and
+`preservation of sickness info during memory reclaim
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=indirect-health-reporting>`_.
+
+5. Kernel Algorithms and Data Structures
+========================================
+
+This section discusses the key algorithms and data structures of the kernel
+code that provide the ability to check and repair metadata while the system
+is running.
+The first chapters in this section reveal the pieces that provide the
+foundation for checking metadata.
+The remainder of this section presents the mechanisms through which XFS
+regenerates itself.
+
+Self Describing Metadata
+------------------------
+
+Starting with XFS version 5 in 2012, XFS updated the format of nearly every
+ondisk block header to record a magic number, a checksum, a universally
+"unique" identifier (UUID), an owner code, the ondisk address of the block,
+and a log sequence number.
+When loading a block buffer from disk, the magic number, UUID, owner, and
+ondisk address confirm that the retrieved block matches the specific owner of
+the current filesystem, and that the information contained in the block is
+supposed to be found at the ondisk address.
+The first three components enable checking tools to disregard alleged metadata
+that doesn't belong to the filesystem, and the fourth component enables the
+filesystem to detect lost writes.
+
+Whenever a file system operation modifies a block, the change is submitted
+to the log as part of a transaction.
+The log then processes these transactions marking them done once they are
+safely persisted to storage.
+The logging code maintains the checksum and the log sequence number of the last
+transactional update.
+Checksums are useful for detecting torn writes and other discrepancies that can
+be introduced between the computer and its storage devices.
+Sequence number tracking enables log recovery to avoid applying out of date
+log updates to the filesystem.
+
+These two features improve overall runtime resiliency by providing a means for
+the filesystem to detect obvious corruption when reading metadata blocks from
+disk, but these buffer verifiers cannot provide any consistency checking
+between metadata structures.
+
+For more information, please see the documentation for
+Documentation/filesystems/xfs/xfs-self-describing-metadata.rst
+
+Reverse Mapping
+---------------
+
+The original design of XFS (circa 1993) is an improvement upon 1980s Unix
+filesystem design.
+In those days, storage density was expensive, CPU time was scarce, and
+excessive seek time could kill performance.
+For performance reasons, filesystem authors were reluctant to add redundancy to
+the filesystem, even at the cost of data integrity.
+Filesystems designers in the early 21st century choose different strategies to
+increase internal redundancy -- either storing nearly identical copies of
+metadata, or more space-efficient encoding techniques.
+
+For XFS, a different redundancy strategy was chosen to modernize the design:
+a secondary space usage index that maps allocated disk extents back to their
+owners.
+By adding a new index, the filesystem retains most of its ability to scale
+well to heavily threaded workloads involving large datasets, since the primary
+file metadata (the directory tree, the file block map, and the allocation
+groups) remain unchanged.
+Like any system that improves redundancy, the reverse-mapping feature increases
+overhead costs for space mapping activities.
+However, it has two critical advantages: first, the reverse index is key to
+enabling online fsck and other requested functionality such as free space
+defragmentation, better media failure reporting, and filesystem shrinking.
+Second, the different ondisk storage format of the reverse mapping btree
+defeats device-level deduplication because the filesystem requires real
+redundancy.
+
++--------------------------------------------------------------------------+
+| **Sidebar**:                                                             |
++--------------------------------------------------------------------------+
+| A criticism of adding the secondary index is that it does nothing to     |
+| improve the robustness of user data storage itself.                      |
+| This is a valid point, but adding a new index for file data block        |
+| checksums increases write amplification by turning data overwrites into  |
+| copy-writes, which age the filesystem prematurely.                       |
+| In keeping with thirty years of precedent, users who want file data      |
+| integrity can supply as powerful a solution as they require.             |
+| As for metadata, the complexity of adding a new secondary index of space |
+| usage is much less than adding volume management and storage device      |
+| mirroring to XFS itself.                                                 |
+| Perfection of RAID and volume management are best left to existing       |
+| layers in the kernel.                                                    |
++--------------------------------------------------------------------------+
+
+The information captured in a reverse space mapping record is as follows:
+
+.. code-block:: c
+
+	struct xfs_rmap_irec {
+	    xfs_agblock_t    rm_startblock;   /* extent start block */
+	    xfs_extlen_t     rm_blockcount;   /* extent length */
+	    uint64_t         rm_owner;        /* extent owner */
+	    uint64_t         rm_offset;       /* offset within the owner */
+	    unsigned int     rm_flags;        /* state flags */
+	};
+
+The first two fields capture the location and size of the physical space,
+in units of filesystem blocks.
+The owner field tells scrub which metadata structure or file inode have been
+assigned this space.
+For space allocated to files, the offset field tells scrub where the space was
+mapped within the file fork.
+Finally, the flags field provides extra information about the space usage --
+is this an attribute fork extent?  A file mapping btree extent?  Or an
+unwritten data extent?
+
+Online filesystem checking judges the consistency of each primary metadata
+record by comparing its information against all other space indices.
+The reverse mapping index plays a key role in the consistency checking process
+because it contains a centralized alternate copy of all space allocation
+information.
+Program runtime and ease of resource acquisition are the only real limits to
+what online checking can consult.
+For example, a file data extent mapping can be checked against:
+
+* The absence of an entry in the free space information.
+* The absence of an entry in the inode index.
+* The absence of an entry in the reference count data if the file is not
+  marked as having shared extents.
+* The correspondence of an entry in the reverse mapping information.
+
+There are several observations to make about reverse mapping indices:
+
+1. Reverse mappings can provide a positive affirmation of correctness if any of
+   the above primary metadata are in doubt.
+   The checking code for most primary metadata follows a path similar to the
+   one outlined above.
+
+2. Proving the consistency of secondary metadata with the primary metadata is
+   difficult because that requires a full scan of all primary space metadata,
+   which is very time intensive.
+   For example, checking a reverse mapping record for a file extent mapping
+   btree block requires locking the file and searching the entire btree to
+   confirm the block.
+   Instead, scrub relies on rigorous cross-referencing during the primary space
+   mapping structure checks.
+
+3. Consistency scans must use non-blocking lock acquisition primitives if the
+   required locking order is not the same order used by regular filesystem
+   operations.
+   For example, if the filesystem normally takes a file ILOCK before taking
+   the AGF buffer lock but scrub wants to take a file ILOCK while holding
+   an AGF buffer lock, scrub cannot block on that second acquisition.
+   This means that forward progress during this part of a scan of the reverse
+   mapping data cannot be guaranteed if system load is heavy.
+
+In summary, reverse mappings play a key role in reconstruction of primary
+metadata.
+The details of how these records are staged, written to disk, and committed
+into the filesystem are covered in subsequent sections.
+
+Checking and Cross-Referencing
+------------------------------
+
+The first step of checking a metadata structure is to examine every record
+contained within the structure and its relationship with the rest of the
+system.
+XFS contains multiple layers of checking to try to prevent inconsistent
+metadata from wreaking havoc on the system.
+Each of these layers contributes information that helps the kernel to make
+three decisions about the health of a metadata structure:
+
+- Is a part of this structure obviously corrupt (``XFS_SCRUB_OFLAG_CORRUPT``) ?
+- Is this structure inconsistent with the rest of the system
+  (``XFS_SCRUB_OFLAG_XCORRUPT``) ?
+- Is there so much damage around the filesystem that cross-referencing is not
+  possible (``XFS_SCRUB_OFLAG_XFAIL``) ?
+- Can the structure be optimized to improve performance or reduce the size of
+  metadata (``XFS_SCRUB_OFLAG_PREEN``) ?
+- Does the structure contain data that is not inconsistent but deserves review
+  by the system administrator (``XFS_SCRUB_OFLAG_WARNING``) ?
+
+The following sections describe how the metadata scrubbing process works.
+
+Metadata Buffer Verification
+````````````````````````````
+
+The lowest layer of metadata protection in XFS are the metadata verifiers built
+into the buffer cache.
+These functions perform inexpensive internal consistency checking of the block
+itself, and answer these questions:
+
+- Does the block belong to this filesystem?
+
+- Does the block belong to the structure that asked for the read?
+  This assumes that metadata blocks only have one owner, which is always true
+  in XFS.
+
+- Is the type of data stored in the block within a reasonable range of what
+  scrub is expecting?
+
+- Does the physical location of the block match the location it was read from?
+
+- Does the block checksum match the data?
+
+The scope of the protections here are very limited -- verifiers can only
+establish that the filesystem code is reasonably free of gross corruption bugs
+and that the storage system is reasonably competent at retrieval.
+Corruption problems observed at runtime cause the generation of health reports,
+failed system calls, and in the extreme case, filesystem shutdowns if the
+corrupt metadata force the cancellation of a dirty transaction.
+
+Every online fsck scrubbing function is expected to read every ondisk metadata
+block of a structure in the course of checking the structure.
+Corruption problems observed during a check are immediately reported to
+userspace as corruption; during a cross-reference, they are reported as a
+failure to cross-reference once the full examination is complete.
+Reads satisfied by a buffer already in cache (and hence already verified)
+bypass these checks.
+
+Internal Consistency Checks
+```````````````````````````
+
+After the buffer cache, the next level of metadata protection is the internal
+record verification code built into the filesystem.
+These checks are split between the buffer verifiers, the in-filesystem users of
+the buffer cache, and the scrub code itself, depending on the amount of higher
+level context required.
+The scope of checking is still internal to the block.
+These higher level checking functions answer these questions:
+
+- Does the type of data stored in the block match what scrub is expecting?
+
+- Does the block belong to the owning structure that asked for the read?
+
+- If the block contains records, do the records fit within the block?
+
+- If the block tracks internal free space information, is it consistent with
+  the record areas?
+
+- Are the records contained inside the block free of obvious corruptions?
+
+Record checks in this category are more rigorous and more time-intensive.
+For example, block pointers and inumbers are checked to ensure that they point
+within the dynamically allocated parts of an allocation group and within
+the filesystem.
+Names are checked for invalid characters, and flags are checked for invalid
+combinations.
+Other record attributes are checked for sensible values.
+Btree records spanning an interval of the btree keyspace are checked for
+correct order and lack of mergeability (except for file fork mappings).
+For performance reasons, regular code may skip some of these checks unless
+debugging is enabled or a write is about to occur.
+Scrub functions, of course, must check all possible problems.
+
+Validation of Userspace-Controlled Record Attributes
+````````````````````````````````````````````````````
+
+Various pieces of filesystem metadata are directly controlled by userspace.
+Because of this nature, validation work cannot be more precise than checking
+that a value is within the possible range.
+These fields include:
+
+- Superblock fields controlled by mount options
+- Filesystem labels
+- File timestamps
+- File permissions
+- File size
+- File flags
+- Names present in directory entries, extended attribute keys, and filesystem
+  labels
+- Extended attribute key namespaces
+- Extended attribute values
+- File data block contents
+- Quota limits
+- Quota timer expiration (if resource usage exceeds the soft limit)
+
+Cross-Referencing Space Metadata
+````````````````````````````````
+
+After internal block checks, the next higher level of checking is
+cross-referencing records between metadata structures.
+For regular runtime code, the cost of these checks is considered to be
+prohibitively expensive, but as scrub is dedicated to rooting out
+inconsistencies, it must pursue all avenues of inquiry.
+The exact set of cross-referencing is highly dependent on the context of the
+data structure being checked.
+
+The XFS btree code has keyspace scanning functions that online fsck uses to
+cross reference one structure with another.
+Specifically, scrub can scan the key space of an index to determine if that
+keyspace is fully, sparsely, or not at all mapped to records.
+For the reverse mapping btree, it is possible to mask parts of the key for the
+purposes of performing a keyspace scan so that scrub can decide if the rmap
+btree contains records mapping a certain extent of physical space without the
+sparsenses of the rest of the rmap keyspace getting in the way.
+
+Btree blocks undergo the following checks before cross-referencing:
+
+- Does the type of data stored in the block match what scrub is expecting?
+
+- Does the block belong to the owning structure that asked for the read?
+
+- Do the records fit within the block?
+
+- Are the records contained inside the block free of obvious corruptions?
+
+- Are the name hashes in the correct order?
+
+- Do node pointers within the btree point to valid block addresses for the type
+  of btree?
+
+- Do child pointers point towards the leaves?
+
+- Do sibling pointers point across the same level?
+
+- For each node block record, does the record key accurate reflect the contents
+  of the child block?
+
+Space allocation records are cross-referenced as follows:
+
+1. Any space mentioned by any metadata structure are cross-referenced as
+   follows:
+
+   - Does the reverse mapping index list only the appropriate owner as the
+     owner of each block?
+
+   - Are none of the blocks claimed as free space?
+
+   - If these aren't file data blocks, are none of the blocks claimed as space
+     shared by different owners?
+
+2. Btree blocks are cross-referenced as follows:
+
+   - Everything in class 1 above.
+
+   - If there's a parent node block, do the keys listed for this block match the
+     keyspace of this block?
+
+   - Do the sibling pointers point to valid blocks?  Of the same level?
+
+   - Do the child pointers point to valid blocks?  Of the next level down?
+
+3. Free space btree records are cross-referenced as follows:
+
+   - Everything in class 1 and 2 above.
+
+   - Does the reverse mapping index list no owners of this space?
+
+   - Is this space not claimed by the inode index for inodes?
+
+   - Is it not mentioned by the reference count index?
+
+   - Is there a matching record in the other free space btree?
+
+4. Inode btree records are cross-referenced as follows:
+
+   - Everything in class 1 and 2 above.
+
+   - Is there a matching record in free inode btree?
+
+   - Do cleared bits in the holemask correspond with inode clusters?
+
+   - Do set bits in the freemask correspond with inode records with zero link
+     count?
+
+5. Inode records are cross-referenced as follows:
+
+   - Everything in class 1.
+
+   - Do all the fields that summarize information about the file forks actually
+     match those forks?
+
+   - Does each inode with zero link count correspond to a record in the free
+     inode btree?
+
+6. File fork space mapping records are cross-referenced as follows:
+
+   - Everything in class 1 and 2 above.
+
+   - Is this space not mentioned by the inode btrees?
+
+   - If this is a CoW fork mapping, does it correspond to a CoW entry in the
+     reference count btree?
+
+7. Reference count records are cross-referenced as follows:
+
+   - Everything in class 1 and 2 above.
+
+   - Within the space subkeyspace of the rmap btree (that is to say, all
+     records mapped to a particular space extent and ignoring the owner info),
+     are there the same number of reverse mapping records for each block as the
+     reference count record claims?
+
+Proposed patchsets are the series to find gaps in
+`refcount btree
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=scrub-detect-refcount-gaps>`_,
+`inode btree
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=scrub-detect-inobt-gaps>`_, and
+`rmap btree
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=scrub-detect-rmapbt-gaps>`_ records;
+to find
+`mergeable records
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=scrub-detect-mergeable-records>`_;
+and to
+`improve cross referencing with rmap
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=scrub-strengthen-rmap-checking>`_
+before starting a repair.
+
+Checking Extended Attributes
+````````````````````````````
+
+Extended attributes implement a key-value store that enable fragments of data
+to be attached to any file.
+Both the kernel and userspace can access the keys and values, subject to
+namespace and privilege restrictions.
+Most typically these fragments are metadata about the file -- origins, security
+contexts, user-supplied labels, indexing information, etc.
+
+Names can be as long as 255 bytes and can exist in several different
+namespaces.
+Values can be as large as 64KB.
+A file's extended attributes are stored in blocks mapped by the attr fork.
+The mappings point to leaf blocks, remote value blocks, or dabtree blocks.
+Block 0 in the attribute fork is always the top of the structure, but otherwise
+each of the three types of blocks can be found at any offset in the attr fork.
+Leaf blocks contain attribute key records that point to the name and the value.
+Names are always stored elsewhere in the same leaf block.
+Values that are less than 3/4 the size of a filesystem block are also stored
+elsewhere in the same leaf block.
+Remote value blocks contain values that are too large to fit inside a leaf.
+If the leaf information exceeds a single filesystem block, a dabtree (also
+rooted at block 0) is created to map hashes of the attribute names to leaf
+blocks in the attr fork.
+
+Checking an extended attribute structure is not so straightforward due to the
+lack of separation between attr blocks and index blocks.
+Scrub must read each block mapped by the attr fork and ignore the non-leaf
+blocks:
+
+1. Walk the dabtree in the attr fork (if present) to ensure that there are no
+   irregularities in the blocks or dabtree mappings that do not point to
+   attr leaf blocks.
+
+2. Walk the blocks of the attr fork looking for leaf blocks.
+   For each entry inside a leaf:
+
+   a. Validate that the name does not contain invalid characters.
+
+   b. Read the attr value.
+      This performs a named lookup of the attr name to ensure the correctness
+      of the dabtree.
+      If the value is stored in a remote block, this also validates the
+      integrity of the remote value block.
+
+Checking and Cross-Referencing Directories
+``````````````````````````````````````````
+
+The filesystem directory tree is a directed acylic graph structure, with files
+constituting the nodes, and directory entries (dirents) constituting the edges.
+Directories are a special type of file containing a set of mappings from a
+255-byte sequence (name) to an inumber.
+These are called directory entries, or dirents for short.
+Each directory file must have exactly one directory pointing to the file.
+A root directory points to itself.
+Directory entries point to files of any type.
+Each non-directory file may have multiple directories point to it.
+
+In XFS, directories are implemented as a file containing up to three 32GB
+partitions.
+The first partition contains directory entry data blocks.
+Each data block contains variable-sized records associating a user-provided
+name with an inumber and, optionally, a file type.
+If the directory entry data grows beyond one block, the second partition (which
+exists as post-EOF extents) is populated with a block containing free space
+information and an index that maps hashes of the dirent names to directory data
+blocks in the first partition.
+This makes directory name lookups very fast.
+If this second partition grows beyond one block, the third partition is
+populated with a linear array of free space information for faster
+expansions.
+If the free space has been separated and the second partition grows again
+beyond one block, then a dabtree is used to map hashes of dirent names to
+directory data blocks.
+
+Checking a directory is pretty straightforward:
+
+1. Walk the dabtree in the second partition (if present) to ensure that there
+   are no irregularities in the blocks or dabtree mappings that do not point to
+   dirent blocks.
+
+2. Walk the blocks of the first partition looking for directory entries.
+   Each dirent is checked as follows:
+
+   a. Does the name contain no invalid characters?
+
+   b. Does the inumber correspond to an actual, allocated inode?
+
+   c. Does the child inode have a nonzero link count?
+
+   d. If a file type is included in the dirent, does it match the type of the
+      inode?
+
+   e. If the child is a subdirectory, does the child's dotdot pointer point
+      back to the parent?
+
+   f. If the directory has a second partition, perform a named lookup of the
+      dirent name to ensure the correctness of the dabtree.
+
+3. Walk the free space list in the third partition (if present) to ensure that
+   the free spaces it describes are really unused.
+
+Checking operations involving :ref:`parents <dirparent>` and
+:ref:`file link counts <nlinks>` are discussed in more detail in later
+sections.
+
+Checking Directory/Attribute Btrees
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+As stated in previous sections, the directory/attribute btree (dabtree) index
+maps user-provided names to improve lookup times by avoiding linear scans.
+Internally, it maps a 32-bit hash of the name to a block offset within the
+appropriate file fork.
+
+The internal structure of a dabtree closely resembles the btrees that record
+fixed-size metadata records -- each dabtree block contains a magic number, a
+checksum, sibling pointers, a UUID, a tree level, and a log sequence number.
+The format of leaf and node records are the same -- each entry points to the
+next level down in the hierarchy, with dabtree node records pointing to dabtree
+leaf blocks, and dabtree leaf records pointing to non-dabtree blocks elsewhere
+in the fork.
+
+Checking and cross-referencing the dabtree is very similar to what is done for
+space btrees:
+
+- Does the type of data stored in the block match what scrub is expecting?
+
+- Does the block belong to the owning structure that asked for the read?
+
+- Do the records fit within the block?
+
+- Are the records contained inside the block free of obvious corruptions?
+
+- Are the name hashes in the correct order?
+
+- Do node pointers within the dabtree point to valid fork offsets for dabtree
+  blocks?
+
+- Do leaf pointers within the dabtree point to valid fork offsets for directory
+  or attr leaf blocks?
+
+- Do child pointers point towards the leaves?
+
+- Do sibling pointers point across the same level?
+
+- For each dabtree node record, does the record key accurate reflect the
+  contents of the child dabtree block?
+
+- For each dabtree leaf record, does the record key accurate reflect the
+  contents of the directory or attr block?
+
+Cross-Referencing Summary Counters
+``````````````````````````````````
+
+XFS maintains three classes of summary counters: available resources, quota
+resource usage, and file link counts.
+
+In theory, the amount of available resources (data blocks, inodes, realtime
+extents) can be found by walking the entire filesystem.
+This would make for very slow reporting, so a transactional filesystem can
+maintain summaries of this information in the superblock.
+Cross-referencing these values against the filesystem metadata should be a
+simple matter of walking the free space and inode metadata in each AG and the
+realtime bitmap, but there are complications that will be discussed in
+:ref:`more detail <fscounters>` later.
+
+:ref:`Quota usage <quotacheck>` and :ref:`file link count <nlinks>`
+checking are sufficiently complicated to warrant separate sections.
+
+Post-Repair Reverification
+``````````````````````````
+
+After performing a repair, the checking code is run a second time to validate
+the new structure, and the results of the health assessment are recorded
+internally and returned to the calling process.
+This step is critical for enabling system administrator to monitor the status
+of the filesystem and the progress of any repairs.
+For developers, it is a useful means to judge the efficacy of error detection
+and correction in the online and offline checking tools.
+
+Eventual Consistency vs. Online Fsck
+------------------------------------
+
+Complex operations can make modifications to multiple per-AG data structures
+with a chain of transactions.
+These chains, once committed to the log, are restarted during log recovery if
+the system crashes while processing the chain.
+Because the AG header buffers are unlocked between transactions within a chain,
+online checking must coordinate with chained operations that are in progress to
+avoid incorrectly detecting inconsistencies due to pending chains.
+Furthermore, online repair must not run when operations are pending because
+the metadata are temporarily inconsistent with each other, and rebuilding is
+not possible.
+
+Only online fsck has this requirement of total consistency of AG metadata, and
+should be relatively rare as compared to filesystem change operations.
+Online fsck coordinates with transaction chains as follows:
+
+* For each AG, maintain a count of intent items targeting that AG.
+  The count should be bumped whenever a new item is added to the chain.
+  The count should be dropped when the filesystem has locked the AG header
+  buffers and finished the work.
+
+* When online fsck wants to examine an AG, it should lock the AG header
+  buffers to quiesce all transaction chains that want to modify that AG.
+  If the count is zero, proceed with the checking operation.
+  If it is nonzero, cycle the buffer locks to allow the chain to make forward
+  progress.
+
+This may lead to online fsck taking a long time to complete, but regular
+filesystem updates take precedence over background checking activity.
+Details about the discovery of this situation are presented in the
+:ref:`next section <chain_coordination>`, and details about the solution
+are presented :ref:`after that<intent_drains>`.
+
+.. _chain_coordination:
+
+Discovery of the Problem
+````````````````````````
+
+Midway through the development of online scrubbing, the fsstress tests
+uncovered a misinteraction between online fsck and compound transaction chains
+created by other writer threads that resulted in false reports of metadata
+inconsistency.
+The root cause of these reports is the eventual consistency model introduced by
+the expansion of deferred work items and compound transaction chains when
+reverse mapping and reflink were introduced.
