- 26 Sep, 2022 40 commits
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Filipe Manana authored
One of the most expensive tasks performed during fiemap is to check if an extent is shared. This task has two major steps: 1) Check if the data extent is shared. This implies checking the extent item in the extent tree, checking delayed references, etc. If we find the data extent is directly shared, we terminate immediately; 2) If the data extent is not directly shared (its extent item has a refcount of 1), then it may be shared if we have snapshots that share subtrees of the inode's subvolume b+tree. So we check if the leaf containing the file extent item is shared, then its parent node, then the parent node of the parent node, etc, until we reach the root node or we find one of them is shared - in which case we stop immediately. During fiemap we process the extents of a file from left to right, from file offset 0 to EOF. This means that we iterate b+tree leaves from left to right, and has the implication that we keep repeating that second step above several times for the same b+tree path of the inode's subvolume b+tree. For example, if we have two file extent items in leaf X, and the path to leaf X is A -> B -> C -> X, then when we try to determine if the data extent referenced by the first extent item is shared, we check if the data extent is shared - if it's not, then we check if leaf X is shared, if not, then we check if node C is shared, if not, then check if node B is shared, if not than check if node A is shared. When we move to the next file extent item, after determining the data extent is not shared, we repeat the checks for X, C, B and A - doing all the expensive searches in the extent tree, delayed refs, etc. If we have thousands of tile extents, then we keep repeating the sharedness checks for the same paths over and over. On a file that has no shared extents or only a small portion, it's easy to see that this scales terribly with the number of extents in the file and the sizes of the extent and subvolume b+trees. This change eliminates the repeated sharedness check on extent buffers by caching the results of the last path used. The results can be used as long as no snapshots were created since they were cached (for not shared extent buffers) or no roots were dropped since they were cached (for shared extent buffers). This greatly reduces the time spent by fiemap for files with thousands of extents and/or large extent and subvolume b+trees. Example performance test: $ cat fiemap-perf-test.sh #!/bin/bash DEV=/dev/sdi MNT=/mnt/sdi mkfs.btrfs -f $DEV mount -o compress=lzo $DEV $MNT # 40G gives 327680 128K file extents (due to compression). xfs_io -f -c "pwrite -S 0xab -b 1M 0 40G" $MNT/foobar umount $MNT mount -o compress=lzo $DEV $MNT start=$(date +%s%N) filefrag $MNT/foobar end=$(date +%s%N) dur=$(( (end - start) / 1000000 )) echo "fiemap took $dur milliseconds (metadata not cached)" start=$(date +%s%N) filefrag $MNT/foobar end=$(date +%s%N) dur=$(( (end - start) / 1000000 )) echo "fiemap took $dur milliseconds (metadata cached)" umount $MNT Before this patch: $ ./fiemap-perf-test.sh (...) /mnt/sdi/foobar: 327680 extents found fiemap took 3597 milliseconds (metadata not cached) /mnt/sdi/foobar: 327680 extents found fiemap took 2107 milliseconds (metadata cached) After this patch: $ ./fiemap-perf-test.sh (...) /mnt/sdi/foobar: 327680 extents found fiemap took 1646 milliseconds (metadata not cached) /mnt/sdi/foobar: 327680 extents found fiemap took 698 milliseconds (metadata cached) That's about 2.2x faster when no metadata is cached, and about 3x faster when all metadata is cached. On a real filesystem with many other files, data, directories, etc, the b+trees will be 2 or 3 levels higher, therefore this optimization will have a higher impact. Several reports of a slow fiemap show up often, the two Link tags below refer to two recent reports of such slowness. This patch, together with the next ones in the series, is meant to address that. Link: https://lore.kernel.org/linux-btrfs/21dd32c6-f1f9-f44a-466a-e18fdc6788a7@virtuozzo.com/ Link: https://lore.kernel.org/linux-btrfs/Ysace25wh5BbLd5f@atmark-techno.com/Reviewed-by:Josef Bacik <josef@toxicpanda.com> Signed-off-by:
Filipe Manana <fdmanana@suse.com> Signed-off-by:
David Sterba <dsterba@suse.com>
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Filipe Manana authored
The function btrfs_check_shared() is supposed to be used to check if a data extent is shared, but its name is too generic, may easily cause confusion in the sense that it may be used for metadata extents. So rename it to btrfs_is_data_extent_shared(), which will also make it less confusing after the next change that adds a backref lookup cache for the b+tree nodes that lead to the leaf that contains the file extent item that points to the target data extent. Reviewed-by:
Qu Wenruo <wqu@suse.com> Signed-off-by:
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Filipe Manana authored
Doing fiemap on a file with a very large number of extents can take a very long time, and we have reports of it being too slow (two recent examples in the Link tags below), so make it interruptible. Link: https://lore.kernel.org/linux-btrfs/21dd32c6-f1f9-f44a-466a-e18fdc6788a7@virtuozzo.com/ Link: https://lore.kernel.org/linux-btrfs/Ysace25wh5BbLd5f@atmark-techno.com/Reviewed-by:
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Filipe Manana authored
If the flag FIEMAP_FLAG_SYNC is passed to fiemap, it means all delalloc should be flushed and writeback complete. We call the generic helper fiemap_prep() which does a filemap_write_and_wait() in case that flag is given, however that is not enough if we have compression. Because a single filemap_fdatawrite_range() only starts compression (in an async thread) and therefore returns before the compression is done and writeback is started. So make btrfs_fiemap(), actually wait for all writeback to start and complete if FIEMAP_FLAG_SYNC is set. We start and wait for writeback on the whole possible file range, from 0 to LLONG_MAX, because that is what the generic code at fiemap_prep() does. Reviewed-by:
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Filipe Manana authored
There's no point to check for a 0 length at extent_fiemap(), as before calling it, we called fiemap_prep() at btrfs_fiemap(), which already checks for a zero length and returns the same -EINVAL error. So remove the pointless check. Reviewed-by:
