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* xfs: add FITRIM supportChristoph Hellwig2011-01-127-12/+259
| | | | | | | | | | Allow manual discards from userspace using the FITRIM ioctl. This is not intended to be run during normal workloads, as the freepsace btree walks can cause large performance degradation. Signed-off-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Dave Chinner <dchinner@redhat.com> Signed-off-by: Alex Elder <aelder@sgi.com>
* xfs: ensure log covering transactions are synchronousDave Chinner2011-01-124-12/+13
| | | | | | | | | | | | | | | | | | To ensure the log is covered and the filesystem idles correctly, we need to ensure that dummy transactions hit the disk and do not stay pinned in memory. If the superblock is pinned in memory, it can't be flushed so the log covering cannot make progress. The result is dependent on timing - more oftent han not we continue to issues a log covering transaction every 36s rather than idling after ~90s. Fix this by making the log covering transaction synchronous. To avoid additional log force from xfssyncd, make the log covering transaction take the place of the existing log force in the xfssyncd background sync process. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de> Signed-off-by: Alex Elder <aelder@sgi.com>
* xfs: serialise unaligned direct IOsDave Chinner2011-01-111-10/+28
| | | | | | | | | | | | | | | | | | | | | | | | When two concurrent unaligned, non-overlapping direct IOs are issued to the same block, the direct Io layer will race to zero the block. The result is that one of the concurrent IOs will overwrite data written by the other IO with zeros. This is demonstrated by the xfsqa test 240. To avoid this problem, serialise all unaligned direct IOs to an inode with a big hammer. We need a big hammer approach as we need to serialise AIO as well, so we can't just block writes on locks. Hence, the big hammer is calling xfs_ioend_wait() while holding out other unaligned direct IOs from starting. We don't bother trying to serialised aligned vs unaligned IOs as they are overlapping IO and the result of concurrent overlapping IOs is undefined - the result of either IO is a valid result so we let them race. Hence we only penalise unaligned IO, which already has a major overhead compared to aligned IO so this isn't a major problem. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Alex Elder <aelder@sgi.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: factor common write setup codeDave Chinner2011-01-111-67/+56
| | | | | | | | | | The buffered IO and direct IO write paths share a common set of checks and limiting code prior to issuing the write. Factor that into a common helper function. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Alex Elder <aelder@sgi.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: split buffered IO write path from xfs_file_aio_writeDave Chinner2011-01-111-77/+69
| | | | | | | | | | | Complete the split of the different write IO paths by splitting the buffered IO write path out of xfs_file_aio_write(). This makes the different mechanisms of the write patchs easier to follow. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Alex Elder <aelder@sgi.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: split direct IO write path from xfs_file_aio_writeDave Chinner2011-01-111-63/+116
| | | | | | | | | | | | | | | The current xfs_file_aio_write code is a mess of locking shenanigans to handle the different locking requirements of buffered and direct IO. Start to clean this up by disentangling the direct IO path from the mess. This also removes the failed direct IO fallback path to buffered IO. XFS handles all direct IO cases without needing to fall back to buffered IO, so we can safely remove this unused path. This greatly simplifies the logic and locking needed in the write path. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: introduce xfs_rw_lock() helpers for locking the inodeDave Chinner2011-01-121-51/+80
| | | | | | | | | | | | | | | | We need to obtain the i_mutex, i_iolock and i_ilock during the read and write paths. Add a set of wrapper functions to neatly encapsulate the lock ordering and shared/exclusive semantics to make the locking easier to follow and get right. Note that this changes some of the exclusive locking serialisation in that serialisation will occur against the i_mutex instead of the XFS_IOLOCK_EXCL. This does not change any behaviour, and it is arguably more efficient to use the mutex for such serialisation than the rw_sem. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: factor post-write newsize updatesDave Chinner2011-01-111-22/+21
| | | | | | Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Alex Elder <aelder@sgi.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: factor common post-write isize handling codeDave Chinner2011-01-111-26/+28
| | | | | | Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Alex Elder <aelder@sgi.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: ensure sync write errors are returnedDave Chinner2011-01-111-26/+23
| | | | | | | | | | | | | xfs_file_aio_write() only returns the error from synchronous flushing of the data and inode if error == 0. At the point where error is being checked, it is guaranteed to be > 0. Therefore any errors returned by the data or fsync flush will never be returned. Fix the checks so we overwrite the current error once and only if an error really occurred. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Alex Elder <aelder@sgi.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: convert grant head manipulations to lockless algorithmDave Chinner2010-12-212-77/+49
| | | | | | | | | | | | | | | | | | | | | | | The only thing that the grant lock remains to protect is the grant head manipulations when adding or removing space from the log. These calculations are already based on atomic variables, so we can already update them safely without locks. However, the grant head manpulations require atomic multi-step calculations to be executed, which the algorithms currently don't allow. To make these multi-step calculations atomic, convert the algorithms to compare-and-exchange loops on the atomic variables. That is, we sample the old value, perform the calculation and use atomic64_cmpxchg() to attempt to update the head with the new value. If the head has not changed since we sampled it, it will succeed and we are done. Otherwise, we rerun the calculation again from a new sample of the head. This allows us to remove the grant lock from around all the grant head space manipulations, and that effectively removes the grant lock from the log completely. Hence we can remove the grant lock completely from the log at this point. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: introduce new locks for the log grant ticket wait queuesDave Chinner2010-12-213-60/+97
| | | | | | | | | | | | | | | | | | | | | | | The log grant ticket wait queues are currently protected by the log grant lock. However, the queues are functionally independent from each other, and operations on them only require serialisation against other queue operations now that all of the other log variables they use are atomic values. Hence, we can make them independent of the grant lock by introducing new locks just to protect the lists operations. because the lists are independent, we can use a lock per list and ensure that reserve and write head queuing do not contend. To ensure forced shutdowns work correctly in conjunction with the new fast paths, ensure that we check whether the log has been shut down in the grant functions once we hold the relevant spin locks but before we go to sleep. This is needed to co-ordinate correctly with the wakeups that are issued on the ticket queues so we don't leave any processes sleeping on the queues during a shutdown. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: convert log grant heads to atomic variablesDave Chinner2010-12-032-10/+10
| | | | | | | | | | Convert the log grant heads to atomic64_t types in preparation for converting the accounting algorithms to atomic operations. his patch just converts the variables; the algorithmic changes are in a separate patch for clarity. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: convert l_tail_lsn to an atomic variable.Dave Chinner2010-12-214-46/+63
| | | | | | | | | | | | | | | | | | | log->l_tail_lsn is currently protected by the log grant lock. The lock is only needed for serialising readers against writers, so we don't really need the lock if we make the l_tail_lsn variable an atomic. Converting the l_tail_lsn variable to an atomic64_t means we can start to peel back the grant lock from various operations. Also, provide functions to safely crack an atomic LSN variable into it's component pieces and to recombined the components into an atomic variable. Use them where appropriate. This also removes the need for explicitly holding a spinlock to read the l_tail_lsn on 32 bit platforms. Signed-off-by: Dave Chinner <dchinner@redhat.com>
* xfs: convert l_last_sync_lsn to an atomic variableDave Chinner2010-12-033-34/+36