+
+Originally, transaction chains were added to XFS to avoid deadlocks when
+unmapping space from files.
+Deadlock avoidance rules require that AGs only be locked in increasing order,
+which makes it impossible (say) to use a single transaction to free a space
+extent in AG 7 and then try to free a now superfluous block mapping btree block
+in AG 3.
+To avoid these kinds of deadlocks, XFS creates Extent Freeing Intent (EFI) log
+items to commit to freeing some space in one transaction while deferring the
+actual metadata updates to a fresh transaction.
+The transaction sequence looks like this:
+
+1. The first transaction contains a physical update to the file's block mapping
+   structures to remove the mapping from the btree blocks.
+   It then attaches to the in-memory transaction an action item to schedule
+   deferred freeing of space.
+   Concretely, each transaction maintains a list of ``struct
+   xfs_defer_pending`` objects, each of which maintains a list of ``struct
+   xfs_extent_free_item`` objects.
+   Returning to the example above, the action item tracks the freeing of both
+   the unmapped space from AG 7 and the block mapping btree (BMBT) block from
+   AG 3.
+   Deferred frees recorded in this manner are committed in the log by creating
+   an EFI log item from the ``struct xfs_extent_free_item`` object and
+   attaching the log item to the transaction.
+   When the log is persisted to disk, the EFI item is written into the ondisk
+   transaction record.
+   EFIs can list up to 16 extents to free, all sorted in AG order.
+
+2. The second transaction contains a physical update to the free space btrees
+   of AG 3 to release the former BMBT block and a second physical update to the
+   free space btrees of AG 7 to release the unmapped file space.
+   Observe that the physical updates are resequenced in the correct order
+   when possible.
+   Attached to the transaction is a an extent free done (EFD) log item.
+   The EFD contains a pointer to the EFI logged in transaction #1 so that log
+   recovery can tell if the EFI needs to be replayed.
+
+If the system goes down after transaction #1 is written back to the filesystem
+but before #2 is committed, a scan of the filesystem metadata would show
+inconsistent filesystem metadata because there would not appear to be any owner
+of the unmapped space.
+Happily, log recovery corrects this inconsistency for us -- when recovery finds
+an intent log item but does not find a corresponding intent done item, it will
+reconstruct the incore state of the intent item and finish it.
+In the example above, the log must replay both frees described in the recovered
+EFI to complete the recovery phase.
+
+There are subtleties to XFS' transaction chaining strategy to consider:
+
+* Log items must be added to a transaction in the correct order to prevent
+  conflicts with principal objects that are not held by the transaction.
+  In other words, all per-AG metadata updates for an unmapped block must be
+  completed before the last update to free the extent, and extents should not
+  be reallocated until that last update commits to the log.
+
+* AG header buffers are released between each transaction in a chain.
+  This means that other threads can observe an AG in an intermediate state,
+  but as long as the first subtlety is handled, this should not affect the
+  correctness of filesystem operations.
+
+* Unmounting the filesystem flushes all pending work to disk, which means that
+  offline fsck never sees the temporary inconsistencies caused by deferred
+  work item processing.
+
+In this manner, XFS employs a form of eventual consistency to avoid deadlocks
+and increase parallelism.
+
+During the design phase of the reverse mapping and reflink features, it was
+decided that it was impractical to cram all the reverse mapping updates for a
+single filesystem change into a single transaction because a single file
+mapping operation can explode into many small updates:
+
+* The block mapping update itself
+* A reverse mapping update for the block mapping update
+* Fixing the freelist
+* A reverse mapping update for the freelist fix
+
+* A shape change to the block mapping btree
+* A reverse mapping update for the btree update
+* Fixing the freelist (again)
+* A reverse mapping update for the freelist fix
+
+* An update to the reference counting information
+* A reverse mapping update for the refcount update
+* Fixing the freelist (a third time)
+* A reverse mapping update for the freelist fix
+
+* Freeing any space that was unmapped and not owned by any other file
+* Fixing the freelist (a fourth time)
+* A reverse mapping update for the freelist fix
+
+* Freeing the space used by the block mapping btree
+* Fixing the freelist (a fifth time)
+* A reverse mapping update for the freelist fix
+
+Free list fixups are not usually needed more than once per AG per transaction
+chain, but it is theoretically possible if space is very tight.
+For copy-on-write updates this is even worse, because this must be done once to
+remove the space from a staging area and again to map it into the file!
+
+To deal with this explosion in a calm manner, XFS expands its use of deferred
+work items to cover most reverse mapping updates and all refcount updates.
+This reduces the worst case size of transaction reservations by breaking the
+work into a long chain of small updates, which increases the degree of eventual
+consistency in the system.
+Again, this generally isn't a problem because XFS orders its deferred work
+items carefully to avoid resource reuse conflicts between unsuspecting threads.
+
+However, online fsck changes the rules -- remember that although physical
+updates to per-AG structures are coordinated by locking the buffers for AG
+headers, buffer locks are dropped between transactions.
+Once scrub acquires resources and takes locks for a data structure, it must do
+all the validation work without releasing the lock.
+If the main lock for a space btree is an AG header buffer lock, scrub may have
+interrupted another thread that is midway through finishing a chain.
+For example, if a thread performing a copy-on-write has completed a reverse
+mapping update but not the corresponding refcount update, the two AG btrees
+will appear inconsistent to scrub and an observation of corruption will be
+recorded.  This observation will not be correct.
+If a repair is attempted in this state, the results will be catastrophic!
+
+Several other solutions to this problem were evaluated upon discovery of this
+flaw and rejected:
+
+1. Add a higher level lock to allocation groups and require writer threads to
+   acquire the higher level lock in AG order before making any changes.
+   This would be very difficult to implement in practice because it is
+   difficult to determine which locks need to be obtained, and in what order,
+   without simulating the entire operation.
+   Performing a dry run of a file operation to discover necessary locks would
+   make the filesystem very slow.
+
+2. Make the deferred work coordinator code aware of consecutive intent items
+   targeting the same AG and have it hold the AG header buffers locked across
+   the transaction roll between updates.
+   This would introduce a lot of complexity into the coordinator since it is
+   only loosely coupled with the actual deferred work items.
+   It would also fail to solve the problem because deferred work items can
+   generate new deferred subtasks, but all subtasks must be complete before
+   work can start on a new sibling task.
+
+3. Teach online fsck to walk all transactions waiting for whichever lock(s)
+   protect the data structure being scrubbed to look for pending operations.
+   The checking and repair operations must factor these pending operations into
+   the evaluations being performed.
+   This solution is a nonstarter because it is *extremely* invasive to the main
+   filesystem.
+
+.. _intent_drains:
+
+Intent Drains
+`````````````
+
+Online fsck uses an atomic intent item counter and lock cycling to coordinate
+with transaction chains.
+There are two key properties to the drain mechanism.
+First, the counter is incremented when a deferred work item is *queued* to a
+transaction, and it is decremented after the associated intent done log item is
+*committed* to another transaction.
+The second property is that deferred work can be added to a transaction without
+holding an AG header lock, but per-AG work items cannot be marked done without
+locking that AG header buffer to log the physical updates and the intent done
+log item.
+The first property enables scrub to yield to running transaction chains, which
+is an explicit deprioritization of online fsck to benefit file operations.
+The second property of the drain is key to the correct coordination of scrub,
+since scrub will always be able to decide if a conflict is possible.
+
+For regular filesystem code, the drain works as follows:
+
+1. Call the appropriate subsystem function to add a deferred work item to a
+   transaction.
+
+2. The function calls ``xfs_defer_drain_bump`` to increase the counter.
+
+3. When the deferred item manager wants to finish the deferred work item, it
+   calls ``->finish_item`` to complete it.
+
+4. The ``->finish_item`` implementation logs some changes and calls
+   ``xfs_defer_drain_drop`` to decrease the sloppy counter and wake up any threads
+   waiting on the drain.
+
+5. The subtransaction commits, which unlocks the resource associated with the
+   intent item.
+
+For scrub, the drain works as follows:
+
+1. Lock the resource(s) associated with the metadata being scrubbed.
+   For example, a scan of the refcount btree would lock the AGI and AGF header
+   buffers.
+
+2. If the counter is zero (``xfs_defer_drain_busy`` returns false), there are no
+   chains in progress and the operation may proceed.
+
+3. Otherwise, release the resources grabbed in step 1.
+
+4. Wait for the intent counter to reach zero (``xfs_defer_drain_intents``), then go
+   back to step 1 unless a signal has been caught.
+
+To avoid polling in step 4, the drain provides a waitqueue for scrub threads to
+be woken up whenever the intent count drops to zero.
+
+The proposed patchset is the
+`scrub intent drain series
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=scrub-drain-intents>`_.
+
+.. _jump_labels:
+
+Static Keys (aka Jump Label Patching)
+`````````````````````````````````````
+
+Online fsck for XFS separates the regular filesystem from the checking and
+repair code as much as possible.
+However, there are a few parts of online fsck (such as the intent drains, and
+later, live update hooks) where it is useful for the online fsck code to know
+what's going on in the rest of the filesystem.
+Since it is not expected that online fsck will be constantly running in the
+background, it is very important to minimize the runtime overhead imposed by
+these hooks when online fsck is compiled into the kernel but not actively
+running on behalf of userspace.
+Taking locks in the hot path of a writer thread to access a data structure only
+to find that no further action is necessary is expensive -- on the author's
+computer, this have an overhead of 40-50ns per access.
+Fortunately, the kernel supports dynamic code patching, which enables XFS to
+replace a static branch to hook code with ``nop`` sleds when online fsck isn't
+running.
+This sled has an overhead of however long it takes the instruction decoder to
+skip past the sled, which seems to be on the order of less than 1ns and
+does not access memory outside of instruction fetching.
+
+When online fsck enables the static key, the sled is replaced with an
+unconditional branch to call the hook code.
+The switchover is quite expensive (~22000ns) but is paid entirely by the
+program that invoked online fsck, and can be amortized if multiple threads
+enter online fsck at the same time, or if multiple filesystems are being
+checked at the same time.
+Changing the branch direction requires taking the CPU hotplug lock, and since
+CPU initialization requires memory allocation, online fsck must be careful not
+to change a static key while holding any locks or resources that could be
+accessed in the memory reclaim paths.
+To minimize contention on the CPU hotplug lock, care should be taken not to
+enable or disable static keys unnecessarily.
+
+Because static keys are intended to minimize hook overhead for regular
+filesystem operations when xfs_scrub is not running, the intended usage
+patterns are as follows:
+
+- The hooked part of XFS should declare a static-scoped static key that
+  defaults to false.
+  The ``DEFINE_STATIC_KEY_FALSE`` macro takes care of this.
+  The static key itself should be declared as a ``static`` variable.
+
+- When deciding to invoke code that's only used by scrub, the regular
+  filesystem should call the ``static_branch_unlikely`` predicate to avoid the
+  scrub-only hook code if the static key is not enabled.
+
+- The regular filesystem should export helper functions that call
+  ``static_branch_inc`` to enable and ``static_branch_dec`` to disable the
+  static key.
+  Wrapper functions make it easy to compile out the relevant code if the kernel
+  distributor turns off online fsck at build time.
+
+- Scrub functions wanting to turn on scrub-only XFS functionality should call
+  the ``xchk_fsgates_enable`` from the setup function to enable a specific
+  hook.
+  This must be done before obtaining any resources that are used by memory
+  reclaim.
+  Callers had better be sure they really need the functionality gated by the
+  static key; the ``TRY_HARDER`` flag is useful here.
+
+Online scrub has resource acquisition helpers (e.g. ``xchk_perag_lock``) to
+handle locking AGI and AGF buffers for all scrubber functions.
+If it detects a conflict between scrub and the running transactions, it will
+try to wait for intents to complete.
+If the caller of the helper has not enabled the static key, the helper will
+return -EDEADLOCK, which should result in the scrub being restarted with the
+``TRY_HARDER`` flag set.
+The scrub setup function should detect that flag, enable the static key, and
+try the scrub again.
+Scrub teardown disables all static keys obtained by ``xchk_fsgates_enable``.
+
+For more information, please see the kernel documentation of
+Documentation/staging/static-keys.rst.
+
+.. _xfile:
+
+Pageable Kernel Memory
+----------------------
+
+Some online checking functions work by scanning the filesystem to build a
+shadow copy of an ondisk metadata structure in memory and comparing the two
+copies.
+For online repair to rebuild a metadata structure, it must compute the record
+set that will be stored in the new structure before it can persist that new
+structure to disk.
+Ideally, repairs complete with a single atomic commit that introduces
+a new data structure.
+To meet these goals, the kernel needs to collect a large amount of information
+in a place that doesn't require the correct operation of the filesystem.
+
+Kernel memory isn't suitable because:
+
+* Allocating a contiguous region of memory to create a C array is very
+  difficult, especially on 32-bit systems.
+
+* Linked lists of records introduce double pointer overhead which is very high
+  and eliminate the possibility of indexed lookups.
+
+* Kernel memory is pinned, which can drive the system into OOM conditions.
+
+* The system might not have sufficient memory to stage all the information.
+
+At any given time, online fsck does not need to keep the entire record set in
+memory, which means that individual records can be paged out if necessary.
+Continued development of online fsck demonstrated that the ability to perform
+indexed data storage would also be very useful.
+Fortunately, the Linux kernel already has a facility for byte-addressable and
+pageable storage: tmpfs.
+In-kernel graphics drivers (most notably i915) take advantage of tmpfs files
+to store intermediate data that doesn't need to be in memory at all times, so
+that usage precedent is already established.
+Hence, the ``xfile`` was born!
+
++--------------------------------------------------------------------------+
+| **Historical Sidebar**:                                                  |
++--------------------------------------------------------------------------+
+| The first edition of online repair inserted records into a new btree as  |
+| it found them, which failed because filesystem could shut down with a    |
+| built data structure, which would be live after recovery finished.       |
+|                                                                          |
+| The second edition solved the half-rebuilt structure problem by storing  |
+| everything in memory, but frequently ran the system out of memory.       |
+|                                                                          |
+| The third edition solved the OOM problem by using linked lists, but the  |
+| memory overhead of the list pointers was extreme.                        |
++--------------------------------------------------------------------------+
+
+xfile Access Models
+```````````````````
+
+A survey of the intended uses of xfiles suggested these use cases:
+
+1. Arrays of fixed-sized records (space management btrees, directory and
+   extended attribute entries)
+
+2. Sparse arrays of fixed-sized records (quotas and link counts)
+
+3. Large binary objects (BLOBs) of variable sizes (directory and extended
+   attribute names and values)
+
+4. Staging btrees in memory (reverse mapping btrees)
+
+5. Arbitrary contents (realtime space management)
+
+To support the first four use cases, high level data structures wrap the xfile
+to share functionality between online fsck functions.
+The rest of this section discusses the interfaces that the xfile presents to
+four of those five higher level data structures.
+The fifth use case is discussed in the :ref:`realtime summary <rtsummary>` case
+study.
+
+The most general storage interface supported by the xfile enables the reading
+and writing of arbitrary quantities of data at arbitrary offsets in the xfile.
+This capability is provided by ``xfile_pread`` and ``xfile_pwrite`` functions,
+which behave similarly to their userspace counterparts.
+XFS is very record-based, which suggests that the ability to load and store
+complete records is important.
+To support these cases, a pair of ``xfile_obj_load`` and ``xfile_obj_store``
+functions are provided to read and persist objects into an xfile.
+They are internally the same as pread and pwrite, except that they treat any
+error as an out of memory error.
+For online repair, squashing error conditions in this manner is an acceptable
+behavior because the only reaction is to abort the operation back to userspace.
+All five xfile usecases can be serviced by these four functions.
+
+However, no discussion of file access idioms is complete without answering the
+question, "But what about mmap?"
+It is convenient to access storage directly with pointers, just like userspace
+code does with regular memory.
+Online fsck must not drive the system into OOM conditions, which means that
+xfiles must be responsive to memory reclamation.
+tmpfs can only push a pagecache folio to the swap cache if the folio is neither
+pinned nor locked, which means the xfile must not pin too many folios.
+
+Short term direct access to xfile contents is done by locking the pagecache
+folio and mapping it into kernel address space.
+Programmatic access (e.g. pread and pwrite) uses this mechanism.
+Folio locks are not supposed to be held for long periods of time, so long
+term direct access to xfile contents is done by bumping the folio refcount,
+mapping it into kernel address space, and dropping the folio lock.
+These long term users *must* be responsive to memory reclaim by hooking into
+the shrinker infrastructure to know when to release folios.
+
+The ``xfile_get_page`` and ``xfile_put_page`` functions are provided to
+retrieve the (locked) folio that backs part of an xfile and to release it.
+The only code to use these folio lease functions are the xfarray
+:ref:`sorting<xfarray_sort>` algorithms and the :ref:`in-memory
+btrees<xfbtree>`.
+
+xfile Access Coordination
+`````````````````````````
+
+For security reasons, xfiles must be owned privately by the kernel.
+They are marked ``S_PRIVATE`` to prevent interference from the security system,
+must never be mapped into process file descriptor tables, and their pages must
+never be mapped into userspace processes.
+
+To avoid locking recursion issues with the VFS, all accesses to the shmfs file
+are performed by manipulating the page cache directly.
+xfile writers call the ``->write_begin`` and ``->write_end`` functions of the
+xfile's address space to grab writable pages, copy the caller's buffer into the
+page, and release the pages.
+xfile readers call ``shmem_read_mapping_page_gfp`` to grab pages directly
+before copying the contents into the caller's buffer.
+In other words, xfiles ignore the VFS read and write code paths to avoid
+having to create a dummy ``struct kiocb`` and to avoid taking inode and
+freeze locks.
+tmpfs cannot be frozen, and xfiles must not be exposed to userspace.
+
+If an xfile is shared between threads to stage repairs, the caller must provide
+its own locks to coordinate access.
+For example, if a scrub function stores scan results in an xfile and needs
+other threads to provide updates to the scanned data, the scrub function must
+provide a lock for all threads to share.
+
+.. _xfarray:
+
+Arrays of Fixed-Sized Records
+`````````````````````````````
+
+In XFS, each type of indexed space metadata (free space, inodes, reference
+counts, file fork space, and reverse mappings) consists of a set of fixed-size
+records indexed with a classic B+ tree.
+Directories have a set of fixed-size dirent records that point to the names,
+and extended attributes have a set of fixed-size attribute keys that point to
+names and values.
+Quota counters and file link counters index records with numbers.
+During a repair, scrub needs to stage new records during the gathering step and
+retrieve them during the btree building step.
+
+Although this requirement can be satisfied by calling the read and write
+methods of the xfile directly, it is simpler for callers for there to be a
+higher level abstraction to take care of computing array offsets, to provide
+iterator functions, and to deal with sparse records and sorting.
+The ``xfarray`` abstraction presents a linear array for fixed-size records atop
+the byte-accessible xfile.
+
+.. _xfarray_access_patterns:
+
+Array Access Patterns
+^^^^^^^^^^^^^^^^^^^^^
+
+Array access patterns in online fsck tend to fall into three categories.
+Iteration of records is assumed to be necessary for all cases and will be
+covered in the next section.
+
+The first type of caller handles records that are indexed by position.
+Gaps may exist between records, and a record may be updated multiple times
+during the collection step.
+In other words, these callers want a sparse linearly addressed table file.
+The typical use case are quota records or file link count records.
+Access to array elements is performed programmatically via ``xfarray_load`` and
+``xfarray_store`` functions, which wrap the similarly-named xfile functions to
+provide loading and storing of array elements at arbitrary array indices.
+Gaps are defined to be null records, and null records are defined to be a
+sequence of all zero bytes.
+Null records are detected by calling ``xfarray_element_is_null``.
+They are created either by calling ``xfarray_unset`` to null out an existing
+record or by never storing anything to an array index.
+
+The second type of caller handles records that are not indexed by position
+and do not require multiple updates to a record.
+The typical use case here is rebuilding space btrees and key/value btrees.
+These callers can add records to the array without caring about array indices
+via the ``xfarray_append`` function, which stores a record at the end of the
+array.
+For callers that require records to be presentable in a specific order (e.g.
+rebuilding btree data), the ``xfarray_sort`` function can arrange the sorted
+records; this function will be covered later.
+
+The third type of caller is a bag, which is useful for counting records.
+The typical use case here is constructing space extent reference counts from
+reverse mapping information.
+Records can be put in the bag in any order, they can be removed from the bag
+at any time, and uniqueness of records is left to callers.
+The ``xfarray_store_anywhere`` function is used to insert a record in any
+null record slot in the bag; and the ``xfarray_unset`` function removes a
+record from the bag.
+
+The proposed patchset is the
+`big in-memory array
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=big-array>`_.
+
+Iterating Array Elements
+^^^^^^^^^^^^^^^^^^^^^^^^
+
+Most users of the xfarray require the ability to iterate the records stored in
+the array.
+Callers can probe every possible array index with the following:
+
+.. code-block:: c
+
+	xfarray_idx_t i;
+	foreach_xfarray_idx(array, i) {
+	    xfarray_load(array, i, &rec);
+
+	    /* do something with rec */
+	}
+
+All users of this idiom must be prepared to handle null records or must already
+know that there aren't any.
+
+For xfarray users that want to iterate a sparse array, the ``xfarray_iter``
+function ignores indices in the xfarray that have never been written to by
+calling ``xfile_seek_data`` (which internally uses ``SEEK_DATA``) to skip areas
+of the array that are not populated with memory pages.
+Once it finds a page, it will skip the zeroed areas of the page.
+
+.. code-block:: c
+
+	xfarray_idx_t i = XFARRAY_CURSOR_INIT;
+	while ((ret = xfarray_iter(array, &i, &rec)) == 1) {
+	    /* do something with rec */
+	}
+
+.. _xfarray_sort:
+
+Sorting Array Elements
+^^^^^^^^^^^^^^^^^^^^^^
+
+During the fourth demonstration of online repair, a community reviewer remarked
+that for performance reasons, online repair ought to load batches of records
+into btree record blocks instead of inserting records into a new btree one at a
+time.
+The btree insertion code in XFS is responsible for maintaining correct ordering
+of the records, so naturally the xfarray must also support sorting the record
+set prior to bulk loading.
+
+Case Study: Sorting xfarrays
+~~~~~~~~~~~~~~~~~~~~~~~~~~~~
+
+The sorting algorithm used in the xfarray is actually a combination of adaptive
+quicksort and a heapsort subalgorithm in the spirit of
+`Sedgewick <https://algs4.cs.princeton.edu/23quicksort/>`_ and
+`pdqsort <https://github.com/orlp/pdqsort>`_, with customizations for the Linux
+kernel.
+To sort records in a reasonably short amount of time, ``xfarray`` takes
+advantage of the binary subpartitioning offered by quicksort, but it also uses
+heapsort to hedge against performance collapse if the chosen quicksort pivots
+are poor.
+Both algorithms are (in general) O(n * lg(n)), but there is a wide performance
+gulf between the two implementations.
+
+The Linux kernel already contains a reasonably fast implementation of heapsort.
+It only operates on regular C arrays, which limits the scope of its usefulness.
+There are two key places where the xfarray uses it:
+
+* Sorting any record subset backed by a single xfile page.
+
+* Loading a small number of xfarray records from potentially disparate parts
+  of the xfarray into a memory buffer, and sorting the buffer.
+
+In other words, ``xfarray`` uses heapsort to constrain the nested recursion of
+quicksort, thereby mitigating quicksort's worst runtime behavior.
+
+Choosing a quicksort pivot is a tricky business.
+A good pivot splits the set to sort in half, leading to the divide and conquer
+behavior that is crucial to  O(n * lg(n)) performance.
+A poor pivot barely splits the subset at all, leading to O(n\ :sup:`2`)
+runtime.
+The xfarray sort routine tries to avoid picking a bad pivot by sampling nine
+records into a memory buffer and using the kernel heapsort to identify the
+median of the nine.
+
+Most modern quicksort implementations employ Tukey's "ninther" to select a
+pivot from a classic C array.
+Typical ninther implementations pick three unique triads of records, sort each
+of the triads, and then sort the middle value of each triad to determine the
+ninther value.
+As stated previously, however, xfile accesses are not entirely cheap.
+It turned out to be much more performant to read the nine elements into a
+memory buffer, run the kernel's in-memory heapsort on the buffer, and choose
+the 4th element of that buffer as the pivot.
+Tukey's ninthers are described in J. W. Tukey, `The ninther, a technique for
+low-effort robust (resistant) location in large samples`, in *Contributions to
+Survey Sampling and Applied Statistics*, edited by H. David, (Academic Press,
+1978), pp. 251–257.
+
+The partitioning of quicksort is fairly textbook -- rearrange the record
+subset around the pivot, then set up the current and next stack frames to
+sort with the larger and the smaller halves of the pivot, respectively.
+This keeps the stack space requirements to log2(record count).
+
+As a final performance optimization, the hi and lo scanning phase of quicksort
+keeps examined xfile pages mapped in the kernel for as long as possible to
+reduce map/unmap cycles.
+Surprisingly, this reduces overall sort runtime by nearly half again after
+accounting for the application of heapsort directly onto xfile pages.
+
+.. _xfblob:
+
+Blob Storage
+````````````
+
+Extended attributes and directories add an additional requirement for staging
+records: arbitrary byte sequences of finite length.
+Each directory entry record needs to store entry name,
+and each extended attribute needs to store both the attribute name and value.
+The names, keys, and values can consume a large amount of memory, so the
+``xfblob`` abstraction was created to simplify management of these blobs
+atop an xfile.
+
+Blob arrays provide ``xfblob_load`` and ``xfblob_store`` functions to retrieve
+and persist objects.
+The store function returns a magic cookie for every object that it persists.
+Later, callers provide this cookie to the ``xblob_load`` to recall the object.
+The ``xfblob_free`` function frees a specific blob, and the ``xfblob_truncate``
+function frees them all because compaction is not needed.
+
+The details of repairing directories and extended attributes will be discussed
+in a subsequent section about atomic extent swapping.
+However, it should be noted that these repair functions only use blob storage
+to cache a small number of entries before adding them to a temporary ondisk
+file, which is why compaction is not required.
+
+The proposed patchset is at the start of the
+`extended attribute repair
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-xattrs>`_ series.
+
+.. _xfbtree:
+
+In-Memory B+Trees
+`````````````````
+
+The chapter about :ref:`secondary metadata<secondary_metadata>` mentioned that
+checking and repairing of secondary metadata commonly requires coordination
+between a live metadata scan of the filesystem and writer threads that are
+updating that metadata.
+Keeping the scan data up to date requires requires the ability to propagate
+metadata updates from the filesystem into the data being collected by the scan.
+This *can* be done by appending concurrent updates into a separate log file and
+applying them before writing the new metadata to disk, but this leads to
+unbounded memory consumption if the rest of the system is very busy.
+Another option is to skip the side-log and commit live updates from the
+filesystem directly into the scan data, which trades more overhead for a lower
+maximum memory requirement.
+In both cases, the data structure holding the scan results must support indexed
+access to perform well.
+
+Given that indexed lookups of scan data is required for both strategies, online
+fsck employs the second strategy of committing live updates directly into
+scan data.
+Because xfarrays are not indexed and do not enforce record ordering, they
+are not suitable for this task.
+Conveniently, however, XFS has a library to create and maintain ordered reverse
+mapping records: the existing rmap btree code!
+If only there was a means to create one in memory.
+
+Recall that the :ref:`xfile <xfile>` abstraction represents memory pages as a
+regular file, which means that the kernel can create byte or block addressable
+virtual address spaces at will.
+The XFS buffer cache specializes in abstracting IO to block-oriented  address
+spaces, which means that adaptation of the buffer cache to interface with
+xfiles enables reuse of the entire btree library.
+Btrees built atop an xfile are collectively known as ``xfbtrees``.
+The next few sections describe how they actually work.
+
+The proposed patchset is the
+`in-memory btree
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=in-memory-btrees>`_
+series.
+
+Using xfiles as a Buffer Cache Target
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+Two modifications are necessary to support xfiles as a buffer cache target.
+The first is to make it possible for the ``struct xfs_buftarg`` structure to
+host the ``struct xfs_buf`` rhashtable, because normally those are held by a
+per-AG structure.
+The second change is to modify the buffer ``ioapply`` function to "read" cached
+pages from the xfile and "write" cached pages back to the xfile.
+Multiple access to individual buffers is controlled by the ``xfs_buf`` lock,
+since the xfile does not provide any locking on its own.
+With this adaptation in place, users of the xfile-backed buffer cache use
+exactly the same APIs as users of the disk-backed buffer cache.
+The separation between xfile and buffer cache implies higher memory usage since
+they do not share pages, but this property could some day enable transactional
+updates to an in-memory btree.
+Today, however, it simply eliminates the need for new code.
+
+Space Management with an xfbtree
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+Space management for an xfile is very simple -- each btree block is one memory
+page in size.
+These blocks use the same header format as an on-disk btree, but the in-memory
+block verifiers ignore the checksums, assuming that xfile memory is no more
+corruption-prone than regular DRAM.
+Reusing existing code here is more important than absolute memory efficiency.
+
+The very first block of an xfile backing an xfbtree contains a header block.
+The header describes the owner, height, and the block number of the root
+xfbtree block.
+
+To allocate a btree block, use ``xfile_seek_data`` to find a gap in the file.
+If there are no gaps, create one by extending the length of the xfile.
+Preallocate space for the block with ``xfile_prealloc``, and hand back the
+location.
+To free an xfbtree block, use ``xfile_discard`` (which internally uses
+``FALLOC_FL_PUNCH_HOLE``) to remove the memory page from the xfile.
+
+Populating an xfbtree
+^^^^^^^^^^^^^^^^^^^^^
+
+An online fsck function that wants to create an xfbtree should proceed as
+follows:
+
+1. Call ``xfile_create`` to create an xfile.
+
+2. Call ``xfs_alloc_memory_buftarg`` to create a buffer cache target structure
+   pointing to the xfile.
+
+3. Pass the buffer cache target, buffer ops, and other information to
+   ``xfbtree_create`` to write an initial tree header and root block to the
+   xfile.
+   Each btree type should define a wrapper that passes necessary arguments to
+   the creation function.
+   For example, rmap btrees define ``xfs_rmapbt_mem_create`` to take care of
+   all the necessary details for callers.
+   A ``struct xfbtree`` object will be returned.
+
+4. Pass the xfbtree object to the btree cursor creation function for the
+   btree type.
+   Following the example above, ``xfs_rmapbt_mem_cursor`` takes care of this
+   for callers.
+
+5. Pass the btree cursor to the regular btree functions to make queries against
+   and to update the in-memory btree.
+   For example, a btree cursor for an rmap xfbtree can be passed to the
+   ``xfs_rmap_*`` functions just like any other btree cursor.
+   See the :ref:`next section<xfbtree_commit>` for information on dealing with
+   xfbtree updates that are logged to a transaction.
+
+6. When finished, delete the btree cursor, destroy the xfbtree object, free the
+   buffer target, and the destroy the xfile to release all resources.
+
+.. _xfbtree_commit:
+
+Committing Logged xfbtree Buffers
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+Although it is a clever hack to reuse the rmap btree code to handle the staging
+structure, the ephemeral nature of the in-memory btree block storage presents
+some challenges of its own.
+The XFS transaction manager must not commit buffer log items for buffers backed
+by an xfile because the log format does not understand updates for devices
+other than the data device.
+An ephemeral xfbtree probably will not exist by the time the AIL checkpoints
+log transactions back into the filesystem, and certainly won't exist during
+log recovery.
+For these reasons, any code updating an xfbtree in transaction context must
+remove the buffer log items from the transaction and write the updates into the
+backing xfile before committing or cancelling the transaction.
+
+The ``xfbtree_trans_commit`` and ``xfbtree_trans_cancel`` functions implement
+this functionality as follows:
+
+1. Find each buffer log item whose buffer targets the xfile.
+
+2. Record the dirty/ordered status of the log item.
+
+3. Detach the log item from the buffer.
+
+4. Queue the buffer to a special delwri list.
+
+5. Clear the transaction dirty flag if the only dirty log items were the ones
+   that were detached in step 3.
+
+6. Submit the delwri list to commit the changes to the xfile, if the updates
+   are being committed.
+
+After removing xfile logged buffers from the transaction in this manner, the
+transaction can be committed or cancelled.
+
+Bulk Loading of Ondisk B+Trees
+------------------------------
+
+As mentioned previously, early iterations of online repair built new btree
+structures by creating a new btree and adding observations individually.
+Loading a btree one record at a time had a slight advantage of not requiring
+the incore records to be sorted prior to commit, but was very slow and leaked
+blocks if the system went down during a repair.
+Loading records one at a time also meant that repair could not control the
+loading factor of the blocks in the new btree.
+
+Fortunately, the venerable ``xfs_repair`` tool had a more efficient means for
+rebuilding a btree index from a collection of records -- bulk btree loading.
+This was implemented rather inefficiently code-wise, since ``xfs_repair``
+had separate copy-pasted implementations for each btree type.
+
+To prepare for online fsck, each of the four bulk loaders were studied, notes
+were taken, and the four were refactored into a single generic btree bulk
+loading mechanism.
+Those notes in turn have been refreshed and are presented below.
+
+Geometry Computation
+````````````````````
+
+The zeroth step of bulk loading is to assemble the entire record set that will
+be stored in the new btree, and sort the records.
+Next, call ``xfs_btree_bload_compute_geometry`` to compute the shape of the
+btree from the record set, the type of btree, and any load factor preferences.
+This information is required for resource reservation.
+
+First, the geometry computation computes the minimum and maximum records that
+will fit in a leaf block from the size of a btree block and the size of the
+block header.
+Roughly speaking, the maximum number of records is::
+
+        maxrecs = (block_size - header_size) / record_size
+
+The XFS design specifies that btree blocks should be merged when possible,
+which means the minimum number of records is half of maxrecs::
+
+        minrecs = maxrecs / 2
+
+The next variable to determine is the desired loading factor.
+This must be at least minrecs and no more than maxrecs.
+Choosing minrecs is undesirable because it wastes half the block.
+Choosing maxrecs is also undesirable because adding a single record to each
+newly rebuilt leaf block will cause a tree split, which causes a noticeable
+drop in performance immediately afterwards.
+The default loading factor was chosen to be 75% of maxrecs, which provides a
+reasonably compact structure without any immediate split penalties::
+
+        default_load_factor = (maxrecs + minrecs) / 2
+
+If space is tight, the loading factor will be set to maxrecs to try to avoid
+running out of space::
+
+        leaf_load_factor = enough space ? default_load_factor : maxrecs
+
+Load factor is computed for btree node blocks using the combined size of the
+btree key and pointer as the record size::
+
+        maxrecs = (block_size - header_size) / (key_size + ptr_size)
+        minrecs = maxrecs / 2
+        node_load_factor = enough space ? default_load_factor : maxrecs
+
+Once that's done, the number of leaf blocks required to store the record set
+can be computed as::
+
+        leaf_blocks = ceil(record_count / leaf_load_factor)
+
+The number of node blocks needed to point to the next level down in the tree
+is computed as::
+
+        n_blocks = (n == 0 ? leaf_blocks : node_blocks[n])
+        node_blocks[n + 1] = ceil(n_blocks / node_load_factor)
+
+The entire computation is performed recursively until the current level only
+needs one block.
+The resulting geometry is as follows:
+
+- For AG-rooted btrees, this level is the root level, so the height of the new
+  tree is ``level + 1`` and the space needed is the summation of the number of
+  blocks on each level.
+
+- For inode-rooted btrees where the records in the top level do not fit in the
+  inode fork area, the height is ``level + 2``, the space needed is the
+  summation of the number of blocks on each level, and the inode fork points to
+  the root block.
+
+- For inode-rooted btrees where the records in the top level can be stored in
+  the inode fork area, then the root block can be stored in the inode, the
+  height is ``level + 1``, and the space needed is one less than the summation
+  of the number of blocks on each level.
+  This only becomes relevant when non-bmap btrees gain the ability to root in
+  an inode, which is a future patchset and only included here for completeness.
+
+.. _newbt:
+
+Reserving New B+Tree Blocks
+```````````````````````````
+
+Once repair knows the number of blocks needed for the new btree, it allocates
+those blocks using the free space information.
+Each reserved extent is tracked separately by the btree builder state data.
+To improve crash resilience, the reservation code also logs an Extent Freeing
+Intent (EFI) item in the same transaction as each space allocation and attaches
+its in-memory ``struct xfs_extent_free_item`` object to the space reservation.
+If the system goes down, log recovery will use the unfinished EFIs to free the
+unused space, the free space, leaving the filesystem unchanged.
+
+Each time the btree builder claims a block for the btree from a reserved
+extent, it updates the in-memory reservation to reflect the claimed space.
+Block reservation tries to allocate as much contiguous space as possible to
+reduce the number of EFIs in play.
+
+While repair is writing these new btree blocks, the EFIs created for the space
+reservations pin the tail of the ondisk log.
+It's possible that other parts of the system will remain busy and push the head
+of the log towards the pinned tail.
+To avoid livelocking the filesystem, the EFIs must not pin the tail of the log
+for too long.
+To alleviate this problem, the dynamic relogging capability of the deferred ops
+mechanism is reused here to commit a transaction at the log head containing an
+EFD for the old EFI and new EFI at the head.
+This enables the log to release the old EFI to keep the log moving forwards.
+
+EFIs have a role to play during the commit and reaping phases; please see the
+next section and the section about :ref:`reaping<reaping>` for more details.
+
+Proposed patchsets are the
+`bitmap rework
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-bitmap-rework>`_
+and the
+`preparation for bulk loading btrees
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-prep-for-bulk-loading>`_.
+
+
+Writing the New Tree
+````````````````````
+
+This part is pretty simple -- the btree builder (``xfs_btree_bulkload``) claims
+a block from the reserved list, writes the new btree block header, fills the
+rest of the block with records, and adds the new leaf block to a list of
+written blocks::
+
+  ┌────┐
+  │leaf│
+  │RRR │
+  └────┘
+
+Sibling pointers are set every time a new block is added to the level::
+
+  ┌────┐ ┌────┐ ┌────┐ ┌────┐
+  │leaf│→│leaf│→│leaf│→│leaf│
+  │RRR │←│RRR │←│RRR │←│RRR │
+  └────┘ └────┘ └────┘ └────┘
+
+When it finishes writing the record leaf blocks, it moves on to the node
+blocks
+To fill a node block, it walks each block in the next level down in the tree
+to compute the relevant keys and write them into the parent node::
+
+      ┌────┐       ┌────┐
+      │node│──────→│node│
+      │PP  │←──────│PP  │
+      └────┘       └────┘
+      ↙   ↘         ↙   ↘
+  ┌────┐ ┌────┐ ┌────┐ ┌────┐
+  │leaf│→│leaf│→│leaf│→│leaf│
+  │RRR │←│RRR │←│RRR │←│RRR │
+  └────┘ └────┘ └────┘ └────┘
+
+When it reaches the root level, it is ready to commit the new btree!::
+
+          ┌─────────┐
+          │  root   │
+          │   PP    │
+          └─────────┘
+          ↙         ↘
+      ┌────┐       ┌────┐
+      │node│──────→│node│
+      │PP  │←──────│PP  │
+      └────┘       └────┘
+      ↙   ↘         ↙   ↘
+  ┌────┐ ┌────┐ ┌────┐ ┌────┐
+  │leaf│→│leaf│→│leaf│→│leaf│
+  │RRR │←│RRR │←│RRR │←│RRR │
+  └────┘ └────┘ └────┘ └────┘
+
+The first step to commit the new btree is to persist the btree blocks to disk
+synchronously.
+This is a little complicated because a new btree block could have been freed
+in the recent past, so the builder must use ``xfs_buf_delwri_queue_here`` to
+remove the (stale) buffer from the AIL list before it can write the new blocks
+to disk.
+Blocks are queued for IO using a delwri list and written in one large batch
+with ``xfs_buf_delwri_submit``.
+
+Once the new blocks have been persisted to disk, control returns to the
+individual repair function that called the bulk loader.
+The repair function must log the location of the new root in a transaction,
+clean up the space reservations that were made for the new btree, and reap the
+old metadata blocks:
+
+1. Commit the location of the new btree root.
+
+2. For each incore reservation:
+
+   a. Log Extent Freeing Done (EFD) items for all the space that was consumed
+      by the btree builder.  The new EFDs must point to the EFIs attached to
+      the reservation to prevent log recovery from freeing the new blocks.
+
+   b. For unclaimed portions of incore reservations, create a regular deferred
+      extent free work item to be free the unused space later in the
+      transaction chain.
+
+   c. The EFDs and EFIs logged in steps 2a and 2b must not overrun the
+      reservation of the committing transaction.
+      If the btree loading code suspects this might be about to happen, it must
+      call ``xrep_defer_finish`` to clear out the deferred work and obtain a
+      fresh transaction.
+
+3. Clear out the deferred work a second time to finish the commit and clean
+   the repair transaction.
+
+The transaction rolling in steps 2c and 3 represent a weakness in the repair
+algorithm, because a log flush and a crash before the end of the reap step can
+result in space leaking.
+Online repair functions minimize the chances of this occurring by using very
+large transactions, which each can accommodate many thousands of block freeing
+instructions.
+Repair moves on to reaping the old blocks, which will be presented in a
+subsequent :ref:`section<reaping>` after a few case studies of bulk loading.
+
+Case Study: Rebuilding the Inode Index
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+The high level process to rebuild the inode index btree is:
+
+1. Walk the reverse mapping records to generate ``struct xfs_inobt_rec``
+   records from the inode chunk information and a bitmap of the old inode btree
+   blocks.
+
+2. Append the records to an xfarray in inode order.
+
+3. Use the ``xfs_btree_bload_compute_geometry`` function to compute the number
+   of blocks needed for the inode btree.
+   If the free space inode btree is enabled, call it again to estimate the
+   geometry of the finobt.
+
+4. Allocate the number of blocks computed in the previous step.
+
+5. Use ``xfs_btree_bload`` to write the xfarray records to btree blocks and
+   generate the internal node blocks.
+   If the free space inode btree is enabled, call it again to load the finobt.
+
+6. Commit the location of the new btree root block(s) to the AGI.
+
+7. Reap the old btree blocks using the bitmap created in step 1.
+
+Details are as follows.
+
+The inode btree maps inumbers to the ondisk location of the associated
+inode records, which means that the inode btrees can be rebuilt from the
+reverse mapping information.
+Reverse mapping records with an owner of ``XFS_RMAP_OWN_INOBT`` marks the
+location of the old inode btree blocks.
+Each reverse mapping record with an owner of ``XFS_RMAP_OWN_INODES`` marks the
+location of at least one inode cluster buffer.
+A cluster is the smallest number of ondisk inodes that can be allocated or
+freed in a single transaction; it is never smaller than 1 fs block or 4 inodes.
+
+For the space represented by each inode cluster, ensure that there are no
+records in the free space btrees nor any records in the reference count btree.
+If there are, the space metadata inconsistencies are reason enough to abort the
+operation.
+Otherwise, read each cluster buffer to check that its contents appear to be
+ondisk inodes and to decide if the file is allocated
+(``xfs_dinode.i_mode != 0``) or free (``xfs_dinode.i_mode == 0``).
+Accumulate the results of successive inode cluster buffer reads until there is
+enough information to fill a single inode chunk record, which is 64 consecutive
+numbers in the inumber keyspace.
+If the chunk is sparse, the chunk record may include holes.
+
+Once the repair function accumulates one chunk's worth of data, it calls
+``xfarray_append`` to add the inode btree record to the xfarray.
+This xfarray is walked twice during the btree creation step -- once to populate
+the inode btree with all inode chunk records, and a second time to populate the
+free inode btree with records for chunks that have free non-sparse inodes.
+The number of records for the inode btree is the number of xfarray records,
+but the record count for the free inode btree has to be computed as inode chunk
+records are stored in the xfarray.
+
+The proposed patchset is the
+`AG btree repair
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-ag-btrees>`_
+series.
+
+Case Study: Rebuilding the Space Reference Counts
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+Reverse mapping records are used to rebuild the reference count information.
+Reference counts are required for correct operation of copy on write for shared
+file data.
+Imagine the reverse mapping entries as rectangles representing extents of
+physical blocks, and that the rectangles can be laid down to allow them to
+overlap each other.
+From the diagram below, it is apparent that a reference count record must start
+or end wherever the height of the stack changes.
+In other words, the record emission stimulus is level-triggered::
+
+                        █    ███
+              ██      █████ ████   ███        ██████
+        ██   ████     ███████████ ████     █████████
+        ████████████████████████████████ ███████████
+        ^ ^  ^^ ^^    ^ ^^ ^^^  ^^^^  ^ ^^ ^  ^     ^
+        2 1  23 21    3 43 234  2123  1 01 2  3     0
+
+The ondisk reference count btree does not store the refcount == 0 cases because
+the free space btree already records which blocks are free.
+Extents being used to stage copy-on-write operations should be the only records
+with refcount == 1.
+Single-owner file blocks aren't recorded in either the free space or the
+reference count btrees.
+
+The high level process to rebuild the reference count btree is:
+
+1. Walk the reverse mapping records to generate ``struct xfs_refcount_irec``
+   records for any space having more than one reverse mapping and add them to
+   the xfarray.
+   Any records owned by ``XFS_RMAP_OWN_COW`` are also added to the xfarray
+   because these are extents allocated to stage a copy on write operation and
+   are tracked in the refcount btree.
+
+   Use any records owned by ``XFS_RMAP_OWN_REFC`` to create a bitmap of old
+   refcount btree blocks.
+
+2. Sort the records in physical extent order, putting the CoW staging extents
+   at the end of the xfarray.
+   This matches the sorting order of records in the refcount btree.
+
+3. Use the ``xfs_btree_bload_compute_geometry`` function to compute the number
+   of blocks needed for the new tree.
+
+4. Allocate the number of blocks computed in the previous step.
+
+5. Use ``xfs_btree_bload`` to write the xfarray records to btree blocks and
+   generate the internal node blocks.
+
+6. Commit the location of new btree root block to the AGF.
+
+7. Reap the old btree blocks using the bitmap created in step 1.
+
+Details are as follows; the same algorithm is used by ``xfs_repair`` to
+generate refcount information from reverse mapping records.
+
+- Until the reverse mapping btree runs out of records:
+
+  - Retrieve the next record from the btree and put it in a bag.
+
+  - Collect all records with the same starting block from the btree and put
+    them in the bag.