| | | | | | | | | | | | | | | log->l_last_sync_lsn is updated in only one critical spot - log buffer Io completion - and is protected by the grant lock here. This requires the grant lock to be taken for every log buffer IO completion. Converting the l_last_sync_lsn variable to an atomic64_t means that we do not need to take the grant lock in log buffer IO completion to update it. This also removes the need for explicitly holding a spinlock to read the l_last_sync_lsn on 32 bit platforms. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: make AIL tail pushing independent of the grant lockDave Chinner2010-12-211-57/+54
| | | | | | | | | | | | | | | | | | | | | The xlog_grant_push_ail() currently takes the grant lock internally to sample the tail lsn, last sync lsn and the reserve grant head. Most of the callers already hold the grant lock but have to drop it before calling xlog_grant_push_ail(). This is a left over from when the AIL tail pushing was done in line and hence xlog_grant_push_ail had to drop the grant lock. AIL push is now done in another thread and hence we can safely hold the grant lock over the entire xlog_grant_push_ail call. Push the grant lock outside of xlog_grant_push_ail() to simplify the locking and synchronisation needed for tail pushing. This will reduce traffic on the grant lock by itself, but this is only one step in preparing for the complete removal of the grant lock. While there, clean up the formatting of xlog_grant_push_ail() to match the rest of the XFS code. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: use wait queues directly for the log wait queuesDave Chinner2010-12-216-106/+52
| | | | | | | | | | | | | | The log grant queues are one of the few places left using sv_t constructs for waiting. Given we are touching this code, we should convert them to plain wait queues. While there, convert all the other sv_t users in the log code as well. Seeing as this removes the last users of the sv_t type, remove the header file defining the wrapper and the fragments that still reference it. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: combine grant heads into a single 64 bit integerDave Chinner2010-12-214-91/+119
| | | | | | | | | | | Prepare for switching the grant heads to atomic variables by combining the two 32 bit values that make up the grant head into a single 64 bit variable. Provide wrapper functions to combine and split the grant heads appropriately for calculations and use them as necessary. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: rework log grant space calculationsDave Chinner2010-12-211-47/+48
| | | | | | | | | | | | The log grant space calculations are repeated for both write and reserve grant heads. To make it simpler to convert the calculations toa different algorithm, factor them so both the gratn heads use the same calculation functions. Once this is done we can drop the wrappers that are used in only a couple of place to update both grant heads at once as they don't provide any particular value. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: fact out common grant head/log tail verification codeDave Chinner2010-12-211-29/+22
| | | | | | | | | Factor repeated debug code out of grant head manipulation functions into a separate function. This removes ifdef DEBUG spagetti from the code and makes the code easier to follow. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: convert log grant ticket queues to list headsDave Chinner2010-12-213-97/+53
| | | | | | | | | | | The grant write and reserve queues use a roll-your-own double linked list, so convert it to a standard list_head structure and convert all the list traversals to use list_for_each_entry(). We can also get rid of the XLOG_TIC_IN_Q flag as we can use the list_empty() check to tell if the ticket is in a list or not. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: use AIL bulk delete function to implement single deleteDave Chinner2010-12-202-72/+11
| | | | | | | | | | | We now have two copies of AIL delete operations that are mostly duplicate functionality. The single log item deletes can be implemented via the bulk updates by turning xfs_trans_ail_delete() into a simple wrapper. This removes all the duplicate delete functionality and associated helpers. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: use AIL bulk update function to implement single updatesDave Chinner2010-12-203-96/+13
| | | | | | | | | | | We now have two copies of AIL insert operations that are mostly duplicate functionality. The single log item updates can be implemented via the bulk updates by turning xfs_trans_ail_update() into a simple wrapper. This removes all the duplicate insert functionality and associated helpers. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: remove all the inodes on a buffer from the AIL in bulkDave Chinner2010-12-203-16/+151