+
+  - While the bag isn't empty:
+
+    - Among the mappings in the bag, compute the lowest block number where the
+      reference count changes.
+      This position will be either the starting block number of the next
+      unprocessed reverse mapping or the next block after the shortest mapping
+      in the bag.
+
+    - Remove all mappings from the bag that end at this position.
+
+    - Collect all reverse mappings that start at this position from the btree
+      and put them in the bag.
+
+    - If the size of the bag changed and is greater than one, create a new
+      refcount record associating the block number range that we just walked to
+      the size of the bag.
+
+The bag-like structure in this case is a type 2 xfarray as discussed in the
+:ref:`xfarray access patterns<xfarray_access_patterns>` section.
+Reverse mappings are added to the bag using ``xfarray_store_anywhere`` and
+removed via ``xfarray_unset``.
+Bag members are examined through ``xfarray_iter`` loops.
+
+The proposed patchset is the
+`AG btree repair
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-ag-btrees>`_
+series.
+
+Case Study: Rebuilding File Fork Mapping Indices
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+The high level process to rebuild a data/attr fork mapping btree is:
+
+1. Walk the reverse mapping records to generate ``struct xfs_bmbt_rec``
+   records from the reverse mapping records for that inode and fork.
+   Append these records to an xfarray.
+   Compute the bitmap of the old bmap btree blocks from the ``BMBT_BLOCK``
+   records.
+
+2. Use the ``xfs_btree_bload_compute_geometry`` function to compute the number
+   of blocks needed for the new tree.
+
+3. Sort the records in file offset order.
+
+4. If the extent records would fit in the inode fork immediate area, commit the
+   records to that immediate area and skip to step 8.
+
+5. Allocate the number of blocks computed in the previous step.
+
+6. Use ``xfs_btree_bload`` to write the xfarray records to btree blocks and
+   generate the internal node blocks.
+
+7. Commit the new btree root block to the inode fork immediate area.
+
+8. Reap the old btree blocks using the bitmap created in step 1.
+
+There are some complications here:
+First, it's possible to move the fork offset to adjust the sizes of the
+immediate areas if the data and attr forks are not both in BMBT format.
+Second, if there are sufficiently few fork mappings, it may be possible to use
+EXTENTS format instead of BMBT, which may require a conversion.
+Third, the incore extent map must be reloaded carefully to avoid disturbing
+any delayed allocation extents.
+
+The proposed patchset is the
+`file mapping repair
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-file-mappings>`_
+series.
+
+.. _reaping:
+
+Reaping Old Metadata Blocks
+---------------------------
+
+Whenever online fsck builds a new data structure to replace one that is
+suspect, there is a question of how to find and dispose of the blocks that
+belonged to the old structure.
+The laziest method of course is not to deal with them at all, but this slowly
+leads to service degradations as space leaks out of the filesystem.
+Hopefully, someone will schedule a rebuild of the free space information to
+plug all those leaks.
+Offline repair rebuilds all space metadata after recording the usage of
+the files and directories that it decides not to clear, hence it can build new
+structures in the discovered free space and avoid the question of reaping.
+
+As part of a repair, online fsck relies heavily on the reverse mapping records
+to find space that is owned by the corresponding rmap owner yet truly free.
+Cross referencing rmap records with other rmap records is necessary because
+there may be other data structures that also think they own some of those
+blocks (e.g. crosslinked trees).
+Permitting the block allocator to hand them out again will not push the system
+towards consistency.
+
+For space metadata, the process of finding extents to dispose of generally
+follows this format:
+
+1. Create a bitmap of space used by data structures that must be preserved.
+   The space reservations used to create the new metadata can be used here if
+   the same rmap owner code is used to denote all of the objects being rebuilt.
+
+2. Survey the reverse mapping data to create a bitmap of space owned by the
+   same ``XFS_RMAP_OWN_*`` number for the metadata that is being preserved.
+
+3. Use the bitmap disunion operator to subtract (1) from (2).
+   The remaining set bits represent candidate extents that could be freed.
+   The process moves on to step 4 below.
+
+Repairs for file-based metadata such as extended attributes, directories,
+symbolic links, quota files and realtime bitmaps are performed by building a
+new structure attached to a temporary file and swapping the forks.
+Afterward, the mappings in the old file fork are the candidate blocks for
+disposal.
+
+The process for disposing of old extents is as follows:
+
+4. For each candidate extent, count the number of reverse mapping records for
+   the first block in that extent that do not have the same rmap owner for the
+   data structure being repaired.
+
+   - If zero, the block has a single owner and can be freed.
+
+   - If not, the block is part of a crosslinked structure and must not be
+     freed.
+
+5. Starting with the next block in the extent, figure out how many more blocks
+   have the same zero/nonzero other owner status as that first block.
+
+6. If the region is crosslinked, delete the reverse mapping entry for the
+   structure being repaired and move on to the next region.
+
+7. If the region is to be freed, mark any corresponding buffers in the buffer
+   cache as stale to prevent log writeback.
+
+8. Free the region and move on.
+
+However, there is one complication to this procedure.
+Transactions are of finite size, so the reaping process must be careful to roll
+the transactions to avoid overruns.
+Overruns come from two sources:
+
+a. EFIs logged on behalf of space that is no longer occupied
+
+b. Log items for buffer invalidations
+
+This is also a window in which a crash during the reaping process can leak
+blocks.
+As stated earlier, online repair functions use very large transactions to
+minimize the chances of this occurring.
+
+The proposed patchset is the
+`preparation for bulk loading btrees
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-prep-for-bulk-loading>`_
+series.
+
+Case Study: Reaping After a Regular Btree Repair
+````````````````````````````````````````````````
+
+Old reference count and inode btrees are the easiest to reap because they have
+rmap records with special owner codes: ``XFS_RMAP_OWN_REFC`` for the refcount
+btree, and ``XFS_RMAP_OWN_INOBT`` for the inode and free inode btrees.
+Creating a list of extents to reap the old btree blocks is quite simple,
+conceptually:
+
+1. Lock the relevant AGI/AGF header buffers to prevent allocation and frees.
+
+2. For each reverse mapping record with an rmap owner corresponding to the
+   metadata structure being rebuilt, set the corresponding range in a bitmap.
+
+3. Walk the current data structures that have the same rmap owner.
+   For each block visited, clear that range in the above bitmap.
+
+4. Each set bit in the bitmap represents a block that could be a block from the
+   old data structures and hence is a candidate for reaping.
+   In other words, ``(rmap_records_owned_by & ~blocks_reachable_by_walk)``
+   are the blocks that might be freeable.
+
+If it is possible to maintain the AGF lock throughout the repair (which is the
+common case), then step 2 can be performed at the same time as the reverse
+mapping record walk that creates the records for the new btree.
+
+Case Study: Rebuilding the Free Space Indices
+`````````````````````````````````````````````
+
+The high level process to rebuild the free space indices is:
+
+1. Walk the reverse mapping records to generate ``struct xfs_alloc_rec_incore``
+   records from the gaps in the reverse mapping btree.
+
+2. Append the records to an xfarray.
+
+3. Use the ``xfs_btree_bload_compute_geometry`` function to compute the number
+   of blocks needed for each new tree.
+
+4. Allocate the number of blocks computed in the previous step from the free
+   space information collected.
+
+5. Use ``xfs_btree_bload`` to write the xfarray records to btree blocks and
+   generate the internal node blocks for the free space by length index.
+   Call it again for the free space by block number index.
+
+6. Commit the locations of the new btree root blocks to the AGF.
+
+7. Reap the old btree blocks by looking for space that is not recorded by the
+   reverse mapping btree, the new free space btrees, or the AGFL.
+
+Repairing the free space btrees has three key complications over a regular
+btree repair:
+
+First, free space is not explicitly tracked in the reverse mapping records.
+Hence, the new free space records must be inferred from gaps in the physical
+space component of the keyspace of the reverse mapping btree.
+
+Second, free space repairs cannot use the common btree reservation code because
+new blocks are reserved out of the free space btrees.
+This is impossible when repairing the free space btrees themselves.
+However, repair holds the AGF buffer lock for the duration of the free space
+index reconstruction, so it can use the collected free space information to
+supply the blocks for the new free space btrees.
+It is not necessary to back each reserved extent with an EFI because the new
+free space btrees are constructed in what the ondisk filesystem thinks is
+unowned space.
+However, if reserving blocks for the new btrees from the collected free space
+information changes the number of free space records, repair must re-estimate
+the new free space btree geometry with the new record count until the
+reservation is sufficient.
+As part of committing the new btrees, repair must ensure that reverse mappings
+are created for the reserved blocks and that unused reserved blocks are
+inserted into the free space btrees.
+Deferrred rmap and freeing operations are used to ensure that this transition
+is atomic, similar to the other btree repair functions.
+
+Third, finding the blocks to reap after the repair is not overly
+straightforward.
+Blocks for the free space btrees and the reverse mapping btrees are supplied by
+the AGFL.
+Blocks put onto the AGFL have reverse mapping records with the owner
+``XFS_RMAP_OWN_AG``.
+This ownership is retained when blocks move from the AGFL into the free space
+btrees or the reverse mapping btrees.
+When repair walks reverse mapping records to synthesize free space records, it
+creates a bitmap (``ag_owner_bitmap``) of all the space claimed by
+``XFS_RMAP_OWN_AG`` records.
+The repair context maintains a second bitmap corresponding to the rmap btree
+blocks and the AGFL blocks (``rmap_agfl_bitmap``).
+When the walk is complete, the bitmap disunion operation ``(ag_owner_bitmap &
+~rmap_agfl_bitmap)`` computes the extents that are used by the old free space
+btrees.
+These blocks can then be reaped using the methods outlined above.
+
+The proposed patchset is the
+`AG btree repair
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-ag-btrees>`_
+series.
+
+.. _rmap_reap:
+
+Case Study: Reaping After Repairing Reverse Mapping Btrees
+``````````````````````````````````````````````````````````
+
+Old reverse mapping btrees are less difficult to reap after a repair.
+As mentioned in the previous section, blocks on the AGFL, the two free space
+btree blocks, and the reverse mapping btree blocks all have reverse mapping
+records with ``XFS_RMAP_OWN_AG`` as the owner.
+The full process of gathering reverse mapping records and building a new btree
+are described in the case study of
+:ref:`live rebuilds of rmap data <rmap_repair>`, but a crucial point from that
+discussion is that the new rmap btree will not contain any records for the old
+rmap btree, nor will the old btree blocks be tracked in the free space btrees.
+The list of candidate reaping blocks is computed by setting the bits
+corresponding to the gaps in the new rmap btree records, and then clearing the
+bits corresponding to extents in the free space btrees and the current AGFL
+blocks.
+The result ``(new_rmapbt_gaps & ~(agfl | bnobt_records))`` are reaped using the
+methods outlined above.
+
+The rest of the process of rebuildng the reverse mapping btree is discussed
+in a separate :ref:`case study<rmap_repair>`.
+
+The proposed patchset is the
+`AG btree repair
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-ag-btrees>`_
+series.
+
+Case Study: Rebuilding the AGFL
+```````````````````````````````
+
+The allocation group free block list (AGFL) is repaired as follows:
+
+1. Create a bitmap for all the space that the reverse mapping data claims is
+   owned by ``XFS_RMAP_OWN_AG``.
+
+2. Subtract the space used by the two free space btrees and the rmap btree.
+
+3. Subtract any space that the reverse mapping data claims is owned by any
+   other owner, to avoid re-adding crosslinked blocks to the AGFL.
+
+4. Once the AGFL is full, reap any blocks leftover.
+
+5. The next operation to fix the freelist will right-size the list.
+
+See `fs/xfs/scrub/agheader_repair.c <https://git.kernel.org/pub/scm/linux/kernel/git/torvalds/linux.git/tree/fs/xfs/scrub/agheader_repair.c>`_ for more details.
+
+Inode Record Repairs
+--------------------
+
+Inode records must be handled carefully, because they have both ondisk records
+("dinodes") and an in-memory ("cached") representation.
+There is a very high potential for cache coherency issues if online fsck is not
+careful to access the ondisk metadata *only* when the ondisk metadata is so
+badly damaged that the filesystem cannot load the in-memory representation.
+When online fsck wants to open a damaged file for scrubbing, it must use
+specialized resource acquisition functions that return either the in-memory
+representation *or* a lock on whichever object is necessary to prevent any
+update to the ondisk location.
+
+The only repairs that should be made to the ondisk inode buffers are whatever
+is necessary to get the in-core structure loaded.
+This means fixing whatever is caught by the inode cluster buffer and inode fork
+verifiers, and retrying the ``iget`` operation.
+If the second ``iget`` fails, the repair has failed.
+
+Once the in-memory representation is loaded, repair can lock the inode and can
+subject it to comprehensive checks, repairs, and optimizations.
+Most inode attributes are easy to check and constrain, or are user-controlled
+arbitrary bit patterns; these are both easy to fix.
+Dealing with the data and attr fork extent counts and the file block counts is
+more complicated, because computing the correct value requires traversing the
+forks, or if that fails, leaving the fields invalid and waiting for the fork
+fsck functions to run.
+
+The proposed patchset is the
+`inode
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-inodes>`_
+repair series.
+
+Quota Record Repairs
+--------------------
+
+Similar to inodes, quota records ("dquots") also have both ondisk records and
+an in-memory representation, and hence are subject to the same cache coherency
+issues.
+Somewhat confusingly, both are known as dquots in the XFS codebase.
+
+The only repairs that should be made to the ondisk quota record buffers are
+whatever is necessary to get the in-core structure loaded.
+Once the in-memory representation is loaded, the only attributes needing
+checking are obviously bad limits and timer values.
+
+Quota usage counters are checked, repaired, and discussed separately in the
+section about :ref:`live quotacheck <quotacheck>`.
+
+The proposed patchset is the
+`quota
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-quota>`_
+repair series.
+
+.. _fscounters:
+
+Freezing to Fix Summary Counters
+--------------------------------
+
+Filesystem summary counters track availability of filesystem resources such
+as free blocks, free inodes, and allocated inodes.
+This information could be compiled by walking the free space and inode indexes,
+but this is a slow process, so XFS maintains a copy in the ondisk superblock
+that should reflect the ondisk metadata, at least when the filesystem has been
+unmounted cleanly.
+For performance reasons, XFS also maintains incore copies of those counters,
+which are key to enabling resource reservations for active transactions.
+Writer threads reserve the worst-case quantities of resources from the
+incore counter and give back whatever they don't use at commit time.
+It is therefore only necessary to serialize on the superblock when the
+superblock is being committed to disk.
+
+The lazy superblock counter feature introduced in XFS v5 took this even further
+by training log recovery to recompute the summary counters from the AG headers,
+which eliminated the need for most transactions even to touch the superblock.
+The only time XFS commits the summary counters is at filesystem unmount.
+To reduce contention even further, the incore counter is implemented as a
+percpu counter, which means that each CPU is allocated a batch of blocks from a
+global incore counter and can satisfy small allocations from the local batch.
+
+The high-performance nature of the summary counters makes it difficult for
+online fsck to check them, since there is no way to quiesce a percpu counter
+while the system is running.
+Although online fsck can read the filesystem metadata to compute the correct
+values of the summary counters, there's no way to hold the value of a percpu
+counter stable, so it's quite possible that the counter will be out of date by
+the time the walk is complete.
+Earlier versions of online scrub would return to userspace with an incomplete
+scan flag, but this is not a satisfying outcome for a system administrator.
+For repairs, the in-memory counters must be stabilized while walking the
+filesystem metadata to get an accurate reading and install it in the percpu
+counter.
+
+To satisfy this requirement, online fsck must prevent other programs in the
+system from initiating new writes to the filesystem, it must disable background
+garbage collection threads, and it must wait for existing writer programs to
+exit the kernel.
+Once that has been established, scrub can walk the AG free space indexes, the
+inode btrees, and the realtime bitmap to compute the correct value of all
+four summary counters.
+This is very similar to a filesystem freeze, though not all of the pieces are
+necessary:
+
+- The final freeze state is set one higher than ``SB_FREEZE_COMPLETE`` to
+  prevent other threads from thawing the filesystem, or other scrub threads
+  from initiating another fscounters freeze.
+
+- It does not quiesce the log.
+
+With this code in place, it is now possible to pause the filesystem for just
+long enough to check and correct the summary counters.
+
++--------------------------------------------------------------------------+
+| **Historical Sidebar**:                                                  |
++--------------------------------------------------------------------------+
+| The initial implementation used the actual VFS filesystem freeze         |
+| mechanism to quiesce filesystem activity.                                |
+| With the filesystem frozen, it is possible to resolve the counter values |
+| with exact precision, but there are many problems with calling the VFS   |
+| methods directly:                                                        |
+|                                                                          |
+| - Other programs can unfreeze the filesystem without our knowledge.      |
+|   This leads to incorrect scan results and incorrect repairs.            |
+|                                                                          |
+| - Adding an extra lock to prevent others from thawing the filesystem     |
+|   required the addition of a ``->freeze_super`` function to wrap         |
+|   ``freeze_fs()``.                                                       |
+|   This in turn caused other subtle problems because it turns out that    |
+|   the VFS ``freeze_super`` and ``thaw_super`` functions can drop the     |
+|   last reference to the VFS superblock, and any subsequent access        |
+|   becomes a UAF bug!                                                     |
+|   This can happen if the filesystem is unmounted while the underlying    |
+|   block device has frozen the filesystem.                                |
+|   This problem could be solved by grabbing extra references to the       |
+|   superblock, but it felt suboptimal given the other inadequacies of     |
+|   this approach.                                                         |
+|                                                                          |
+| - The log need not be quiesced to check the summary counters, but a VFS  |
+|   freeze initiates one anyway.                                           |
+|   This adds unnecessary runtime to live fscounter fsck operations.       |
+|                                                                          |
+| - Quiescing the log means that XFS flushes the (possibly incorrect)      |
+|   counters to disk as part of cleaning the log.                          |
+|                                                                          |
+| - A bug in the VFS meant that freeze could complete even when            |
+|   sync_filesystem fails to flush the filesystem and returns an error.    |
+|   This bug was fixed in Linux 5.17.                                      |
++--------------------------------------------------------------------------+
+
+The proposed patchset is the
+`summary counter cleanup
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-fscounters>`_
+series.
+
+Full Filesystem Scans
+---------------------
+
+Certain types of metadata can only be checked by walking every file in the
+entire filesystem to record observations and comparing the observations against
+what's recorded on disk.
+Like every other type of online repair, repairs are made by writing those
+observations to disk in a replacement structure and committing it atomically.
+However, it is not practical to shut down the entire filesystem to examine
+hundreds of billions of files because the downtime would be excessive.
+Therefore, online fsck must build the infrastructure to manage a live scan of
+all the files in the filesystem.
+There are two questions that need to be solved to perform a live walk:
+
+- How does scrub manage the scan while it is collecting data?
+
+- How does the scan keep abreast of changes being made to the system by other
+  threads?
+
+.. _iscan:
+
+Coordinated Inode Scans
+```````````````````````
+
+In the original Unix filesystems of the 1970s, each directory entry contained
+an index number (*inumber*) which was used as an index into on ondisk array
+(*itable*) of fixed-size records (*inodes*) describing a file's attributes and
+its data block mapping.
+This system is described by J. Lions, `"inode (5659)"
+<http://www.lemis.com/grog/Documentation/Lions/>`_ in *Lions' Commentary on
+UNIX, 6th Edition*, (Dept. of Computer Science, the University of New South
+Wales, November 1977), pp. 18-2; and later by D. Ritchie and K. Thompson,
+`"Implementation of the File System"
+<https://archive.org/details/bstj57-6-1905/page/n8/mode/1up>`_, from *The UNIX
+Time-Sharing System*, (The Bell System Technical Journal, July 1978), pp.
+1913-4.
+
+XFS retains most of this design, except now inumbers are search keys over all
+the space in the data section filesystem.
+They form a continuous keyspace that can be expressed as a 64-bit integer,
+though the inodes themselves are sparsely distributed within the keyspace.
+Scans proceed in a linear fashion across the inumber keyspace, starting from
+``0x0`` and ending at ``0xFFFFFFFFFFFFFFFF``.
+Naturally, a scan through a keyspace requires a scan cursor object to track the
+scan progress.
+Because this keyspace is sparse, this cursor contains two parts.
+The first part of this scan cursor object tracks the inode that will be
+examined next; call this the examination cursor.
+Somewhat less obviously, the scan cursor object must also track which parts of
+the keyspace have already been visited, which is critical for deciding if a
+concurrent filesystem update needs to be incorporated into the scan data.
+Call this the visited inode cursor.
+
+Advancing the scan cursor is a multi-step process encapsulated in
+``xchk_iscan_iter``:
+
+1. Lock the AGI buffer of the AG containing the inode pointed to by the visited
+   inode cursor.
+   This guarantee that inodes in this AG cannot be allocated or freed while
+   advancing the cursor.
+
+2. Use the per-AG inode btree to look up the next inumber after the one that
+   was just visited, since it may not be keyspace adjacent.
+
+3. If there are no more inodes left in this AG:
+
+   a. Move the examination cursor to the point of the inumber keyspace that
+      corresponds to the start of the next AG.
+
+   b. Adjust the visited inode cursor to indicate that it has "visited" the
+      last possible inode in the current AG's inode keyspace.
+      XFS inumbers are segmented, so the cursor needs to be marked as having
+      visited the entire keyspace up to just before the start of the next AG's
+      inode keyspace.
+
+   c. Unlock the AGI and return to step 1 if there are unexamined AGs in the
+      filesystem.
+
+   d. If there are no more AGs to examine, set both cursors to the end of the
+      inumber keyspace.
+      The scan is now complete.
+
+4. Otherwise, there is at least one more inode to scan in this AG:
+
+   a. Move the examination cursor ahead to the next inode marked as allocated
+      by the inode btree.
+
+   b. Adjust the visited inode cursor to point to the inode just prior to where
+      the examination cursor is now.
+      Because the scanner holds the AGI buffer lock, no inodes could have been
+      created in the part of the inode keyspace that the visited inode cursor
+      just advanced.
+
+5. Get the incore inode for the inumber of the examination cursor.
+   By maintaining the AGI buffer lock until this point, the scanner knows that
+   it was safe to advance the examination cursor across the entire keyspace,
+   and that it has stabilized this next inode so that it cannot disappear from
+   the filesystem until the scan releases the incore inode.
+
+6. Drop the AGI lock and return the incore inode to the caller.
+
+Online fsck functions scan all files in the filesystem as follows:
+
+1. Start a scan by calling ``xchk_iscan_start``.
+
+2. Advance the scan cursor (``xchk_iscan_iter``) to get the next inode.
+   If one is provided:
+
+   a. Lock the inode to prevent updates during the scan.
+
+   b. Scan the inode.
+
+   c. While still holding the inode lock, adjust the visited inode cursor
+      (``xchk_iscan_mark_visited``) to point to this inode.
+
+   d. Unlock and release the inode.
+
+8. Call ``xchk_iscan_teardown`` to complete the scan.
+
+There are subtleties with the inode cache that complicate grabbing the incore
+inode for the caller.
+Obviously, it is an absolute requirement that the inode metadata be consistent
+enough to load it into the inode cache.
+Second, if the incore inode is stuck in some intermediate state, the scan
+coordinator must release the AGI and push the main filesystem to get the inode
+back into a loadable state.
+
+The proposed patches are the
+`inode scanner
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=scrub-iscan>`_
+series.
+The first user of the new functionality is the
+`online quotacheck
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-quotacheck>`_
+series.
+
+Inode Management
+````````````````
+
+In regular filesystem code, references to allocated XFS incore inodes are
+always obtained (``xfs_iget``) outside of transaction context because the
+creation of the incore context for an existing file does not require metadata
+updates.
+However, it is important to note that references to incore inodes obtained as
+part of file creation must be performed in transaction context because the
+filesystem must ensure the atomicity of the ondisk inode btree index updates
+and the initialization of the actual ondisk inode.
+
+References to incore inodes are always released (``xfs_irele``) outside of
+transaction context because there are a handful of activities that might
+require ondisk updates:
+
+- The VFS may decide to kick off writeback as part of a ``DONTCACHE`` inode
+  release.
+
+- Speculative preallocations need to be unreserved.
+
+- An unlinked file may have lost its last reference, in which case the entire
+  file must be inactivated, which involves releasing all of its resources in
+  the ondisk metadata and freeing the inode.
+
+These activities are collectively called inode inactivation.
+Inactivation has two parts -- the VFS part, which initiates writeback on all
+dirty file pages, and the XFS part, which cleans up XFS-specific information
+and frees the inode if it was unlinked.
+If the inode is unlinked (or unconnected after a file handle operation), the
+kernel drops the inode into the inactivation machinery immediately.
+
+During normal operation, resource acquisition for an update follows this order
+to avoid deadlocks:
+
+1. Inode reference (``iget``).
+
+2. Filesystem freeze protection, if repairing (``mnt_want_write_file``).
+
+3. Inode ``IOLOCK`` (VFS ``i_rwsem``) lock to control file IO.
+
+4. Inode ``MMAPLOCK`` (page cache ``invalidate_lock``) lock for operations that
+   can update page cache mappings.
+
+5. Log feature enablement.
+
+6. Transaction log space grant.
+
+7. Space on the data and realtime devices for the transaction.
+
+8. Incore dquot references, if a file is being repaired.
+   Note that they are not locked, merely acquired.
+
+9. Inode ``ILOCK`` for file metadata updates.
+
+10. AG header buffer locks / Realtime metadata inode ILOCK.
+
+11. Realtime metadata buffer locks, if applicable.
+
+12. Extent mapping btree blocks, if applicable.
+
+Resources are often released in the reverse order, though this is not required.
+However, online fsck differs from regular XFS operations because it may examine
+an object that normally is acquired in a later stage of the locking order, and
+then decide to cross-reference the object with an object that is acquired
+earlier in the order.
+The next few sections detail the specific ways in which online fsck takes care
+to avoid deadlocks.
+
+iget and irele During a Scrub
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+An inode scan performed on behalf of a scrub operation runs in transaction
+context, and possibly with resources already locked and bound to it.
+This isn't much of a problem for ``iget`` since it can operate in the context
+of an existing transaction, as long as all of the bound resources are acquired
+before the inode reference in the regular filesystem.
+
+When the VFS ``iput`` function is given a linked inode with no other
+references, it normally puts the inode on an LRU list in the hope that it can
+save time if another process re-opens the file before the system runs out
+of memory and frees it.
+Filesystem callers can short-circuit the LRU process by setting a ``DONTCACHE``
+flag on the inode to cause the kernel to try to drop the inode into the
+inactivation machinery immediately.
+
+In the past, inactivation was always done from the process that dropped the
+inode, which was a problem for scrub because scrub may already hold a
+transaction, and XFS does not support nesting transactions.
+On the other hand, if there is no scrub transaction, it is desirable to drop
+otherwise unused inodes immediately to avoid polluting caches.
+To capture these nuances, the online fsck code has a separate ``xchk_irele``
+function to set or clear the ``DONTCACHE`` flag to get the required release
+behavior.
+
+Proposed patchsets include fixing
+`scrub iget usage
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=scrub-iget-fixes>`_ and
+`dir iget usage
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=scrub-dir-iget-fixes>`_.
+
+.. _ilocking:
+
+Locking Inodes
+^^^^^^^^^^^^^^
+
+In regular filesystem code, the VFS and XFS will acquire multiple IOLOCK locks
+in a well-known order: parent → child when updating the directory tree, and
+in numerical order of the addresses of their ``struct inode`` object otherwise.
+For regular files, the MMAPLOCK can be acquired after the IOLOCK to stop page
+faults.
+If two MMAPLOCKs must be acquired, they are acquired in numerical order of
+the addresses of their ``struct address_space`` objects.
+Due to the structure of existing filesystem code, IOLOCKs and MMAPLOCKs must be
+acquired before transactions are allocated.
+If two ILOCKs must be acquired, they are acquired in inumber order.
+
+Inode lock acquisition must be done carefully during a coordinated inode scan.
+Online fsck cannot abide these conventions, because for a directory tree
+scanner, the scrub process holds the IOLOCK of the file being scanned and it
+needs to take the IOLOCK of the file at the other end of the directory link.
+If the directory tree is corrupt because it contains a cycle, ``xfs_scrub``
+cannot use the regular inode locking functions and avoid becoming trapped in an
+ABBA deadlock.
+
+Solving both of these problems is straightforward -- any time online fsck
+needs to take a second lock of the same class, it uses trylock to avoid an ABBA
+deadlock.
+If the trylock fails, scrub drops all inode locks and use trylock loops to
+(re)acquire all necessary resources.
+Trylock loops enable scrub to check for pending fatal signals, which is how
+scrub avoids deadlocking the filesystem or becoming an unresponsive process.
+However, trylock loops means that online fsck must be prepared to measure the
+resource being scrubbed before and after the lock cycle to detect changes and
+react accordingly.
+
+.. _dirparent:
+
+Case Study: Finding a Directory Parent
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+Consider the directory parent pointer repair code as an example.
+Online fsck must verify that the dotdot dirent of a directory points up to a
+parent directory, and that the parent directory contains exactly one dirent
+pointing down to the child directory.
+Fully validating this relationship (and repairing it if possible) requires a
+walk of every directory on the filesystem while holding the child locked, and
+while updates to the directory tree are being made.
+The coordinated inode scan provides a way to walk the filesystem without the
+possibility of missing an inode.
+The child directory is kept locked to prevent updates to the dotdot dirent, but
+if the scanner fails to lock a parent, it can drop and relock both the child
+and the prospective parent.
+If the dotdot entry changes while the directory is unlocked, then a move or
+rename operation must have changed the child's parentage, and the scan can
+exit early.
+
+The proposed patchset is the
+`directory repair
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-dirs>`_
+series.
+
+.. _fshooks:
+
+Filesystem Hooks
+`````````````````
+
+The second piece of support that online fsck functions need during a full
+filesystem scan is the ability to stay informed about updates being made by
+other threads in the filesystem, since comparisons against the past are useless
+in a dynamic environment.
+Two pieces of Linux kernel infrastructure enable online fsck to monitor regular
+filesystem operations: filesystem hooks and :ref:`static keys<jump_labels>`.
+
+Filesystem hooks convey information about an ongoing filesystem operation to
+a downstream consumer.
+In this case, the downstream consumer is always an online fsck function.
+Because multiple fsck functions can run in parallel, online fsck uses the Linux
+notifier call chain facility to dispatch updates to any number of interested
+fsck processes.
+Call chains are a dynamic list, which means that they can be configured at
+run time.
+Because these hooks are private to the XFS module, the information passed along
+contains exactly what the checking function needs to update its observations.
+
+The current implementation of XFS hooks uses SRCU notifier chains to reduce the
+impact to highly threaded workloads.
+Regular blocking notifier chains use a rwsem and seem to have a much lower
+overhead for single-threaded applications.
+However, it may turn out that the combination of blocking chains and static
+keys are a more performant combination; more study is needed here.
+
+The following pieces are necessary to hook a certain point in the filesystem:
+
+- A ``struct xfs_hooks`` object must be embedded in a convenient place such as
+  a well-known incore filesystem object.
+
+- Each hook must define an action code and a structure containing more context
+  about the action.
+
+- Hook providers should provide appropriate wrapper functions and structs
+  around the ``xfs_hooks`` and ``xfs_hook`` objects to take advantage of type
+  checking to ensure correct usage.
+
+- A callsite in the regular filesystem code must be chosen to call
+  ``xfs_hooks_call`` with the action code and data structure.
+  This place should be adjacent to (and not earlier than) the place where
+  the filesystem update is committed to the transaction.
+  In general, when the filesystem calls a hook chain, it should be able to
+  handle sleeping and should not be vulnerable to memory reclaim or locking
+  recursion.
+  However, the exact requirements are very dependent on the context of the hook
+  caller and the callee.
+
+- The online fsck function should define a structure to hold scan data, a lock
+  to coordinate access to the scan data, and a ``struct xfs_hook`` object.
+  The scanner function and the regular filesystem code must acquire resources
+  in the same order; see the next section for details.
+
+- The online fsck code must contain a C function to catch the hook action code
+  and data structure.
+  If the object being updated has already been visited by the scan, then the
+  hook information must be applied to the scan data.
+
+- Prior to unlocking inodes to start the scan, online fsck must call
+  ``xfs_hooks_setup`` to initialize the ``struct xfs_hook``, and
+  ``xfs_hooks_add`` to enable the hook.
+
+- Online fsck must call ``xfs_hooks_del`` to disable the hook once the scan is
+  complete.
+
+The number of hooks should be kept to a minimum to reduce complexity.
+Static keys are used to reduce the overhead of filesystem hooks to nearly
+zero when online fsck is not running.
+
+.. _liveupdate:
+
+Live Updates During a Scan
+``````````````````````````
+
+The code paths of the online fsck scanning code and the :ref:`hooked<fshooks>`
+filesystem code look like this::
+
+            other program
+                  ↓
+            inode lock ←────────────────────┐
+                  ↓                         │
+            AG header lock                  │
+                  ↓                         │
+            filesystem function             │
+                  ↓                         │
+            notifier call chain             │    same
+                  ↓                         ├─── inode
+            scrub hook function             │    lock
+                  ↓                         │
+            scan data mutex ←──┐    same    │
+                  ↓            ├─── scan    │
+            update scan data   │    lock    │
+                  ↑            │            │
+            scan data mutex ←──┘            │
+                  ↑                         │
+            inode lock ←────────────────────┘
+                  ↑
+            scrub function
+                  ↑
+            inode scanner
+                  ↑
+            xfs_scrub
+
+These rules must be followed to ensure correct interactions between the
+checking code and the code making an update to the filesystem:
+
+- Prior to invoking the notifier call chain, the filesystem function being
+  hooked must acquire the same lock that the scrub scanning function acquires
+  to scan the inode.
+
+- The scanning function and the scrub hook function must coordinate access to
+  the scan data by acquiring a lock on the scan data.
+
+- Scrub hook function must not add the live update information to the scan
+  observations unless the inode being updated has already been scanned.
+  The scan coordinator has a helper predicate (``xchk_iscan_want_live_update``)
+  for this.
+
+- Scrub hook functions must not change the caller's state, including the
+  transaction that it is running.
+  They must not acquire any resources that might conflict with the filesystem
+  function being hooked.
+
+- The hook function can abort the inode scan to avoid breaking the other rules.
+
+The inode scan APIs are pretty simple:
+
+- ``xchk_iscan_start`` starts a scan
+
+- ``xchk_iscan_iter`` grabs a reference to the next inode in the scan or
+  returns zero if there is nothing left to scan
+
+- ``xchk_iscan_want_live_update`` to decide if an inode has already been
+  visited in the scan.
+  This is critical for hook functions to decide if they need to update the
+  in-memory scan information.
+
+- ``xchk_iscan_mark_visited`` to mark an inode as having been visited in the
+  scan
+
+- ``xchk_iscan_teardown`` to finish the scan
+
+This functionality is also a part of the
+`inode scanner
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=scrub-iscan>`_
+series.
+
+.. _quotacheck:
+
+Case Study: Quota Counter Checking
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+It is useful to compare the mount time quotacheck code to the online repair
+quotacheck code.
+Mount time quotacheck does not have to contend with concurrent operations, so
+it does the following:
+
+1. Make sure the ondisk dquots are in good enough shape that all the incore
+   dquots will actually load, and zero the resource usage counters in the
+   ondisk buffer.
+
+2. Walk every inode in the filesystem.
+   Add each file's resource usage to the incore dquot.
+
+3. Walk each incore dquot.
+   If the incore dquot is not being flushed, add the ondisk buffer backing the
+   incore dquot to a delayed write (delwri) list.
+
+4. Write the buffer list to disk.
+
+Like most online fsck functions, online quotacheck can't write to regular
+filesystem objects until the newly collected metadata reflect all filesystem
+state.
+Therefore, online quotacheck records file resource usage to a shadow dquot
+index implemented with a sparse ``xfarray``, and only writes to the real dquots
+once the scan is complete.
+Handling transactional updates is tricky because quota resource usage updates
+are handled in phases to minimize contention on dquots:
+
+1. The inodes involved are joined and locked to a transaction.
+
+2. For each dquot attached to the file:
+
+   a. The dquot is locked.
+
+   b. A quota reservation is added to the dquot's resource usage.
+      The reservation is recorded in the transaction.
+
+   c. The dquot is unlocked.
+
+3. Changes in actual quota usage are tracked in the transaction.
+
+4. At transaction commit time, each dquot is examined again:
+
+   a. The dquot is locked again.
+
+   b. Quota usage changes are logged and unused reservation is given back to
+      the dquot.
+
+   c. The dquot is unlocked.
+
+For online quotacheck, hooks are placed in steps 2 and 4.
+The step 2 hook creates a shadow version of the transaction dquot context
+(``dqtrx``) that operates in a similar manner to the regular code.
+The step 4 hook commits the shadow ``dqtrx`` changes to the shadow dquots.
+Notice that both hooks are called with the inode locked, which is how the
+live update coordinates with the inode scanner.
+
+The quotacheck scan looks like this:
+
+1. Set up a coordinated inode scan.
+
+2. For each inode returned by the inode scan iterator:
+
+   a. Grab and lock the inode.
+
+   b. Determine that inode's resource usage (data blocks, inode counts,
+      realtime blocks) and add that to the shadow dquots for the user, group,
+      and project ids associated with the inode.
+
+   c. Unlock and release the inode.
+
+3. For each dquot in the system:
+
+   a. Grab and lock the dquot.
+
+   b. Check the dquot against the shadow dquots created by the scan and updated
+      by the live hooks.
+
+Live updates are key to being able to walk every quota record without
+needing to hold any locks for a long duration.
+If repairs are desired, the real and shadow dquots are locked and their
+resource counts are set to the values in the shadow dquot.
+
+The proposed patchset is the
+`online quotacheck
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-quotacheck>`_
+series.
+
+.. _nlinks:
+
+Case Study: File Link Count Checking
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+File link count checking also uses live update hooks.
+The coordinated inode scanner is used to visit all directories on the
+filesystem, and per-file link count records are stored in a sparse ``xfarray``
+indexed by inumber.
+During the scanning phase, each entry in a directory generates observation
+data as follows:
+
+1. If the entry is a dotdot (``'..'``) entry of the root directory, the
+   directory's parent link count is bumped because the root directory's dotdot
+   entry is self referential.
+
+2. If the entry is a dotdot entry of a subdirectory, the parent's backref
+   count is bumped.
+
+3. If the entry is neither a dot nor a dotdot entry, the target file's parent
+   count is bumped.
+
+4. If the target is a subdirectory, the parent's child link count is bumped.
+
+A crucial point to understand about how the link count inode scanner interacts
+with the live update hooks is that the scan cursor tracks which *parent*
+directories have been scanned.
+In other words, the live updates ignore any update about ``A → B`` when A has
+not been scanned, even if B has been scanned.
+Furthermore, a subdirectory A with a dotdot entry pointing back to B is
+accounted as a backref counter in the shadow data for A, since child dotdot
+entries affect the parent's link count.
+Live update hooks are carefully placed in all parts of the filesystem that
+create, change, or remove directory entries, since those operations involve
+bumplink and droplink.
+
+For any file, the correct link count is the number of parents plus the number
+of child subdirectories.
+Non-directories never have children of any kind.
+The backref information is used to detect inconsistencies in the number of
+links pointing to child subdirectories and the number of dotdot entries
+pointing back.
+
+After the scan completes, the link count of each file can be checked by locking
+both the inode and the shadow data, and comparing the link counts.
+A second coordinated inode scan cursor is used for comparisons.
+Live updates are key to being able to walk every inode without needing to hold
+any locks between inodes.
+If repairs are desired, the inode's link count is set to the value in the
+shadow information.
+If no parents are found, the file must be :ref:`reparented <orphanage>` to the
+orphanage to prevent the file from being lost forever.
+
+The proposed patchset is the
+`file link count repair
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=scrub-nlinks>`_
+series.
+
+.. _rmap_repair:
+
+Case Study: Rebuilding Reverse Mapping Records
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+Most repair functions follow the same pattern: lock filesystem resources,
+walk the surviving ondisk metadata looking for replacement metadata records,
+and use an :ref:`in-memory array <xfarray>` to store the gathered observations.
+The primary advantage of this approach is the simplicity and modularity of the
+repair code -- code and data are entirely contained within the scrub module,
+do not require hooks in the main filesystem, and are usually the most efficient
+in memory use.
+A secondary advantage of this repair approach is atomicity -- once the kernel
+decides a structure is corrupt, no other threads can access the metadata until
+the kernel finishes repairing and revalidating the metadata.
+
+For repairs going on within a shard of the filesystem, these advantages
+outweigh the delays inherent in locking the shard while repairing parts of the
+shard.
+Unfortunately, repairs to the reverse mapping btree cannot use the "standard"
+btree repair strategy because it must scan every space mapping of every fork of
+every file in the filesystem, and the filesystem cannot stop.
+Therefore, rmap repair foregoes atomicity between scrub and repair.
+It combines a :ref:`coordinated inode scanner <iscan>`, :ref:`live update hooks
+<liveupdate>`, and an :ref:`in-memory rmap btree <xfbtree>` to complete the
+scan for reverse mapping records.
+
+1. Set up an xfbtree to stage rmap records.
+
+2. While holding the locks on the AGI and AGF buffers acquired during the
+   scrub, generate reverse mappings for all AG metadata: inodes, btrees, CoW
+   staging extents, and the internal log.
+
+3. Set up an inode scanner.
+
+4. Hook into rmap updates for the AG being repaired so that the live scan data
+   can receive updates to the rmap btree from the rest of the filesystem during
+   the file scan.
+
+5. For each space mapping found in either fork of each file scanned,
+   decide if the mapping matches the AG of interest.
+   If so:
+
+   a. Create a btree cursor for the in-memory btree.
+
+   b. Use the rmap code to add the record to the in-memory btree.
+
+   c. Use the :ref:`special commit function <xfbtree_commit>` to write the
+      xfbtree changes to the xfile.
+
+6. For each live update received via the hook, decide if the owner has already
+   been scanned.
+   If so, apply the live update into the scan data:
+
+   a. Create a btree cursor for the in-memory btree.
+
+   b. Replay the operation into the in-memory btree.
+
+   c. Use the :ref:`special commit function <xfbtree_commit>` to write the
+      xfbtree changes to the xfile.
+      This is performed with an empty transaction to avoid changing the
+      caller's state.
+
+7. When the inode scan finishes, create a new scrub transaction and relock the
+   two AG headers.
+
+8. Compute the new btree geometry using the number of rmap records in the
+   shadow btree, like all other btree rebuilding functions.
+
+9. Allocate the number of blocks computed in the previous step.
+
+10. Perform the usual btree bulk loading and commit to install the new rmap
+    btree.
+
+11. Reap the old rmap btree blocks as discussed in the case study about how
+    to :ref:`reap after rmap btree repair <rmap_reap>`.
+
+12. Free the xfbtree now that it not needed.
+
+The proposed patchset is the
+`rmap repair
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-rmap-btree>`_
+series.
+
+Staging Repairs with Temporary Files on Disk
+--------------------------------------------
+
+XFS stores a substantial amount of metadata in file forks: directories,
+extended attributes, symbolic link targets, free space bitmaps and summary
+information for the realtime volume, and quota records.
+File forks map 64-bit logical file fork space extents to physical storage space
+extents, similar to how a memory management unit maps 64-bit virtual addresses
+to physical memory addresses.
+Therefore, file-based tree structures (such as directories and extended
+attributes) use blocks mapped in the file fork offset address space that point
+to other blocks mapped within that same address space, and file-based linear
+structures (such as bitmaps and quota records) compute array element offsets in
+the file fork offset address space.
+
+Because file forks can consume as much space as the entire filesystem, repairs
+cannot be staged in memory, even when a paging scheme is available.
+Therefore, online repair of file-based metadata createas a temporary file in
+the XFS filesystem, writes a new structure at the correct offsets into the
+temporary file, and atomically swaps the fork mappings (and hence the fork
+contents) to commit the repair.
+Once the repair is complete, the old fork can be reaped as necessary; if the
+system goes down during the reap, the iunlink code will delete the blocks
+during log recovery.
+
+**Note**: All space usage and inode indices in the filesystem *must* be
+consistent to use a temporary file safely!
+This dependency is the reason why online repair can only use pageable kernel
+memory to stage ondisk space usage information.
+
+Swapping metadata extents with a temporary file requires the owner field of the
+block headers to match the file being repaired and not the temporary file.  The
+directory, extended attribute, and symbolic link functions were all modified to
+allow callers to specify owner numbers explicitly.
+
+There is a downside to the reaping process -- if the system crashes during the
+reap phase and the fork extents are crosslinked, the iunlink processing will
+fail because freeing space will find the extra reverse mappings and abort.
+
+Temporary files created for repair are similar to ``O_TMPFILE`` files created
+by userspace.
+They are not linked into a directory and the entire file will be reaped when
+the last reference to the file is lost.
+The key differences are that these files must have no access permission outside
+the kernel at all, they must be specially marked to prevent them from being
+opened by handle, and they must never be linked into the directory tree.
+
++--------------------------------------------------------------------------+
+| **Historical Sidebar**:                                                  |
++--------------------------------------------------------------------------+
+| In the initial iteration of file metadata repair, the damaged metadata   |
+| blocks would be scanned for salvageable data; the extents in the file    |
+| fork would be reaped; and then a new structure would be built in its     |
+| place.                                                                   |
+| This strategy did not survive the introduction of the atomic repair      |
+| requirement expressed earlier in this document.                          |
+|                                                                          |
+| The second iteration explored building a second structure at a high      |
+| offset in the fork from the salvage data, reaping the old extents, and   |
+| using a ``COLLAPSE_RANGE`` operation to slide the new extents into       |
+| place.                                                                   |
+|                                                                          |
+| This had many drawbacks:                                                 |
+|                                                                          |
+| - Array structures are linearly addressed, and the regular filesystem    |
+|   codebase does not have the concept of a linear offset that could be    |
+|   applied to the record offset computation to build an alternate copy.   |
+|                                                                          |
+| - Extended attributes are allowed to use the entire attr fork offset     |
+|   address space.                                                         |
+|                                                                          |
+| - Even if repair could build an alternate copy of a data structure in a  |
+|   different part of the fork address space, the atomic repair commit     |
+|   requirement means that online repair would have to be able to perform  |
+|   a log assisted ``COLLAPSE_RANGE`` operation to ensure that the old     |
+|   structure was completely replaced.                                     |
+|                                                                          |
+| - A crash after construction of the secondary tree but before the range  |
+|   collapse would leave unreachable blocks in the file fork.              |
+|   This would likely confuse things further.                              |
+|                                                                          |
+| - Reaping blocks after a repair is not a simple operation, and           |
+|   initiating a reap operation from a restarted range collapse operation  |
+|   during log recovery is daunting.                                       |
+|                                                                          |
+| - Directory entry blocks and quota records record the file fork offset   |
+|   in the header area of each block.                                      |
+|   An atomic range collapse operation would have to rewrite this part of  |
+|   each block header.                                                     |
+|   Rewriting a single field in block headers is not a huge problem, but   |
+|   it's something to be aware of.                                         |
+|                                                                          |
+| - Each block in a directory or extended attributes btree index contains  |
+|   sibling and child block pointers.                                      |
+|   Were the atomic commit to use a range collapse operation, each block   |
+|   would have to be rewritten very carefully to preserve the graph        |
+|   structure.                                                             |
+|   Doing this as part of a range collapse means rewriting a large number  |
+|   of blocks repeatedly, which is not conducive to quick repairs.         |
+|                                                                          |
+| This lead to the introduction of temporary file staging.                 |
++--------------------------------------------------------------------------+
+
+Using a Temporary File
+``````````````````````
+
+Online repair code should use the ``xrep_tempfile_create`` function to create a
+temporary file inside the filesystem.