| | | | | | | | | | | | | | | | | | | When inode buffer IO completes, usually all of the inodes are removed from the AIL. This involves processing them one at a time and taking the AIL lock once for every inode. When all CPUs are processing inode IO completions, this causes excessive amount sof contention on the AIL lock. Instead, change the way we process inode IO completion in the buffer IO done callback. Allow the inode IO done callback to walk the list of IO done callbacks and pull all the inodes off the buffer in one go and then process them as a batch. Once all the inodes for removal are collected, take the AIL lock once and do a bulk removal operation to minimise traffic on the AIL lock. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: consume iodone callback items on buffers as they are processedDave Chinner2010-12-031-11/+21
| | | | | | | | | | | | | | | | | To allow buffer iodone callbacks to consume multiple items off the callback list, first we need to convert the xfs_buf_do_callbacks() to consume items and always pull the next item from the head of the list. The means the item list walk is never dependent on knowing the next item on the list and hence allows callbacks to remove items from the list as well. This allows callbacks to do bulk operations by scanning the list for identical callbacks, consuming them all and then processing them in bulk, negating the need for multiple callbacks of that type. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: reduce the number of AIL push wakeupsDave Chinner2010-12-171-4/+16
| | | | | | | | | | | | | | The xfaild often tries to rest to wait for congestion to pass of for IO to complete, but is regularly woken in tail-pushing situations. In severe cases, the xfsaild is getting woken tens of thousands of times a second. Reduce the number needless wakeups by only waking the xfsaild if the new target is larger than the old one. Further make short sleeps uninterruptible as they occur when the xfsaild has decided it needs to back off to allow some IO to complete and being woken early is counter-productive. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: bulk AIL insertion during transaction commitDave Chinner2010-12-204-12/+195
| | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | When inserting items into the AIL from the transaction committed callbacks, we take the AIL lock for every single item that is to be inserted. For a CIL checkpoint commit, this can be tens of thousands of individual inserts, yet almost all of the items will be inserted at the same point in the AIL because they have the same index. To reduce the overhead and contention on the AIL lock for such operations, introduce a "bulk insert" operation which allows a list of log items with the same LSN to be inserted in a single operation via a list splice. To do this, we need to pre-sort the log items being committed into a temporary list for insertion. The complexity is that not every log item will end up with the same LSN, and not every item is actually inserted into the AIL. Items that don't match the commit LSN will be inserted and unpinned as per the current one-at-a-time method (relatively rare), while items that are not to be inserted will be unpinned and freed immediately. Items that are to be inserted at the given commit lsn are placed in a temporary array and inserted into the AIL in bulk each time the array fills up. As a result of this, we trade off AIL hold time for a significant reduction in traffic. lock_stat output shows that the worst case hold time is unchanged, but contention from AIL inserts drops by an order of magnitude and the number of lock traversal decreases significantly. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: clean up xfs_ail_delete()Dave Chinner2010-12-031-20/+7
| | | | | | | | | | | | | | | | xfs_ail_delete() has a needlessly complex interface. It returns the log item that was passed in for deletion (which the callers then assert is identical to the one passed in), and callers of xfs_ail_delete() still need to invalidate current traversal cursors. Make xfs_ail_delete() return void, move the cursor invalidation inside it, and clean up the callers just to use the log item pointer they passed in. While cleaning up, remove the messy and unnecessary "/* ARGUSED */" comments around all these functions. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: Pull EFI/EFD handling out from under the AIL lockDave Chinner2010-12-204-49/+59