+This allocates an inode, marks the in-core inode private, and attaches it to
+the scrub context.
+These files are hidden from userspace, may not be added to the directory tree,
+and must be kept private.
+
+Temporary files only use two inode locks: the IOLOCK and the ILOCK.
+The MMAPLOCK is not needed here, because there must not be page faults from
+userspace for data fork blocks.
+The usage patterns of these two locks are the same as for any other XFS file --
+access to file data are controlled via the IOLOCK, and access to file metadata
+are controlled via the ILOCK.
+Locking helpers are provided so that the temporary file and its lock state can
+be cleaned up by the scrub context.
+To comply with the nested locking strategy laid out in the :ref:`inode
+locking<ilocking>` section, it is recommended that scrub functions use the
+xrep_tempfile_ilock*_nowait lock helpers.
+
+Data can be written to a temporary file by two means:
+
+1. ``xrep_tempfile_copyin`` can be used to set the contents of a regular
+   temporary file from an xfile.
+
+2. The regular directory, symbolic link, and extended attribute functions can
+   be used to write to the temporary file.
+
+Once a good copy of a data file has been constructed in a temporary file, it
+must be conveyed to the file being repaired, which is the topic of the next
+section.
+
+The proposed patches are in the
+`repair temporary files
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-tempfiles>`_
+series.
+
+Atomic Extent Swapping
+----------------------
+
+Once repair builds a temporary file with a new data structure written into
+it, it must commit the new changes into the existing file.
+It is not possible to swap the inumbers of two files, so instead the new
+metadata must replace the old.
+This suggests the need for the ability to swap extents, but the existing extent
+swapping code used by the file defragmenting tool ``xfs_fsr`` is not sufficient
+for online repair because:
+
+a. When the reverse-mapping btree is enabled, the swap code must keep the
+   reverse mapping information up to date with every exchange of mappings.
+   Therefore, it can only exchange one mapping per transaction, and each
+   transaction is independent.
+
+b. Reverse-mapping is critical for the operation of online fsck, so the old
+   defragmentation code (which swapped entire extent forks in a single
+   operation) is not useful here.
+
+c. Defragmentation is assumed to occur between two files with identical
+   contents.
+   For this use case, an incomplete exchange will not result in a user-visible
+   change in file contents, even if the operation is interrupted.
+
+d. Online repair needs to swap the contents of two files that are by definition
+   *not* identical.
+   For directory and xattr repairs, the user-visible contents might be the
+   same, but the contents of individual blocks may be very different.
+
+e. Old blocks in the file may be cross-linked with another structure and must
+   not reappear if the system goes down mid-repair.
+
+These problems are overcome by creating a new deferred operation and a new type
+of log intent item to track the progress of an operation to exchange two file
+ranges.
+The new deferred operation type chains together the same transactions used by
+the reverse-mapping extent swap code.
+The new log item records the progress of the exchange to ensure that once an
+exchange begins, it will always run to completion, even there are
+interruptions.
+The new ``XFS_SB_FEAT_INCOMPAT_LOG_ATOMIC_SWAP`` log-incompatible feature flag
+in the superblock protects these new log item records from being replayed on
+old kernels.
+
+The proposed patchset is the
+`atomic extent swap
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=atomic-file-updates>`_
+series.
+
++--------------------------------------------------------------------------+
+| **Sidebar: Using Log-Incompatible Feature Flags**                        |
++--------------------------------------------------------------------------+
+| Starting with XFS v5, the superblock contains a                          |
+| ``sb_features_log_incompat`` field to indicate that the log contains     |
+| records that might not readable by all kernels that could mount this     |
+| filesystem.                                                              |
+| In short, log incompat features protect the log contents against kernels |
+| that will not understand the contents.                                   |
+| Unlike the other superblock feature bits, log incompat bits are          |
+| ephemeral because an empty (clean) log does not need protection.         |
+| The log cleans itself after its contents have been committed into the    |
+| filesystem, either as part of an unmount or because the system is        |
+| otherwise idle.                                                          |
+| Because upper level code can be working on a transaction at the same     |
+| time that the log cleans itself, it is necessary for upper level code to |
+| communicate to the log when it is going to use a log incompatible        |
+| feature.                                                                 |
+|                                                                          |
+| The log coordinates access to incompatible features through the use of   |
+| one ``struct rw_semaphore`` for each feature.                            |
+| The log cleaning code tries to take this rwsem in exclusive mode to      |
+| clear the bit; if the lock attempt fails, the feature bit remains set.   |
+| Filesystem code signals its intention to use a log incompat feature in a |
+| transaction by calling ``xlog_use_incompat_feat``, which takes the rwsem |
+| in shared mode.                                                          |
+| The code supporting a log incompat feature should create wrapper         |
+| functions to obtain the log feature and call                             |
+| ``xfs_add_incompat_log_feature`` to set the feature bits in the primary  |
+| superblock.                                                              |
+| The superblock update is performed transactionally, so the wrapper to    |
+| obtain log assistance must be called just prior to the creation of the   |
+| transaction that uses the functionality.                                 |
+| For a file operation, this step must happen after taking the IOLOCK      |
+| and the MMAPLOCK, but before allocating the transaction.                 |
+| When the transaction is complete, the ``xlog_drop_incompat_feat``        |
+| function is called to release the feature.                               |
+| The feature bit will not be cleared from the superblock until the log    |
+| becomes clean.                                                           |
+|                                                                          |
+| Log-assisted extended attribute updates and atomic extent swaps both use |
+| log incompat features and provide convenience wrappers around the        |
+| functionality.                                                           |
++--------------------------------------------------------------------------+
+
+Mechanics of an Atomic Extent Swap
+``````````````````````````````````
+
+Swapping entire file forks is a complex task.
+The goal is to exchange all file fork mappings between two file fork offset
+ranges.
+There are likely to be many extent mappings in each fork, and the edges of
+the mappings aren't necessarily aligned.
+Furthermore, there may be other updates that need to happen after the swap,
+such as exchanging file sizes, inode flags, or conversion of fork data to local
+format.
+This is roughly the format of the new deferred extent swap work item:
+
+.. code-block:: c
+
+	struct xfs_swapext_intent {
+	    /* Inodes participating in the operation. */
+	    struct xfs_inode    *sxi_ip1;
+	    struct xfs_inode    *sxi_ip2;
+
+	    /* File offset range information. */
+	    xfs_fileoff_t       sxi_startoff1;
+	    xfs_fileoff_t       sxi_startoff2;
+	    xfs_filblks_t       sxi_blockcount;
+
+	    /* Set these file sizes after the operation, unless negative. */
+	    xfs_fsize_t         sxi_isize1;
+	    xfs_fsize_t         sxi_isize2;
+
+	    /* XFS_SWAP_EXT_* log operation flags */
+	    uint64_t            sxi_flags;
+	};
+
+The new log intent item contains enough information to track two logical fork
+offset ranges: ``(inode1, startoff1, blockcount)`` and ``(inode2, startoff2,
+blockcount)``.
+Each step of a swap operation exchanges the largest file range mapping possible
+from one file to the other.
+After each step in the swap operation, the two startoff fields are incremented
+and the blockcount field is decremented to reflect the progress made.
+The flags field captures behavioral parameters such as swapping the attr fork
+instead of the data fork and other work to be done after the extent swap.
+The two isize fields are used to swap the file size at the end of the operation
+if the file data fork is the target of the swap operation.
+
+When the extent swap is initiated, the sequence of operations is as follows:
+
+1. Create a deferred work item for the extent swap.
+   At the start, it should contain the entirety of the file ranges to be
+   swapped.
+
+2. Call ``xfs_defer_finish`` to process the exchange.
+   This is encapsulated in ``xrep_tempswap_contents`` for scrub operations.
+   This will log an extent swap intent item to the transaction for the deferred
+   extent swap work item.
+
+3. Until ``sxi_blockcount`` of the deferred extent swap work item is zero,
+
+   a. Read the block maps of both file ranges starting at ``sxi_startoff1`` and
+      ``sxi_startoff2``, respectively, and compute the longest extent that can
+      be swapped in a single step.
+      This is the minimum of the two ``br_blockcount`` s in the mappings.
+      Keep advancing through the file forks until at least one of the mappings
+      contains written blocks.
+      Mutual holes, unwritten extents, and extent mappings to the same physical
+      space are not exchanged.
+
+      For the next few steps, this document will refer to the mapping that came
+      from file 1 as "map1", and the mapping that came from file 2 as "map2".
+
+   b. Create a deferred block mapping update to unmap map1 from file 1.
+
+   c. Create a deferred block mapping update to unmap map2 from file 2.
+
+   d. Create a deferred block mapping update to map map1 into file 2.
+
+   e. Create a deferred block mapping update to map map2 into file 1.
+
+   f. Log the block, quota, and extent count updates for both files.
+
+   g. Extend the ondisk size of either file if necessary.
+
+   h. Log an extent swap done log item for the extent swap intent log item
+      that was read at the start of step 3.
+
+   i. Compute the amount of file range that has just been covered.
+      This quantity is ``(map1.br_startoff + map1.br_blockcount -
+      sxi_startoff1)``, because step 3a could have skipped holes.
+
+   j. Increase the starting offsets of ``sxi_startoff1`` and ``sxi_startoff2``
+      by the number of blocks computed in the previous step, and decrease
+      ``sxi_blockcount`` by the same quantity.
+      This advances the cursor.
+
+   k. Log a new extent swap intent log item reflecting the advanced state of
+      the work item.
+
+   l. Return the proper error code (EAGAIN) to the deferred operation manager
+      to inform it that there is more work to be done.
+      The operation manager completes the deferred work in steps 3b-3e before
+      moving back to the start of step 3.
+
+4. Perform any post-processing.
+   This will be discussed in more detail in subsequent sections.
+
+If the filesystem goes down in the middle of an operation, log recovery will
+find the most recent unfinished extent swap log intent item and restart from
+there.
+This is how extent swapping guarantees that an outside observer will either see
+the old broken structure or the new one, and never a mismash of both.
+
+Preparation for Extent Swapping
+```````````````````````````````
+
+There are a few things that need to be taken care of before initiating an
+atomic extent swap operation.
+First, regular files require the page cache to be flushed to disk before the
+operation begins, and directio writes to be quiesced.
+Like any filesystem operation, extent swapping must determine the maximum
+amount of disk space and quota that can be consumed on behalf of both files in
+the operation, and reserve that quantity of resources to avoid an unrecoverable
+out of space failure once it starts dirtying metadata.
+The preparation step scans the ranges of both files to estimate:
+
+- Data device blocks needed to handle the repeated updates to the fork
+  mappings.
+- Change in data and realtime block counts for both files.
+- Increase in quota usage for both files, if the two files do not share the
+  same set of quota ids.
+- The number of extent mappings that will be added to each file.
+- Whether or not there are partially written realtime extents.
+  User programs must never be able to access a realtime file extent that maps
+  to different extents on the realtime volume, which could happen if the
+  operation fails to run to completion.
+
+The need for precise estimation increases the run time of the swap operation,
+but it is very important to maintain correct accounting.
+The filesystem must not run completely out of free space, nor can the extent
+swap ever add more extent mappings to a fork than it can support.
+Regular users are required to abide the quota limits, though metadata repairs
+may exceed quota to resolve inconsistent metadata elsewhere.
+
+Special Features for Swapping Metadata File Extents
+```````````````````````````````````````````````````
+
+Extended attributes, symbolic links, and directories can set the fork format to
+"local" and treat the fork as a literal area for data storage.
+Metadata repairs must take extra steps to support these cases:
+
+- If both forks are in local format and the fork areas are large enough, the
+  swap is performed by copying the incore fork contents, logging both forks,
+  and committing.
+  The atomic extent swap mechanism is not necessary, since this can be done
+  with a single transaction.
+
+- If both forks map blocks, then the regular atomic extent swap is used.
+
+- Otherwise, only one fork is in local format.
+  The contents of the local format fork are converted to a block to perform the
+  swap.
+  The conversion to block format must be done in the same transaction that
+  logs the initial extent swap intent log item.
+  The regular atomic extent swap is used to exchange the mappings.
+  Special flags are set on the swap operation so that the transaction can be
+  rolled one more time to convert the second file's fork back to local format
+  so that the second file will be ready to go as soon as the ILOCK is dropped.
+
+Extended attributes and directories stamp the owning inode into every block,
+but the buffer verifiers do not actually check the inode number!
+Although there is no verification, it is still important to maintain
+referential integrity, so prior to performing the extent swap, online repair
+builds every block in the new data structure with the owner field of the file
+being repaired.
+
+After a successful swap operation, the repair operation must reap the old fork
+blocks by processing each fork mapping through the standard :ref:`file extent
+reaping <reaping>` mechanism that is done post-repair.
+If the filesystem should go down during the reap part of the repair, the
+iunlink processing at the end of recovery will free both the temporary file and
+whatever blocks were not reaped.
+However, this iunlink processing omits the cross-link detection of online
+repair, and is not completely foolproof.
+
+Swapping Temporary File Extents
+```````````````````````````````
+
+To repair a metadata file, online repair proceeds as follows:
+
+1. Create a temporary repair file.
+
+2. Use the staging data to write out new contents into the temporary repair
+   file.
+   The same fork must be written to as is being repaired.
+
+3. Commit the scrub transaction, since the swap estimation step must be
+   completed before transaction reservations are made.
+
+4. Call ``xrep_tempswap_trans_alloc`` to allocate a new scrub transaction with
+   the appropriate resource reservations, locks, and fill out a ``struct
+   xfs_swapext_req`` with the details of the swap operation.
+
+5. Call ``xrep_tempswap_contents`` to swap the contents.
+
+6. Commit the transaction to complete the repair.
+
+.. _rtsummary:
+
+Case Study: Repairing the Realtime Summary File
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+In the "realtime" section of an XFS filesystem, free space is tracked via a
+bitmap, similar to Unix FFS.
+Each bit in the bitmap represents one realtime extent, which is a multiple of
+the filesystem block size between 4KiB and 1GiB in size.
+The realtime summary file indexes the number of free extents of a given size to
+the offset of the block within the realtime free space bitmap where those free
+extents begin.
+In other words, the summary file helps the allocator find free extents by
+length, similar to what the free space by count (cntbt) btree does for the data
+section.
+
+The summary file itself is a flat file (with no block headers or checksums!)
+partitioned into ``log2(total rt extents)`` sections containing enough 32-bit
+counters to match the number of blocks in the rt bitmap.
+Each counter records the number of free extents that start in that bitmap block
+and can satisfy a power-of-two allocation request.
+
+To check the summary file against the bitmap:
+
+1. Take the ILOCK of both the realtime bitmap and summary files.
+
+2. For each free space extent recorded in the bitmap:
+
+   a. Compute the position in the summary file that contains a counter that
+      represents this free extent.
+
+   b. Read the counter from the xfile.
+
+   c. Increment it, and write it back to the xfile.
+
+3. Compare the contents of the xfile against the ondisk file.
+
+To repair the summary file, write the xfile contents into the temporary file
+and use atomic extent swap to commit the new contents.
+The temporary file is then reaped.
+
+The proposed patchset is the
+`realtime summary repair
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-rtsummary>`_
+series.
+
+Case Study: Salvaging Extended Attributes
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+In XFS, extended attributes are implemented as a namespaced name-value store.
+Values are limited in size to 64KiB, but there is no limit in the number of
+names.
+The attribute fork is unpartitioned, which means that the root of the attribute
+structure is always in logical block zero, but attribute leaf blocks, dabtree
+index blocks, and remote value blocks are intermixed.
+Attribute leaf blocks contain variable-sized records that associate
+user-provided names with the user-provided values.
+Values larger than a block are allocated separate extents and written there.
+If the leaf information expands beyond a single block, a directory/attribute
+btree (``dabtree``) is created to map hashes of attribute names to entries
+for fast lookup.
+
+Salvaging extended attributes is done as follows:
+
+1. Walk the attr fork mappings of the file being repaired to find the attribute
+   leaf blocks.
+   When one is found,
+
+   a. Walk the attr leaf block to find candidate keys.
+      When one is found,
+
+      1. Check the name for problems, and ignore the name if there are.
+
+      2. Retrieve the value.
+         If that succeeds, add the name and value to the staging xfarray and
+         xfblob.
+
+2. If the memory usage of the xfarray and xfblob exceed a certain amount of
+   memory or there are no more attr fork blocks to examine, unlock the file and
+   add the staged extended attributes to the temporary file.
+
+3. Use atomic extent swapping to exchange the new and old extended attribute
+   structures.
+   The old attribute blocks are now attached to the temporary file.
+
+4. Reap the temporary file.
+
+The proposed patchset is the
+`extended attribute repair
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-xattrs>`_
+series.
+
+Fixing Directories
+------------------
+
+Fixing directories is difficult with currently available filesystem features,
+since directory entries are not redundant.
+The offline repair tool scans all inodes to find files with nonzero link count,
+and then it scans all directories to establish parentage of those linked files.
+Damaged files and directories are zapped, and files with no parent are
+moved to the ``/lost+found`` directory.
+It does not try to salvage anything.
+
+The best that online repair can do at this time is to read directory data
+blocks and salvage any dirents that look plausible, correct link counts, and
+move orphans back into the directory tree.
+The salvage process is discussed in the case study at the end of this section.
+The :ref:`file link count fsck <nlinks>` code takes care of fixing link counts
+and moving orphans to the ``/lost+found`` directory.
+
+Case Study: Salvaging Directories
+`````````````````````````````````
+
+Unlike extended attributes, directory blocks are all the same size, so
+salvaging directories is straightforward:
+
+1. Find the parent of the directory.
+   If the dotdot entry is not unreadable, try to confirm that the alleged
+   parent has a child entry pointing back to the directory being repaired.
+   Otherwise, walk the filesystem to find it.
+
+2. Walk the first partition of data fork of the directory to find the directory
+   entry data blocks.
+   When one is found,
+
+   a. Walk the directory data block to find candidate entries.
+      When an entry is found:
+
+      i. Check the name for problems, and ignore the name if there are.
+
+      ii. Retrieve the inumber and grab the inode.
+          If that succeeds, add the name, inode number, and file type to the
+          staging xfarray and xblob.
+
+3. If the memory usage of the xfarray and xfblob exceed a certain amount of
+   memory or there are no more directory data blocks to examine, unlock the
+   directory and add the staged dirents into the temporary directory.
+   Truncate the staging files.
+
+4. Use atomic extent swapping to exchange the new and old directory structures.
+   The old directory blocks are now attached to the temporary file.
+
+5. Reap the temporary file.
+
+**Future Work Question**: Should repair revalidate the dentry cache when
+rebuilding a directory?
+
+*Answer*: Yes, it should.
+
+In theory it is necessary to scan all dentry cache entries for a directory to
+ensure that one of the following apply:
+
+1. The cached dentry reflects an ondisk dirent in the new directory.
+
+2. The cached dentry no longer has a corresponding ondisk dirent in the new
+   directory and the dentry can be purged from the cache.
+
+3. The cached dentry no longer has an ondisk dirent but the dentry cannot be
+   purged.
+   This is the problem case.
+
+Unfortunately, the current dentry cache design doesn't provide a means to walk
+every child dentry of a specific directory, which makes this a hard problem.
+There is no known solution.
+
+The proposed patchset is the
+`directory repair
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-dirs>`_
+series.
+
+Parent Pointers
+```````````````
+
+A parent pointer is a piece of file metadata that enables a user to locate the
+file's parent directory without having to traverse the directory tree from the
+root.
+Without them, reconstruction of directory trees is hindered in much the same
+way that the historic lack of reverse space mapping information once hindered
+reconstruction of filesystem space metadata.
+The parent pointer feature, however, makes total directory reconstruction
+possible.
+
+XFS parent pointers include the dirent name and location of the entry within
+the parent directory.
+In other words, child files use extended attributes to store pointers to
+parents in the form ``(parent_inum, parent_gen, dirent_pos) → (dirent_name)``.
+The directory checking process can be strengthened to ensure that the target of
+each dirent also contains a parent pointer pointing back to the dirent.
+Likewise, each parent pointer can be checked by ensuring that the target of
+each parent pointer is a directory and that it contains a dirent matching
+the parent pointer.
+Both online and offline repair can use this strategy.
+
+**Note**: The ondisk format of parent pointers is not yet finalized.
+
++--------------------------------------------------------------------------+
+| **Historical Sidebar**:                                                  |
++--------------------------------------------------------------------------+
+| Directory parent pointers were first proposed as an XFS feature more     |
+| than a decade ago by SGI.                                                |
+| Each link from a parent directory to a child file is mirrored with an    |
+| extended attribute in the child that could be used to identify the       |
+| parent directory.                                                        |
+| Unfortunately, this early implementation had major shortcomings and was  |
+| never merged into Linux XFS:                                             |
+|                                                                          |
+| 1. The XFS codebase of the late 2000s did not have the infrastructure to |
+|    enforce strong referential integrity in the directory tree.           |
+|    It did not guarantee that a change in a forward link would always be  |
+|    followed up with the corresponding change to the reverse links.       |
+|                                                                          |
+| 2. Referential integrity was not integrated into offline repair.         |
+|    Checking and repairs were performed on mounted filesystems without    |
+|    taking any kernel or inode locks to coordinate access.                |
+|    It is not clear how this actually worked properly.                    |
+|                                                                          |
+| 3. The extended attribute did not record the name of the directory entry |
+|    in the parent, so the SGI parent pointer implementation cannot be     |
+|    used to reconnect the directory tree.                                 |
+|                                                                          |
+| 4. Extended attribute forks only support 65,536 extents, which means     |
+|    that parent pointer attribute creation is likely to fail at some      |
+|    point before the maximum file link count is achieved.                 |
+|                                                                          |
+| The original parent pointer design was too unstable for something like   |
+| a file system repair to depend on.                                       |
+| Allison Henderson, Chandan Babu, and Catherine Hoang are working on a    |
+| second implementation that solves all shortcomings of the first.         |
+| During 2022, Allison introduced log intent items to track physical       |
+| manipulations of the extended attribute structures.                      |
+| This solves the referential integrity problem by making it possible to   |
+| commit a dirent update and a parent pointer update in the same           |
+| transaction.                                                             |
+| Chandan increased the maximum extent counts of both data and attribute   |
+| forks, thereby ensuring that the extended attribute structure can grow   |
+| to handle the maximum hardlink count of any file.                        |
++--------------------------------------------------------------------------+
+
+Case Study: Repairing Directories with Parent Pointers
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+Directory rebuilding uses a :ref:`coordinated inode scan <iscan>` and
+a :ref:`directory entry live update hook <liveupdate>` as follows:
+
+1. Set up a temporary directory for generating the new directory structure,
+   an xfblob for storing entry names, and an xfarray for stashing directory
+   updates.
+
+2. Set up an inode scanner and hook into the directory entry code to receive
+   updates on directory operations.
+
+3. For each parent pointer found in each file scanned, decide if the parent
+   pointer references the directory of interest.
+   If so:
+
+   a. Stash an addname entry for this dirent in the xfarray for later.
+
+   b. When finished scanning that file, flush the stashed updates to the
+      temporary directory.
+
+4. For each live directory update received via the hook, decide if the child
+   has already been scanned.
+   If so:
+
+   a. Stash an addname or removename entry for this dirent update in the
+      xfarray for later.
+      We cannot write directly to the temporary directory because hook
+      functions are not allowed to modify filesystem metadata.
+      Instead, we stash updates in the xfarray and rely on the scanner thread
+      to apply the stashed updates to the temporary directory.
+
+5. When the scan is complete, atomically swap the contents of the temporary
+   directory and the directory being repaired.
+   The temporary directory now contains the damaged directory structure.
+
+6. Reap the temporary directory.
+
+7. Update the dirent position field of parent pointers as necessary.
+   This may require the queuing of a substantial number of xattr log intent
+   items.
+
+The proposed patchset is the
+`parent pointers directory repair
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=pptrs-online-dir-repair>`_
+series.
+
+**Unresolved Question**: How will repair ensure that the ``dirent_pos`` fields
+match in the reconstructed directory?
+
+*Answer*: There are a few ways to solve this problem:
+
+1. The field could be designated advisory, since the other three values are
+   sufficient to find the entry in the parent.
+   However, this makes indexed key lookup impossible while repairs are ongoing.
+
+2. We could allow creating directory entries at specified offsets, which solves
+   the referential integrity problem but runs the risk that dirent creation
+   will fail due to conflicts with the free space in the directory.
+
+   These conflicts could be resolved by appending the directory entry and
+   amending the xattr code to support updating an xattr key and reindexing the
+   dabtree, though this would have to be performed with the parent directory
+   still locked.
+
+3. Same as above, but remove the old parent pointer entry and add a new one
+   atomically.
+
+4. Change the ondisk xattr format to ``(parent_inum, name) → (parent_gen)``,
+   which would provide the attr name uniqueness that we require, without
+   forcing repair code to update the dirent position.
+   Unfortunately, this requires changes to the xattr code to support attr
+   names as long as 263 bytes.
+
+5. Change the ondisk xattr format to ``(parent_inum, hash(name)) →
+   (name, parent_gen)``.
+   If the hash is sufficiently resistant to collisions (e.g. sha256) then
+   this should provide the attr name uniqueness that we require.
+   Names shorter than 247 bytes could be stored directly.
+
+Discussion is ongoing under the `parent pointers patch deluge
+<https://www.spinics.net/lists/linux-xfs/msg69397.html>`_.
+
+Case Study: Repairing Parent Pointers
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+Online reconstruction of a file's parent pointer information works similarly to
+directory reconstruction:
+
+1. Set up a temporary file for generating a new extended attribute structure,
+   an `xfblob<xfblob>` for storing parent pointer names, and an xfarray for
+   stashing parent pointer updates.
+
+2. Set up an inode scanner and hook into the directory entry code to receive
+   updates on directory operations.
+
+3. For each directory entry found in each directory scanned, decide if the
+   dirent references the file of interest.
+   If so:
+
+   a. Stash an addpptr entry for this parent pointer in the xfblob and xfarray
+      for later.
+
+   b. When finished scanning the directory, flush the stashed updates to the
+      temporary directory.
+
+4. For each live directory update received via the hook, decide if the parent
+   has already been scanned.
+   If so:
+
+   a. Stash an addpptr or removepptr entry for this dirent update in the
+      xfarray for later.
+      We cannot write parent pointers directly to the temporary file because
+      hook functions are not allowed to modify filesystem metadata.
+      Instead, we stash updates in the xfarray and rely on the scanner thread
+      to apply the stashed parent pointer updates to the temporary file.
+
+5. Copy all non-parent pointer extended attributes to the temporary file.
+
+6. When the scan is complete, atomically swap the attribute fork of the
+   temporary file and the file being repaired.
+   The temporary file now contains the damaged extended attribute structure.
+
+7. Reap the temporary file.
+
+The proposed patchset is the
+`parent pointers repair
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=pptrs-online-parent-repair>`_
+series.
+
+Digression: Offline Checking of Parent Pointers
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+Examining parent pointers in offline repair works differently because corrupt
+files are erased long before directory tree connectivity checks are performed.
+Parent pointer checks are therefore a second pass to be added to the existing
+connectivity checks:
+
+1. After the set of surviving files has been established (i.e. phase 6),
+   walk the surviving directories of each AG in the filesystem.
+   This is already performed as part of the connectivity checks.
+
+2. For each directory entry found, record the name in an xfblob, and store
+   ``(child_ag_inum, parent_inum, parent_gen, dirent_pos)`` tuples in a
+   per-AG in-memory slab.
+
+3. For each AG in the filesystem,
+
+   a. Sort the per-AG tuples in order of child_ag_inum, parent_inum, and
+      dirent_pos.
+
+   b. For each inode in the AG,
+
+      1. Scan the inode for parent pointers.
+         Record the names in a per-file xfblob, and store ``(parent_inum,
+         parent_gen, dirent_pos)`` tuples in a per-file slab.
+
+      2. Sort the per-file tuples in order of parent_inum, and dirent_pos.
+
+      3. Position one slab cursor at the start of the inode's records in the
+         per-AG tuple slab.
+         This should be trivial since the per-AG tuples are in child inumber
+         order.
+
+      4. Position a second slab cursor at the start of the per-file tuple slab.
+
+      5. Iterate the two cursors in lockstep, comparing the parent_ino and
+         dirent_pos fields of the records under each cursor.
+
+         a. Tuples in the per-AG list but not the per-file list are missing and
+            need to be written to the inode.
+
+         b. Tuples in the per-file list but not the per-AG list are dangling
+            and need to be removed from the inode.
+
+         c. For tuples in both lists, update the parent_gen and name components
+            of the parent pointer if necessary.
+
+4. Move on to examining link counts, as we do today.
+
+The proposed patchset is the
+`offline parent pointers repair
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfsprogs-dev.git/log/?h=pptrs-repair>`_
+series.
+
+Rebuilding directories from parent pointers in offline repair is very
+challenging because it currently uses a single-pass scan of the filesystem
+during phase 3 to decide which files are corrupt enough to be zapped.
+This scan would have to be converted into a multi-pass scan:
+
+1. The first pass of the scan zaps corrupt inodes, forks, and attributes
+   much as it does now.
+   Corrupt directories are noted but not zapped.
+
+2. The next pass records parent pointers pointing to the directories noted
+   as being corrupt in the first pass.
+   This second pass may have to happen after the phase 4 scan for duplicate
+   blocks, if phase 4 is also capable of zapping directories.
+
+3. The third pass resets corrupt directories to an empty shortform directory.
+   Free space metadata has not been ensured yet, so repair cannot yet use the
+   directory building code in libxfs.
+
+4. At the start of phase 6, space metadata have been rebuilt.
+   Use the parent pointer information recorded during step 2 to reconstruct
+   the dirents and add them to the now-empty directories.
+
+This code has not yet been constructed.
+
+.. _orphanage:
+
+The Orphanage
+-------------
+
+Filesystems present files as a directed, and hopefully acyclic, graph.
+In other words, a tree.
+The root of the filesystem is a directory, and each entry in a directory points
+downwards either to more subdirectories or to non-directory files.
+Unfortunately, a disruption in the directory graph pointers result in a
+disconnected graph, which makes files impossible to access via regular path
+resolution.
+
+Without parent pointers, the directory parent pointer online scrub code can
+detect a dotdot entry pointing to a parent directory that doesn't have a link
+back to the child directory and the file link count checker can detect a file
+that isn't pointed to by any directory in the filesystem.
+If such a file has a positive link count, the file is an orphan.
+
+With parent pointers, directories can be rebuilt by scanning parent pointers
+and parent pointers can be rebuilt by scanning directories.
+This should reduce the incidence of files ending up in ``/lost+found``.
+
+When orphans are found, they should be reconnected to the directory tree.
+Offline fsck solves the problem by creating a directory ``/lost+found`` to
+serve as an orphanage, and linking orphan files into the orphanage by using the
+inumber as the name.
+Reparenting a file to the orphanage does not reset any of its permissions or
+ACLs.
+
+This process is more involved in the kernel than it is in userspace.
+The directory and file link count repair setup functions must use the regular
+VFS mechanisms to create the orphanage directory with all the necessary
+security attributes and dentry cache entries, just like a regular directory
+tree modification.
+
+Orphaned files are adopted by the orphanage as follows:
+
+1. Call ``xrep_orphanage_try_create`` at the start of the scrub setup function
+   to try to ensure that the lost and found directory actually exists.
+   This also attaches the orphanage directory to the scrub context.
+
+2. If the decision is made to reconnect a file, take the IOLOCK of both the
+   orphanage and the file being reattached.
+   The ``xrep_orphanage_iolock_two`` function follows the inode locking
+   strategy discussed earlier.
+
+3. Call ``xrep_orphanage_compute_blkres`` and ``xrep_orphanage_compute_name``
+   to compute the new name in the orphanage and the block reservation required.
+
+4. Use ``xrep_orphanage_adoption_prep`` to reserve resources to the repair
+   transaction.
+
+5. Call ``xrep_orphanage_adopt`` to reparent the orphaned file into the lost
+   and found, and update the kernel dentry cache.
+
+The proposed patches are in the
+`orphanage adoption
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=repair-orphanage>`_
+series.
+
+6. Userspace Algorithms and Data Structures
+===========================================
+
+This section discusses the key algorithms and data structures of the userspace
+program, ``xfs_scrub``, that provide the ability to drive metadata checks and
+repairs in the kernel, verify file data, and look for other potential problems.
+
+.. _scrubcheck:
+
+Checking Metadata
+-----------------
+
+Recall the :ref:`phases of fsck work<scrubphases>` outlined earlier.
+That structure follows naturally from the data dependencies designed into the
+filesystem from its beginnings in 1993.
+In XFS, there are several groups of metadata dependencies:
+
+a. Filesystem summary counts depend on consistency within the inode indices,
+   the allocation group space btrees, and the realtime volume space
+   information.
+
+b. Quota resource counts depend on consistency within the quota file data
+   forks, inode indices, inode records, and the forks of every file on the
+   system.
+
+c. The naming hierarchy depends on consistency within the directory and
+   extended attribute structures.
+   This includes file link counts.
+
+d. Directories, extended attributes, and file data depend on consistency within
+   the file forks that map directory and extended attribute data to physical
+   storage media.
+
+e. The file forks depends on consistency within inode records and the space
+   metadata indices of the allocation groups and the realtime volume.
+   This includes quota and realtime metadata files.
+
+f. Inode records depends on consistency within the inode metadata indices.
+
+g. Realtime space metadata depend on the inode records and data forks of the
+   realtime metadata inodes.
+
+h. The allocation group metadata indices (free space, inodes, reference count,
+   and reverse mapping btrees) depend on consistency within the AG headers and
+   between all the AG metadata btrees.
+
+i. ``xfs_scrub`` depends on the filesystem being mounted and kernel support
+   for online fsck functionality.
+
+Therefore, a metadata dependency graph is a convenient way to schedule checking
+operations in the ``xfs_scrub`` program:
+
+- Phase 1 checks that the provided path maps to an XFS filesystem and detect
+  the kernel's scrubbing abilities, which validates group (i).
+
+- Phase 2 scrubs groups (g) and (h) in parallel using a threaded workqueue.
+
+- Phase 3 scans inodes in parallel.
+  For each inode, groups (f), (e), and (d) are checked, in that order.
+
+- Phase 4 repairs everything in groups (i) through (d) so that phases 5 and 6
+  may run reliably.
+
+- Phase 5 starts by checking groups (b) and (c) in parallel before moving on
+  to checking names.
+
+- Phase 6 depends on groups (i) through (b) to find file data blocks to verify,
+  to read them, and to report which blocks of which files are affected.
+
+- Phase 7 checks group (a), having validated everything else.
+
+Notice that the data dependencies between groups are enforced by the structure
+of the program flow.
+
+Parallel Inode Scans
+--------------------
+
+An XFS filesystem can easily contain hundreds of millions of inodes.
+Given that XFS targets installations with large high-performance storage,
+it is desirable to scrub inodes in parallel to minimize runtime, particularly
+if the program has been invoked manually from a command line.
+This requires careful scheduling to keep the threads as evenly loaded as
+possible.
+
+Early iterations of the ``xfs_scrub`` inode scanner naïvely created a single
+workqueue and scheduled a single workqueue item per AG.
+Each workqueue item walked the inode btree (with ``XFS_IOC_INUMBERS``) to find
+inode chunks and then called bulkstat (``XFS_IOC_BULKSTAT``) to gather enough
+information to construct file handles.
+The file handle was then passed to a function to generate scrub items for each
+metadata object of each inode.
+This simple algorithm leads to thread balancing problems in phase 3 if the
+filesystem contains one AG with a few large sparse files and the rest of the
+AGs contain many smaller files.
+The inode scan dispatch function was not sufficiently granular; it should have
+been dispatching at the level of individual inodes, or, to constrain memory
+consumption, inode btree records.
+
+Thanks to Dave Chinner, bounded workqueues in userspace enable ``xfs_scrub`` to
+avoid this problem with ease by adding a second workqueue.
+Just like before, the first workqueue is seeded with one workqueue item per AG,
+and it uses INUMBERS to find inode btree chunks.
+The second workqueue, however, is configured with an upper bound on the number
+of items that can be waiting to be run.
+Each inode btree chunk found by the first workqueue's workers are queued to the
+second workqueue, and it is this second workqueue that queries BULKSTAT,
+creates a file handle, and passes it to a function to generate scrub items for
+each metadata object of each inode.
+If the second workqueue is too full, the workqueue add function blocks the
+first workqueue's workers until the backlog eases.
+This doesn't completely solve the balancing problem, but reduces it enough to
+move on to more pressing issues.
+
+The proposed patchsets are the scrub
+`performance tweaks
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfsprogs-dev.git/log/?h=scrub-performance-tweaks>`_
+and the
+`inode scan rebalance
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfsprogs-dev.git/log/?h=scrub-iscan-rebalance>`_
+series.
+
+.. _scrubrepair:
+
+Scheduling Repairs
+------------------
+
+During phase 2, corruptions and inconsistencies reported in any AGI header or
+inode btree are repaired immediately, because phase 3 relies on proper
+functioning of the inode indices to find inodes to scan.
+Failed repairs are rescheduled to phase 4.
+Problems reported in any other space metadata are deferred to phase 4.
+Optimization opportunities are always deferred to phase 4, no matter their
+origin.
+
+During phase 3, corruptions and inconsistencies reported in any part of a
+file's metadata are repaired immediately if all space metadata were validated
+during phase 2.
+Repairs that fail or cannot be repaired immediately are scheduled for phase 4.
+
+In the original design of ``xfs_scrub``, it was thought that repairs would be
+so infrequent that the ``struct xfs_scrub_metadata`` objects used to
+communicate with the kernel could also be used as the primary object to
+schedule repairs.
+With recent increases in the number of optimizations possible for a given
+filesystem object, it became much more memory-efficient to track all eligible
+repairs for a given filesystem object with a single repair item.
+Each repair item represents a single lockable object -- AGs, metadata files,
+individual inodes, or a class of summary information.
+
+Phase 4 is responsible for scheduling a lot of repair work in as quick a
+manner as is practical.
+The :ref:`data dependencies <scrubcheck>` outlined earlier still apply, which
+means that ``xfs_scrub`` must try to complete the repair work scheduled by
+phase 2 before trying repair work scheduled by phase 3.
+The repair process is as follows:
+
+1. Start a round of repair with a workqueue and enough workers to keep the CPUs
+   as busy as the user desires.
+
+   a. For each repair item queued by phase 2,
+
+      i.   Ask the kernel to repair everything listed in the repair item for a
+           given filesystem object.
+
+      ii.  Make a note if the kernel made any progress in reducing the number
+           of repairs needed for this object.
+
+      iii. If the object no longer requires repairs, revalidate all metadata
+           associated with this object.
+           If the revalidation succeeds, drop the repair item.
+           If not, requeue the item for more repairs.
+
+   b. If any repairs were made, jump back to 1a to retry all the phase 2 items.
+
+   c. For each repair item queued by phase 3,
+
+      i.   Ask the kernel to repair everything listed in the repair item for a
+           given filesystem object.
+
+      ii.  Make a note if the kernel made any progress in reducing the number
+           of repairs needed for this object.
+
+      iii. If the object no longer requires repairs, revalidate all metadata
+           associated with this object.
+           If the revalidation succeeds, drop the repair item.
+           If not, requeue the item for more repairs.
+
+   d. If any repairs were made, jump back to 1c to retry all the phase 3 items.
+
+2. If step 1 made any repair progress of any kind, jump back to step 1 to start
+   another round of repair.
+
+3. If there are items left to repair, run them all serially one more time.
+   Complain if the repairs were not successful, since this is the last chance
+   to repair anything.
+
+Corruptions and inconsistencies encountered during phases 5 and 7 are repaired
+immediately.
+Corrupt file data blocks reported by phase 6 cannot be recovered by the
+filesystem.
+
+The proposed patchsets are the
+`repair warning improvements
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfsprogs-dev.git/log/?h=scrub-better-repair-warnings>`_,
+refactoring of the
+`repair data dependency
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfsprogs-dev.git/log/?h=scrub-repair-data-deps>`_
+and
+`object tracking
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfsprogs-dev.git/log/?h=scrub-object-tracking>`_,
+and the
+`repair scheduling
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfsprogs-dev.git/log/?h=scrub-repair-scheduling>`_
+improvement series.
+
+Checking Names for Confusable Unicode Sequences
+-----------------------------------------------
+
+If ``xfs_scrub`` succeeds in validating the filesystem metadata by the end of
+phase 4, it moves on to phase 5, which checks for suspicious looking names in
+the filesystem.
+These names consist of the filesystem label, names in directory entries, and
+the names of extended attributes.
+Like most Unix filesystems, XFS imposes the sparest of constraints on the
+contents of a name:
+
+- Slashes and null bytes are not allowed in directory entries.
+
+- Null bytes are not allowed in userspace-visible extended attributes.
+
+- Null bytes are not allowed in the filesystem label.
+
+Directory entries and attribute keys store the length of the name explicitly
+ondisk, which means that nulls are not name terminators.
+For this section, the term "naming domain" refers to any place where names are
+presented together -- all the names in a directory, or all the attributes of a
+file.
+
+Although the Unix naming constraints are very permissive, the reality of most
+modern-day Linux systems is that programs work with Unicode character code
+points to support international languages.
+These programs typically encode those code points in UTF-8 when interfacing
+with the C library because the kernel expects null-terminated names.
+In the common case, therefore, names found in an XFS filesystem are actually
+UTF-8 encoded Unicode data.
+
+To maximize its expressiveness, the Unicode standard defines separate control
+points for various characters that render similarly or identically in writing
+systems around the world.
+For example, the character "Cyrillic Small Letter A" U+0430 "а" often renders
+identically to "Latin Small Letter A" U+0061 "a".
+
+The standard also permits characters to be constructed in multiple ways --
+either by using a defined code point, or by combining one code point with
+various combining marks.
+For example, the character "Angstrom Sign U+212B "Å" can also be expressed
+as "Latin Capital Letter A" U+0041 "A" followed by "Combining Ring Above"
+U+030A "◌̊".
+Both sequences render identically.
+
+Like the standards that preceded it, Unicode also defines various control
+characters to alter the presentation of text.
+For example, the character "Right-to-Left Override" U+202E can trick some
+programs into rendering "moo\\xe2\\x80\\xaegnp.txt" as "mootxt.png".
+A second category of rendering problems involves whitespace characters.
+If the character "Zero Width Space" U+200B is encountered in a file name, the
+name will render identically to a name that does not have the zero width
+space.
+
+If two names within a naming domain have different byte sequences but render
+identically, a user may be confused by it.
+The kernel, in its indifference to upper level encoding schemes, permits this.
+Most filesystem drivers persist the byte sequence names that are given to them
+by the VFS.
+
+Techniques for detecting confusable names are explained in great detail in
+sections 4 and 5 of the
+`Unicode Security Mechanisms <https://unicode.org/reports/tr39/>`_
+document.
+When ``xfs_scrub`` detects UTF-8 encoding in use on a system, it uses the
+Unicode normalization form NFD in conjunction with the confusable name
+detection component of
+`libicu <https://github.com/unicode-org/icu>`_
+to identify names with a directory or within a file's extended attributes that
+could be confused for each other.
+Names are also checked for control characters, non-rendering characters, and
+mixing of bidirectional characters.
+All of these potential issues are reported to the system administrator during
+phase 5.
+
+Media Verification of File Data Extents
+---------------------------------------
+
+The system administrator can elect to initiate a media scan of all file data
+blocks.
+This scan after validation of all filesystem metadata (except for the summary
+counters) as phase 6.
+The scan starts by calling ``FS_IOC_GETFSMAP`` to scan the filesystem space map
+to find areas that are allocated to file data fork extents.
+Gaps between data fork extents that are smaller than 64k are treated as if
+they were data fork extents to reduce the command setup overhead.
+When the space map scan accumulates a region larger than 32MB, a media
+verification request is sent to the disk as a directio read of the raw block
+device.
+
+If the verification read fails, ``xfs_scrub`` retries with single-block reads
+to narrow down the failure to the specific region of the media and recorded.
+When it has finished issuing verification requests, it again uses the space
+mapping ioctl to map the recorded media errors back to metadata structures
+and report what has been lost.
+For media errors in blocks owned by files, parent pointers can be used to
+construct file paths from inode numbers for user-friendly reporting.
+
+7. Conclusion and Future Work
+=============================
+
+It is hoped that the reader of this document has followed the designs laid out
+in this document and now has some familiarity with how XFS performs online
+rebuilding of its metadata indices, and how filesystem users can interact with
+that functionality.
+Although the scope of this work is daunting, it is hoped that this guide will
+make it easier for code readers to understand what has been built, for whom it
+has been built, and why.
+Please feel free to contact the XFS mailing list with questions.
+
+FIEXCHANGE_RANGE
+----------------
+
+As discussed earlier, a second frontend to the atomic extent swap mechanism is
+a new ioctl call that userspace programs can use to commit updates to files
+atomically.
+This frontend has been out for review for several years now, though the
+necessary refinements to online repair and lack of customer demand mean that
+the proposal has not been pushed very hard.
+
+Extent Swapping with Regular User Files
+```````````````````````````````````````
+
+As mentioned earlier, XFS has long had the ability to swap extents between
+files, which is used almost exclusively by ``xfs_fsr`` to defragment files.