| | | | | | | | | | | | | | | | | EFI/EFD interactions are protected from races by the AIL lock. They are the only type of log items that require the the AIL lock to serialise internal state, so they need to be separated from the AIL lock before we can do bulk insert operations on the AIL. To acheive this, convert the counter of the number of extents in the EFI to an atomic so it can be safely manipulated by EFD processing without locks. Also, convert the EFI state flag manipulations to use atomic bit operations so no locks are needed to record state changes. Finally, use the state bits to determine when it is safe to free the EFI and clean up the code to do this neatly. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: fix EFI transaction cancellation.Dave Chinner2010-12-202-12/+9
| | | | | | | | | | | | | | | | | | XFS_EFI_CANCELED has not been set in the code base since xfs_efi_cancel() was removed back in 2006 by commit 065d312e15902976d256ddaf396a7950ec0350a8 ("[XFS] Remove unused iop_abort log item operation), and even then xfs_efi_cancel() was never called. I haven't tracked it back further than that (beyond git history), but it indicates that the handling of EFIs in cancelled transactions has been broken for a long time. Basically, when we get an IOP_UNPIN(lip, 1); call from xfs_trans_uncommit() (i.e. remove == 1), if we don't free the log item descriptor we leak it. Fix the behviour to be correct and kill the XFS_EFI_CANCELED flag. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: connect up buffer reclaim priority hooksDave Chinner2010-12-024-9/+14
| | | | | | | | | | | Now that the buffer reclaim infrastructure can handle different reclaim priorities for different types of buffers, reconnect the hooks in the XFS code that has been sitting dormant since it was ported to Linux. This should finally give use reclaim prioritisation that is on a par with the functionality that Irix provided XFS 15 years ago. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: add a lru to the XFS buffer cacheDave Chinner2010-12-022-22/+150
| | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | Introduce a per-buftarg LRU for memory reclaim to operate on. This is the last piece we need to put in place so that we can fully control the buffer lifecycle. This allows XFS to be responsibile for maintaining the working set of buffers under memory pressure instead of relying on the VM reclaim not to take pages we need out from underneath us. The implementation introduces a b_lru_ref counter into the buffer. This is currently set to 1 whenever the buffer is referenced and so is used to determine if the buffer should be added to the LRU or not when freed. Effectively it allows lazy LRU initialisation of the buffer so we do not need to touch the LRU list and locks in xfs_buf_find(). Instead, when the buffer is being released and we drop the last reference to it, we check the b_lru_ref count and if it is none zero we re-add the buffer reference and add the inode to the LRU. The b_lru_ref counter is decremented by the shrinker, and whenever the shrinker comes across a buffer with a zero b_lru_ref counter, if released the LRU reference on the buffer. In the absence of a lookup race, this will result in the buffer being freed. This counting mechanism is used instead of a reference flag so that it is simple to re-introduce buffer-type specific reclaim reference counts to prioritise reclaim more effectively. We still have all those hooks in the XFS code, so this will provide the infrastructure to re-implement that functionality. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: convert xfsbud shrinker to a per-buftarg shrinker.Dave Chinner2010-11-302-66/+27
| | | | | | | | | | | | | | Before we introduce per-buftarg LRU lists, split the shrinker implementation into per-buftarg shrinker callbacks. At the moment we wake all the xfsbufds to run the delayed write queues to free the dirty buffers and make their pages available for reclaim. However, with an LRU, we want to be able to free clean, unused buffers as well, so we need to separate the xfsbufd from the shrinker callbacks. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Alex Elder <aelder@sgi.com>
* xfs: convert pag_ici_lock to a spin lockDave Chinner2010-12-164-11/+11
| | | | | | | | | | | | now that we are using RCU protection for the inode cache lookups, the lock is only needed on the modification side. Hence it is not necessary for the lock to be a rwlock as there are no read side holders anymore. Convert it to a spin lock to reflect it's exclusive nature. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Alex Elder <aelder@sgi.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: convert inode cache lookups to use RCU lockingDave Chinner2010-12-173-42/+141