+The earliest form of this was the fork swap mechanism, where the entire
+contents of data forks could be exchanged between two files by exchanging the
+raw bytes in each inode fork's immediate area.
+When XFS v5 came along with self-describing metadata, this old mechanism grew
+some log support to continue rewriting the owner fields of BMBT blocks during
+log recovery.
+When the reverse mapping btree was later added to XFS, the only way to maintain
+the consistency of the fork mappings with the reverse mapping index was to
+develop an iterative mechanism that used deferred bmap and rmap operations to
+swap mappings one at a time.
+This mechanism is identical to steps 2-3 from the procedure above except for
+the new tracking items, because the atomic extent swap mechanism is an
+iteration of an existing mechanism and not something totally novel.
+For the narrow case of file defragmentation, the file contents must be
+identical, so the recovery guarantees are not much of a gain.
+
+Atomic extent swapping is much more flexible than the existing swapext
+implementations because it can guarantee that the caller never sees a mix of
+old and new contents even after a crash, and it can operate on two arbitrary
+file fork ranges.
+The extra flexibility enables several new use cases:
+
+- **Atomic commit of file writes**: A userspace process opens a file that it
+  wants to update.
+  Next, it opens a temporary file and calls the file clone operation to reflink
+  the first file's contents into the temporary file.
+  Writes to the original file should instead be written to the temporary file.
+  Finally, the process calls the atomic extent swap system call
+  (``FIEXCHANGE_RANGE``) to exchange the file contents, thereby committing all
+  of the updates to the original file, or none of them.
+
+.. _swapext_if_unchanged:
+
+- **Transactional file updates**: The same mechanism as above, but the caller
+  only wants the commit to occur if the original file's contents have not
+  changed.
+  To make this happen, the calling process snapshots the file modification and
+  change timestamps of the original file before reflinking its data to the
+  temporary file.
+  When the program is ready to commit the changes, it passes the timestamps
+  into the kernel as arguments to the atomic extent swap system call.
+  The kernel only commits the changes if the provided timestamps match the
+  original file.
+
+- **Emulation of atomic block device writes**: Export a block device with a
+  logical sector size matching the filesystem block size to force all writes
+  to be aligned to the filesystem block size.
+  Stage all writes to a temporary file, and when that is complete, call the
+  atomic extent swap system call with a flag to indicate that holes in the
+  temporary file should be ignored.
+  This emulates an atomic device write in software, and can support arbitrary
+  scattered writes.
+
+Vectorized Scrub
+----------------
+
+As it turns out, the :ref:`refactoring <scrubrepair>` of repair items mentioned
+earlier was a catalyst for enabling a vectorized scrub system call.
+Since 2018, the cost of making a kernel call has increased considerably on some
+systems to mitigate the effects of speculative execution attacks.
+This incentivizes program authors to make as few system calls as possible to
+reduce the number of times an execution path crosses a security boundary.
+
+With vectorized scrub, userspace pushes to the kernel the identity of a
+filesystem object, a list of scrub types to run against that object, and a
+simple representation of the data dependencies between the selected scrub
+types.
+The kernel executes as much of the caller's plan as it can until it hits a
+dependency that cannot be satisfied due to a corruption, and tells userspace
+how much was accomplished.
+It is hoped that ``io_uring`` will pick up enough of this functionality that
+online fsck can use that instead of adding a separate vectored scrub system
+call to XFS.
+
+The relevant patchsets are the
+`kernel vectorized scrub
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=vectorized-scrub>`_
+and
+`userspace vectorized scrub
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfsprogs-dev.git/log/?h=vectorized-scrub>`_
+series.
+
+Quality of Service Targets for Scrub
+------------------------------------
+
+One serious shortcoming of the online fsck code is that the amount of time that
+it can spend in the kernel holding resource locks is basically unbounded.
+Userspace is allowed to send a fatal signal to the process which will cause
+``xfs_scrub`` to exit when it reaches a good stopping point, but there's no way
+for userspace to provide a time budget to the kernel.
+Given that the scrub codebase has helpers to detect fatal signals, it shouldn't
+be too much work to allow userspace to specify a timeout for a scrub/repair
+operation and abort the operation if it exceeds budget.
+However, most repair functions have the property that once they begin to touch
+ondisk metadata, the operation cannot be cancelled cleanly, after which a QoS
+timeout is no longer useful.
+
+Defragmenting Free Space
+------------------------
+
+Over the years, many XFS users have requested the creation of a program to
+clear a portion of the physical storage underlying a filesystem so that it
+becomes a contiguous chunk of free space.
+Call this free space defragmenter ``clearspace`` for short.
+
+The first piece the ``clearspace`` program needs is the ability to read the
+reverse mapping index from userspace.
+This already exists in the form of the ``FS_IOC_GETFSMAP`` ioctl.
+The second piece it needs is a new fallocate mode
+(``FALLOC_FL_MAP_FREE_SPACE``) that allocates the free space in a region and
+maps it to a file.
+Call this file the "space collector" file.
+The third piece is the ability to force an online repair.
+
+To clear all the metadata out of a portion of physical storage, clearspace
+uses the new fallocate map-freespace call to map any free space in that region
+to the space collector file.
+Next, clearspace finds all metadata blocks in that region by way of
+``GETFSMAP`` and issues forced repair requests on the data structure.
+This often results in the metadata being rebuilt somewhere that is not being
+cleared.
+After each relocation, clearspace calls the "map free space" function again to
+collect any newly freed space in the region being cleared.
+
+To clear all the file data out of a portion of the physical storage, clearspace
+uses the FSMAP information to find relevant file data blocks.
+Having identified a good target, it uses the ``FICLONERANGE`` call on that part
+of the file to try to share the physical space with a dummy file.
+Cloning the extent means that the original owners cannot overwrite the
+contents; any changes will be written somewhere else via copy-on-write.
+Clearspace makes its own copy of the frozen extent in an area that is not being
+cleared, and uses ``FIEDEUPRANGE`` (or the :ref:`atomic extent swap
+<swapext_if_unchanged>` feature) to change the target file's data extent
+mapping away from the area being cleared.
+When all other mappings have been moved, clearspace reflinks the space into the
+space collector file so that it becomes unavailable.
+
+There are further optimizations that could apply to the above algorithm.
+To clear a piece of physical storage that has a high sharing factor, it is
+strongly desirable to retain this sharing factor.
+In fact, these extents should be moved first to maximize sharing factor after
+the operation completes.
+To make this work smoothly, clearspace needs a new ioctl
+(``FS_IOC_GETREFCOUNTS``) to report reference count information to userspace.
+With the refcount information exposed, clearspace can quickly find the longest,
+most shared data extents in the filesystem, and target them first.
+
+**Future Work Question**: How might the filesystem move inode chunks?
+
+*Answer*: To move inode chunks, Dave Chinner constructed a prototype program
+that creates a new file with the old contents and then locklessly runs around
+the filesystem updating directory entries.
+The operation cannot complete if the filesystem goes down.
+That problem isn't totally insurmountable: create an inode remapping table
+hidden behind a jump label, and a log item that tracks the kernel walking the
+filesystem to update directory entries.
+The trouble is, the kernel can't do anything about open files, since it cannot
+revoke them.
+
+**Future Work Question**: Can static keys be used to minimize the cost of
+supporting ``revoke()`` on XFS files?
+
+*Answer*: Yes.
+Until the first revocation, the bailout code need not be in the call path at
+all.
+
+The relevant patchsets are the
+`kernel freespace defrag
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=defrag-freespace>`_
+and
+`userspace freespace defrag
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfsprogs-dev.git/log/?h=defrag-freespace>`_
+series.
+
+Shrinking Filesystems
+---------------------
+
+Removing the end of the filesystem ought to be a simple matter of evacuating
+the data and metadata at the end of the filesystem, and handing the freed space
+to the shrink code.
+That requires an evacuation of the space at end of the filesystem, which is a
+use of free space defragmentation!
diff --git a/Documentation/filesystems/xfs/xfs-self-describing-metadata.rst b/Documentation/filesystems/xfs/xfs-self-describing-metadata.rst
new file mode 100644
index 000000000000..a10c4ae6955e
--- /dev/null
+++ b/Documentation/filesystems/xfs/xfs-self-describing-metadata.rst
@@ -0,0 +1,353 @@
+.. SPDX-License-Identifier: GPL-2.0
+.. _xfs_self_describing_metadata:
+
+============================
+XFS Self Describing Metadata
+============================
+
+Introduction
+============
+
+The largest scalability problem facing XFS is not one of algorithmic
+scalability, but of verification of the filesystem structure. Scalabilty of the
+structures and indexes on disk and the algorithms for iterating them are
+adequate for supporting PB scale filesystems with billions of inodes, however it
+is this very scalability that causes the verification problem.
+
+Almost all metadata on XFS is dynamically allocated. The only fixed location
+metadata is the allocation group headers (SB, AGF, AGFL and AGI), while all
+other metadata structures need to be discovered by walking the filesystem
+structure in different ways. While this is already done by userspace tools for
+validating and repairing the structure, there are limits to what they can
+verify, and this in turn limits the supportable size of an XFS filesystem.
+
+For example, it is entirely possible to manually use xfs_db and a bit of
+scripting to analyse the structure of a 100TB filesystem when trying to
+determine the root cause of a corruption problem, but it is still mainly a
+manual task of verifying that things like single bit errors or misplaced writes
+weren't the ultimate cause of a corruption event. It may take a few hours to a
+few days to perform such forensic analysis, so for at this scale root cause
+analysis is entirely possible.
+
+However, if we scale the filesystem up to 1PB, we now have 10x as much metadata
+to analyse and so that analysis blows out towards weeks/months of forensic work.
+Most of the analysis work is slow and tedious, so as the amount of analysis goes
+up, the more likely that the cause will be lost in the noise.  Hence the primary
+concern for supporting PB scale filesystems is minimising the time and effort
+required for basic forensic analysis of the filesystem structure.
+
+
+Self Describing Metadata
+========================
+
+One of the problems with the current metadata format is that apart from the
+magic number in the metadata block, we have no other way of identifying what it
+is supposed to be. We can't even identify if it is the right place. Put simply,
+you can't look at a single metadata block in isolation and say "yes, it is
+supposed to be there and the contents are valid".
+
+Hence most of the time spent on forensic analysis is spent doing basic
+verification of metadata values, looking for values that are in range (and hence
+not detected by automated verification checks) but are not correct. Finding and
+understanding how things like cross linked block lists (e.g. sibling
+pointers in a btree end up with loops in them) are the key to understanding what
+went wrong, but it is impossible to tell what order the blocks were linked into
+each other or written to disk after the fact.
+
+Hence we need to record more information into the metadata to allow us to
+quickly determine if the metadata is intact and can be ignored for the purpose
+of analysis. We can't protect against every possible type of error, but we can
+ensure that common types of errors are easily detectable.  Hence the concept of
+self describing metadata.
+
+The first, fundamental requirement of self describing metadata is that the
+metadata object contains some form of unique identifier in a well known
+location. This allows us to identify the expected contents of the block and
+hence parse and verify the metadata object. IF we can't independently identify
+the type of metadata in the object, then the metadata doesn't describe itself
+very well at all!
+
+Luckily, almost all XFS metadata has magic numbers embedded already - only the
+AGFL, remote symlinks and remote attribute blocks do not contain identifying
+magic numbers. Hence we can change the on-disk format of all these objects to
+add more identifying information and detect this simply by changing the magic
+numbers in the metadata objects. That is, if it has the current magic number,
+the metadata isn't self identifying. If it contains a new magic number, it is
+self identifying and we can do much more expansive automated verification of the
+metadata object at runtime, during forensic analysis or repair.
+
+As a primary concern, self describing metadata needs some form of overall
+integrity checking. We cannot trust the metadata if we cannot verify that it has
+not been changed as a result of external influences. Hence we need some form of
+integrity check, and this is done by adding CRC32c validation to the metadata
+block. If we can verify the block contains the metadata it was intended to
+contain, a large amount of the manual verification work can be skipped.
+
+CRC32c was selected as metadata cannot be more than 64k in length in XFS and
+hence a 32 bit CRC is more than sufficient to detect multi-bit errors in
+metadata blocks. CRC32c is also now hardware accelerated on common CPUs so it is
+fast. So while CRC32c is not the strongest of possible integrity checks that
+could be used, it is more than sufficient for our needs and has relatively
+little overhead. Adding support for larger integrity fields and/or algorithms
+does really provide any extra value over CRC32c, but it does add a lot of
+complexity and so there is no provision for changing the integrity checking
+mechanism.
+
+Self describing metadata needs to contain enough information so that the
+metadata block can be verified as being in the correct place without needing to
+look at any other metadata. This means it needs to contain location information.
+Just adding a block number to the metadata is not sufficient to protect against
+mis-directed writes - a write might be misdirected to the wrong LUN and so be
+written to the "correct block" of the wrong filesystem. Hence location
+information must contain a filesystem identifier as well as a block number.
+
+Another key information point in forensic analysis is knowing who the metadata
+block belongs to. We already know the type, the location, that it is valid
+and/or corrupted, and how long ago that it was last modified. Knowing the owner
+of the block is important as it allows us to find other related metadata to
+determine the scope of the corruption. For example, if we have a extent btree
+object, we don't know what inode it belongs to and hence have to walk the entire
+filesystem to find the owner of the block. Worse, the corruption could mean that
+no owner can be found (i.e. it's an orphan block), and so without an owner field
+in the metadata we have no idea of the scope of the corruption. If we have an
+owner field in the metadata object, we can immediately do top down validation to
+determine the scope of the problem.
+
+Different types of metadata have different owner identifiers. For example,
+directory, attribute and extent tree blocks are all owned by an inode, while
+freespace btree blocks are owned by an allocation group. Hence the size and
+contents of the owner field are determined by the type of metadata object we are
+looking at.  The owner information can also identify misplaced writes (e.g.
+freespace btree block written to the wrong AG).
+
+Self describing metadata also needs to contain some indication of when it was
+written to the filesystem. One of the key information points when doing forensic
+analysis is how recently the block was modified. Correlation of set of corrupted
+metadata blocks based on modification times is important as it can indicate
+whether the corruptions are related, whether there's been multiple corruption
+events that lead to the eventual failure, and even whether there are corruptions
+present that the run-time verification is not detecting.
+
+For example, we can determine whether a metadata object is supposed to be free
+space or still allocated if it is still referenced by its owner by looking at
+when the free space btree block that contains the block was last written
+compared to when the metadata object itself was last written.  If the free space
+block is more recent than the object and the object's owner, then there is a
+very good chance that the block should have been removed from the owner.
+
+To provide this "written timestamp", each metadata block gets the Log Sequence
+Number (LSN) of the most recent transaction it was modified on written into it.
+This number will always increase over the life of the filesystem, and the only
+thing that resets it is running xfs_repair on the filesystem. Further, by use of
+the LSN we can tell if the corrupted metadata all belonged to the same log
+checkpoint and hence have some idea of how much modification occurred between
+the first and last instance of corrupt metadata on disk and, further, how much
+modification occurred between the corruption being written and when it was
+detected.
+
+Runtime Validation
+==================
+
+Validation of self-describing metadata takes place at runtime in two places:
+
+	- immediately after a successful read from disk
+	- immediately prior to write IO submission
+
+The verification is completely stateless - it is done independently of the
+modification process, and seeks only to check that the metadata is what it says
+it is and that the metadata fields are within bounds and internally consistent.
+As such, we cannot catch all types of corruption that can occur within a block
+as there may be certain limitations that operational state enforces of the
+metadata, or there may be corruption of interblock relationships (e.g. corrupted
+sibling pointer lists). Hence we still need stateful checking in the main code
+body, but in general most of the per-field validation is handled by the
+verifiers.
+
+For read verification, the caller needs to specify the expected type of metadata
+that it should see, and the IO completion process verifies that the metadata
+object matches what was expected. If the verification process fails, then it
+marks the object being read as EFSCORRUPTED. The caller needs to catch this
+error (same as for IO errors), and if it needs to take special action due to a
+verification error it can do so by catching the EFSCORRUPTED error value. If we
+need more discrimination of error type at higher levels, we can define new
+error numbers for different errors as necessary.
+
+The first step in read verification is checking the magic number and determining
+whether CRC validating is necessary. If it is, the CRC32c is calculated and
+compared against the value stored in the object itself. Once this is validated,
+further checks are made against the location information, followed by extensive
+object specific metadata validation. If any of these checks fail, then the
+buffer is considered corrupt and the EFSCORRUPTED error is set appropriately.
+
+Write verification is the opposite of the read verification - first the object
+is extensively verified and if it is OK we then update the LSN from the last
+modification made to the object, After this, we calculate the CRC and insert it
+into the object. Once this is done the write IO is allowed to continue. If any
+error occurs during this process, the buffer is again marked with a EFSCORRUPTED
+error for the higher layers to catch.
+
+Structures
+==========
+
+A typical on-disk structure needs to contain the following information::
+
+    struct xfs_ondisk_hdr {
+	    __be32  magic;		/* magic number */
+	    __be32  crc;		/* CRC, not logged */
+	    uuid_t  uuid;		/* filesystem identifier */
+	    __be64  owner;		/* parent object */
+	    __be64  blkno;		/* location on disk */
+	    __be64  lsn;		/* last modification in log, not logged */
+    };
+
+Depending on the metadata, this information may be part of a header structure
+separate to the metadata contents, or may be distributed through an existing
+structure. The latter occurs with metadata that already contains some of this
+information, such as the superblock and AG headers.
+
+Other metadata may have different formats for the information, but the same
+level of information is generally provided. For example:
+
+	- short btree blocks have a 32 bit owner (ag number) and a 32 bit block
+	  number for location. The two of these combined provide the same
+	  information as @owner and @blkno in eh above structure, but using 8
+	  bytes less space on disk.
+
+	- directory/attribute node blocks have a 16 bit magic number, and the
+	  header that contains the magic number has other information in it as
+	  well. hence the additional metadata headers change the overall format
+	  of the metadata.
+
+A typical buffer read verifier is structured as follows::
+
+    #define XFS_FOO_CRC_OFF		offsetof(struct xfs_ondisk_hdr, crc)
+
+    static void
+    xfs_foo_read_verify(
+	    struct xfs_buf	*bp)
+    {
+	struct xfs_mount *mp = bp->b_mount;
+
+	    if ((xfs_sb_version_hascrc(&mp->m_sb) &&
+		!xfs_verify_cksum(bp->b_addr, BBTOB(bp->b_length),
+					    XFS_FOO_CRC_OFF)) ||
+		!xfs_foo_verify(bp)) {
+		    XFS_CORRUPTION_ERROR(__func__, XFS_ERRLEVEL_LOW, mp, bp->b_addr);
+		    xfs_buf_ioerror(bp, EFSCORRUPTED);
+	    }
+    }
+
+The code ensures that the CRC is only checked if the filesystem has CRCs enabled
+by checking the superblock of the feature bit, and then if the CRC verifies OK
+(or is not needed) it verifies the actual contents of the block.
+
+The verifier function will take a couple of different forms, depending on
+whether the magic number can be used to determine the format of the block. In
+the case it can't, the code is structured as follows::
+
+    static bool
+    xfs_foo_verify(
+	    struct xfs_buf		*bp)
+    {
+	    struct xfs_mount	*mp = bp->b_mount;
+	    struct xfs_ondisk_hdr	*hdr = bp->b_addr;
+
+	    if (hdr->magic != cpu_to_be32(XFS_FOO_MAGIC))
+		    return false;
+
+	    if (!xfs_sb_version_hascrc(&mp->m_sb)) {
+		    if (!uuid_equal(&hdr->uuid, &mp->m_sb.sb_uuid))
+			    return false;
+		    if (bp->b_bn != be64_to_cpu(hdr->blkno))
+			    return false;
+		    if (hdr->owner == 0)
+			    return false;
+	    }
+
+	    /* object specific verification checks here */
+
+	    return true;
+    }
+
+If there are different magic numbers for the different formats, the verifier
+will look like::
+
+    static bool
+    xfs_foo_verify(
+	    struct xfs_buf		*bp)
+    {
+	    struct xfs_mount	*mp = bp->b_mount;
+	    struct xfs_ondisk_hdr	*hdr = bp->b_addr;
+
+	    if (hdr->magic == cpu_to_be32(XFS_FOO_CRC_MAGIC)) {
+		    if (!uuid_equal(&hdr->uuid, &mp->m_sb.sb_uuid))
+			    return false;
+		    if (bp->b_bn != be64_to_cpu(hdr->blkno))
+			    return false;
+		    if (hdr->owner == 0)
+			    return false;
+	    } else if (hdr->magic != cpu_to_be32(XFS_FOO_MAGIC))
+		    return false;
+
+	    /* object specific verification checks here */
+
+	    return true;
+    }
+
+Write verifiers are very similar to the read verifiers, they just do things in
+the opposite order to the read verifiers. A typical write verifier::
+
+    static void
+    xfs_foo_write_verify(
+	    struct xfs_buf	*bp)
+    {
+	    struct xfs_mount	*mp = bp->b_mount;
+	    struct xfs_buf_log_item	*bip = bp->b_fspriv;
+
+	    if (!xfs_foo_verify(bp)) {
+		    XFS_CORRUPTION_ERROR(__func__, XFS_ERRLEVEL_LOW, mp, bp->b_addr);
+		    xfs_buf_ioerror(bp, EFSCORRUPTED);
+		    return;
+	    }
+
+	    if (!xfs_sb_version_hascrc(&mp->m_sb))
+		    return;
+
+
+	    if (bip) {
+		    struct xfs_ondisk_hdr	*hdr = bp->b_addr;
+		    hdr->lsn = cpu_to_be64(bip->bli_item.li_lsn);
+	    }
+	    xfs_update_cksum(bp->b_addr, BBTOB(bp->b_length), XFS_FOO_CRC_OFF);
+    }
+
+This will verify the internal structure of the metadata before we go any
+further, detecting corruptions that have occurred as the metadata has been
+modified in memory. If the metadata verifies OK, and CRCs are enabled, we then
+update the LSN field (when it was last modified) and calculate the CRC on the
+metadata. Once this is done, we can issue the IO.
+
+Inodes and Dquots
+=================
+
+Inodes and dquots are special snowflakes. They have per-object CRC and
+self-identifiers, but they are packed so that there are multiple objects per
+buffer. Hence we do not use per-buffer verifiers to do the work of per-object
+verification and CRC calculations. The per-buffer verifiers simply perform basic
+identification of the buffer - that they contain inodes or dquots, and that
+there are magic numbers in all the expected spots. All further CRC and
+verification checks are done when each inode is read from or written back to the
+buffer.
+
+The structure of the verifiers and the identifiers checks is very similar to the
+buffer code described above. The only difference is where they are called. For
+example, inode read verification is done in xfs_inode_from_disk() when the inode
+is first read out of the buffer and the struct xfs_inode is instantiated. The
+inode is already extensively verified during writeback in xfs_iflush_int, so the
+only addition here is to add the LSN and CRC to the inode as it is copied back
+into the buffer.
+
+XXX: inode unlinked list modification doesn't recalculate the inode CRC! None of
+the unlinked list modifications check or update CRCs, neither during unlink nor
+log recovery. So, it's gone unnoticed until now. This won't matter immediately -
+repair will probably complain about it - but it needs to be fixed.