| | | | | | | | | | | | | | | | | | | | | With delayed logging greatly increasing the sustained parallelism of inode operations, the inode cache locking is showing significant read vs write contention when inode reclaim runs at the same time as lookups. There is also a lot more write lock acquistions than there are read locks (4:1 ratio) so the read locking is not really buying us much in the way of parallelism. To avoid the read vs write contention, change the cache to use RCU locking on the read side. To avoid needing to RCU free every single inode, use the built in slab RCU freeing mechanism. This requires us to be able to detect lookups of freed inodes, so enѕure that ever freed inode has an inode number of zero and the XFS_IRECLAIM flag set. We already check the XFS_IRECLAIM flag in cache hit lookup path, but also add a check for a zero inode number as well. We canthen convert all the read locking lockups to use RCU read side locking and hence remove all read side locking. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Alex Elder <aelder@sgi.com>
* xfs: rcu free inodesDave Chinner2010-12-161-1/+13
| | | | | | | | | | | | | | | | | | | Introduce RCU freeing of XFS inodes so that we can convert lookup traversals to use rcu_read_lock() protection. This patch only introduces the RCU freeing to minimise the potential conflicts with mainline if this is merged into mainline via a VFS patchset. It abuses the i_dentry list for the RCU callback structure because the VFS patches make this a union so it is safe to use like this and simplifies and merge issues. This patch uses basic RCU freeing rather than SLAB_DESTROY_BY_RCU. The later lookup patches need the same "found free inode" protection regardless of the RCU freeing method used, so once again the RCU freeing method can be dealt with apprpriately at merge time without affecting any other code. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Paul E. McKenney <paulmck@linux.vnet.ibm.com>
* xfs: don't truncate prealloc from frequently accessed inodesDave Chinner2010-12-232-27/+47
| | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | A long standing problem for streaming writeѕ through the NFS server has been that the NFS server opens and closes file descriptors on an inode for every write. The result of this behaviour is that the ->release() function is called on every close and that results in XFS truncating speculative preallocation beyond the EOF. This has an adverse effect on file layout when multiple files are being written at the same time - they interleave their extents and can result in severe fragmentation. To avoid this problem, keep track of ->release calls made on a dirty inode. For most cases, an inode is only going to be opened once for writing and then closed again during it's lifetime in cache. Hence if there are multiple ->release calls when the inode is dirty, there is a good chance that the inode is being accessed by the NFS server. Hence set a flag the first time ->release is called while there are delalloc blocks still outstanding on the inode. If this flag is set when ->release is next called, then do no truncate away the speculative preallocation - leave it there so that subsequent writes do not need to reallocate the delalloc space. This will prevent interleaving of extents of different inodes written concurrently to the same AG. If we get this wrong, it is not a big deal as we truncate speculative allocation beyond EOF anyway in xfs_inactive() when the inode is thrown out of the cache. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: dynamic speculative EOF preallocationDave Chinner2011-01-044-10/+110
| | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | | Currently the size of the speculative preallocation during delayed allocation is fixed by either the allocsize mount option of a default size. We are seeing a lot of cases where we need to recommend using the allocsize mount option to prevent fragmentation when buffered writes land in the same AG. Rather than using a fixed preallocation size by default (up to 64k), make it dynamic by basing it on the current inode size. That way the EOF preallocation will increase as the file size increases. Hence for streaming writes we are much more likely to get large preallocations exactly when we need it to reduce fragementation. For default settings, the size of the initial extents is determined by the number of parallel writers and the amount of memory in the machine. For 4GB RAM and 4 concurrent 32GB file writes: EXT: FILE-OFFSET BLOCK-RANGE AG AG-OFFSET TOTAL 0: [0..1048575]: 1048672..2097247 0 (1048672..2097247) 1048576 1: [1048576..2097151]: 5242976..6291551 0 (5242976..6291551) 1048576 2: [2097152..4194303]: 12583008..14680159 0 (12583008..14680159) 2097152 3: [4194304..8388607]: 25165920..29360223 0 (25165920..29360223) 4194304 4: [8388608..16777215]: 58720352..67108959 0 (58720352..67108959) 8388608 5: [16777216..33554423]: 117440584..134217791 0 (117440584..134217791) 16777208 6: [33554424..50331511]: 184549056..201326143 0 (184549056..201326143) 16777088 7: [50331512..67108599]: 251657408..268434495 0 (251657408..268434495) 16777088 and for 16 concurrent 16GB file writes: EXT: FILE-OFFSET BLOCK-RANGE AG AG-OFFSET TOTAL 0: [0..262143]: 2490472..2752615 0 (2490472..2752615) 262144 1: [262144..524287]: 6291560..6553703 0 (6291560..6553703) 262144 2: [524288..1048575]: 13631592..14155879 0 (13631592..14155879) 524288 3: [1048576..2097151]: 30408808..31457383 0 (30408808..31457383) 1048576 4: [2097152..4194303]: 52428904..54526055 0 (52428904..54526055) 2097152 5: [4194304..8388607]: 104857704..109052007 0 (104857704..109052007) 4194304 6: [8388608..16777215]: 209715304..218103911 0 (209715304..218103911) 8388608 7: [16777216..33554423]: 452984848..469762055 0 (452984848..469762055) 16777208 Because it is hard to take back specualtive preallocation, cases where there are large slow growing log files on a nearly full filesystem may cause premature ENOSPC. Hence as the filesystem nears full, the maximum dynamic prealloc size іs reduced according to this table (based on 4k block size): freespace max prealloc size >5% full extent (8GB) 4-5% 2GB (8GB >> 2) 3-4% 1GB (8GB >> 3) 2-3% 512MB (8GB >> 4) 1-2% 256MB (8GB >> 5) <1% 128MB (8GB >> 6) This should reduce the amount of space held in speculative preallocation for such cases. The allocsize mount option turns off the dynamic behaviour and fixes the prealloc size to whatever the mount option specifies. i.e. the behaviour is unchanged. Signed-off-by: Dave Chinner <dchinner@redhat.com>
* xfs: use KM_NOFS for allocations during attribute list operationsDave Chinner2010-12-231-2/+2
| | | | | | | | | | | | | When listing attributes, we are doiing memory allocations under the inode ilock using only KM_SLEEP. This allows memory allocation to recurse back into the filesystem and do writeback, which may the ilock we already hold on the current inode. THis will deadlock. Hence use KM_NOFS for such allocations outside of transaction context to ensure that reclaim recursion does not occur. Reported-by: Nick Piggin <npiggin@gmail.com> Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: provide a inode iolock lockdep classDave Chinner2010-12-233-0/+23
| | | | | | | | | | | | | | | | | | | | The XFS iolock needs to be re-initialised to a new lock class before it enters reclaim to prevent lockdep false positives. Unfortunately, this is not sufficient protection as inodes in the XFS_IRECLAIMABLE state can be recycled and not re-initialised before being reused. We need to re-initialise the lock state when transfering out of XFS_IRECLAIMABLE state to XFS_INEW, but we need to keep the same class as if the inode was just allocated. Hence we need a specific lockdep class variable for the iolock so that both initialisations use the same class. While there, add a specific class for inodes in the reclaim state so that it is easy to tell from lockdep reports what state the inode was in that generated the report. Signed-off-by: Dave Chinner <dchinner@redhat.com> Reviewed-by: Christoph Hellwig <hch@lst.de>
* xfs: factor duplicate code in xfs_alloc_ag_vextent_near into a helperChristoph Hellwig2010-12-161-180/+113
| | | | | | | | | | | | Add a new xfs_alloc_find_best_extent that does a forward/backward search in the allocation btree. That code previously was existed two times in xfs_alloc_ag_vextent_near, once for each search direction. Based on an earlier patch from Dave Chinner. Signed-off-by: Christoph Hellwig <hch@lst.de> Signed-off-by: Alex Elder <aelder@sgi.com>
* xfs: clean up xfs_alloc_ag_vextent_exactChristoph Hellwig2010-12-162-26/+33
| | | | | | | | | | Use a goto label to consolidate all block not found cases, and add a tracepoint for them. Also clean up a few whitespace issues. Based on an earlier patch from Dave Chinner. Signed-off-by: Christoph Hellwig <hch@lst.de> Signed-off-by: Alex Elder <aelder@sgi.com>
* xfs: simplify xfs_map_at_offsetChristoph Hellwig2010-12-161-11/+4
| | | | | | | | | | | | Move the buffer locking into the callers as they need to do it wether they call xfs_map_at_offset or not. Remove the b_bdev assignment, which is already done by get_blocks. Remove the duplicate extent type asserts in xfs_convert_page just before calling xfs_map_at_offset. Signed-off-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Dave Chinner <dchinner@redhat.com> Signed-off-by: Alex Elder <aelder@sgi.com>
* xfs: refactor xfs_vm_writepageChristoph Hellwig2010-12-161-58/+39
| | | | | | | | | | | | | | | | | After the last patches the code for overwrites is the same as for delayed and unwritten extents except that it doesn't need to call xfs_map_at_offset. Take care of that fact to simplify xfs_vm_writepage. The buffer loop now first checks the type of buffer and checks/sets the ioend type, or continues to the next buffer if it's not interesting to us. Only after that we validate the iomap and perform the block mapping if needed, all in common code for the cases where we have to do work. Signed-off-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Dave Chinner <dchinner@redhat.com> Signed-off-by: Alex Elder <aelder@sgi.com>
* xfs: remove the all_bh flag from xfs_convert_pageChristoph Hellwig2010-12-161-20/+13
| | | | | | | | | | The all_bh flag is always set when entering the page clustering machinery with a regular written extent, which means the check for it is superflous. Signed-off-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Dave Chinner <dchinner@redhat.com> Signed-off-by: Alex Elder <aelder@sgi.com>
* xfs: remove xfs_probe_clusterChristoph Hellwig2010-12-161-107/+4
| | | | | | | | | | | | xfs_map_blocks always calls xfs_bmapi with the XFS_BMAPI_ENTIRE entire flag, which tells it to not cap the extent at the passed in size, but just treat the size as an minimum to map. This means xfs_probe_cluster is entirely useless as we'll always get the whole extent back anyway. Signed-off-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Dave Chinner <dchinner@redhat.com> Signed-off-by: Alex Elder <aelder@sgi.com>
* xfs: simplify xfs_map_blocksChristoph Hellwig2010-12-161-52/+25
| | | | | | | | | | | | | | | | No need to lock the extent map exclusive when performing an overwrite, we know the extent map must already have been loaded by get_blocks. Apply the non-blocking inode semantics to all mapping types instead of just delayed allocations. Remove the handling of not yet allocated blocks for the IO_UNWRITTEN case - if an extent is marked as unwritten allocated in the buffer it must already have an extent on disk. Add asserts to verify all the assumptions above in debug builds. Signed-off-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Dave Chinner <dchinner@redhat.com> Signed-off-by: Alex Elder <aelder@sgi.com>
* xfs: kill xfs_iomapChristoph Hellwig2010-12-165-213/+191
| | | | | | | | | | | | | | | | | Opencode the xfs_iomap code in it's two callers. The overlap of passed flags already was minimal and will be further reduced in the next patch. As a side effect the BMAPI_* flags for xfs_bmapi and the IO_* flags for I/O end processing are merged into a single set of flags, which should be a bit more descriptive of the operation we perform. Also improve the tracing by giving each caller it's own type set of tracepoints. Signed-off-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Dave Chinner <dchinner@redhat.com> Signed-off-by: Alex Elder <aelder@sgi.com>
* xfs: cleanup the xfs_iomap_write_* helpersChristoph Hellwig2010-12-161-30/+15
| | | | | | | | | | | Remove passing the BMAPI_* flags to these helpers, in xfs_iomap_write_direct the check BMAPI_DIRECT was always true, and in the xfs_iomap_write_delay path is was never checked at all. Remove the nmap return value as we never make use of it. Signed-off-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Dave Chinner <dchinner@redhat.com> Signed-off-by: Alex Elder <aelder@sgi.com>
* xfs: a few small tweaks for overwrites in xfs_vm_writepageChristoph Hellwig2010-12-161-8/+7
| | | | | | | | | | | | | | | | | | | | Don't trylock the buffer. We are the only one ever locking it for a regular file address space, and trylock was only copied from the generic code which did it due to the old buffer based writeout in jbd. Also make sure to only write out the buffer if the iomap actually is valid, because we wouldn't have a proper mapping otherwise. In practice we will never get an invalid mapping here as the page lock guarantees truncate doesn't race with us, but better be safe than sorry. Also make sure we allocate a new ioend when crossing boundaries between mappings, just like we do for delalloc and unwritten extents. Again this currently doesn't matter as the I/O end handler only cares for the boundaries for unwritten extents, but this makes the code fully correct and the same as for delalloc/unwritten extents. Signed-off-by: Christoph Hellwig <hch@lst.de> Reviewed-by: Dave Chinner <dchinner@redhat.com> Signed-off-by: Alex Elder <aelder@sgi.com>
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