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Filipe Manana authored
During fiemap we are testing if an extent map has a block start with a value of EXTENT_MAP_LAST_BYTE, but that is never set on an extent map, and never was according to git history. So remove that useless check. Reviewed-by:
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Filipe Manana authored
The current implementation of hole and data seeking for llseek does not scale well in regards to the number of extents and the distance between the start offset and the next hole or extent. This is due to a very high algorithmic complexity. Often we also get reports of btrfs' hole and data seeking (llseek) being too slow, such as at 2017's LSFMM (see the Link tag at the bottom). In order to better understand it, lets consider the case where the start offset is 0, we are seeking for a hole and the file size is 16G. Between file offset 0 and the first hole in the file there are 100K extents - this is common for large files, specially if we have compression enabled, since the maximum extent size is limited to 128K. The steps take by the main loop of the current algorithm are the following: 1) We start by calling btrfs_get_extent_fiemap(), for file offset 0, which calls btrfs_get_extent(). This will first lookup for an extent map in the inode's extent map tree (a red black tree). If the extent map is not loaded in memory, then it will do a lookup for the corresponding file extent item in the subvolume's b+tree, create an extent map based on the contents of the file extent item and then add the extent map to the extent map tree of the inode; 2) The second iteration calls btrfs_get_extent_fiemap() again, this time with a start offset matching the end offset of the previous extent. Again, btrfs_get_extent() will first search the extent map tree, and if it doesn't find an extent map there, it will again search in the b+tree of the subvolume for a matching file extent item, build an extent map based on the file extent item, and add the extent map to to the extent map tree of the inode; 3) This repeats over and over until we find the first hole (when seeking for holes) or until we find the first extent (when seeking for data). If there no extent maps loaded in memory for each iteration, then on each iteration we do 1 extent map tree search, 1 b+tree search, plus 1 more extent map tree traversal to insert an extent map - plus we allocate memory for the extent map. On each iteration we are growing the size of the extent map tree, making each future search slower, and also visiting the same b+tree leaves over and over again - taking into account with the default leaf size of 16K we can fit more than 200 file extent items in a leaf - so we can visit the same b+tree leaf 200+ times, on each visit walking down a path from the root to the leaf. So it's easy to see that what we have now doesn't scale well. Also, it loads an extent map for every file extent item into memory, which is not efficient - we should add extents maps only when doing IO (writing or reading file data). This change implements a new algorithm which scales much better, and works like this: 1) We iterate over the subvolume's b+tree, visiting each leaf that has file extent items once and only once; 2) For any file extent items found, that don't represent holes or prealloc extents, it will not search the extent map tree - there's no need at all for that - an extent map is just an in-memory representation of a file extent item; 3) When a hole is found, or a prealloc extent, it will check if there's delalloc for its range. For this it will search for EXTENT_DELALLOC bits in the inode's io tree and check the extent map tree - this is for accounting for unflushed delalloc and for flushed delalloc (the period between running delalloc and ordered extent completion), respectively. This is similar to what the current implementation does when it finds a hole or prealloc extent, but without creating extent maps and adding them to the extent map tree in case they are not loaded in memory; 4) It never allocates extent maps, or adds extent maps to the inode's extent map tree. This not only saves memory and time (from the tree insertions and allocations), but also eliminates the possibility of -ENOMEM due to allocating too many extent maps. Part of this new code will also be used later for fiemap (which also suffers similar scalability problems). The following test example can be used to quickly measure the efficiency before and after this patch: $ cat test-seek-hole.sh #!/bin/bash DEV=/dev/sdi MNT=/mnt/sdi mkfs.btrfs -f $DEV mount -o compress=lzo $DEV $MNT # 16G file -> 131073 compressed extents. xfs_io -f -c "pwrite -S 0xab -b 1M 0 16G" $MNT/foobar # Leave a 1M hole at file offset 15G. xfs_io -c "fpunch 15G 1M" $MNT/foobar # Unmount and mount again, so that we can test when there's no # metadata cached in memory. umount $MNT mount -o compress=lzo $DEV $MNT # Test seeking for hole from offset 0 (hole is at offset 15G). start=$(date +%s%N) xfs_io -c "seek -h 0" $MNT/foobar end=$(date +%s%N) dur=$(( (end - start) / 1000000 )) echo "Took $dur milliseconds to seek first hole (metadata not cached)" echo start=$(date +%s%N) xfs_io -c "seek -h 0" $MNT/foobar end=$(date +%s%N) dur=$(( (end - start) / 1000000 )) echo "Took $dur milliseconds to seek first hole (metadata cached)" echo umount $MNT Before this change: $ ./test-seek-hole.sh (...) Whence Result HOLE 16106127360 Took 176 milliseconds to seek first hole (metadata not cached) Whence Result HOLE 16106127360 Took 17 milliseconds to seek first hole (metadata cached) After this change: $ ./test-seek-hole.sh (...) Whence Result HOLE 16106127360 Took 43 milliseconds to seek first hole (metadata not cached) Whence Result HOLE 16106127360 Took 13 milliseconds to seek first hole (metadata cached) That's about 4x faster when no metadata is cached and about 30% faster when all metadata is cached. In practice the differences may often be significantly higher, either due to a higher number of extents in a file or because the subvolume's b+tree is much bigger than in this example, where we only have one file. Link: https://lwn.net/Articles/718805/Reviewed-by: -
Filipe Manana authored
Doing hole or data seeking on a file with a very large number of extents can take a long time, and we have reports of it being too slow (such as at LSFMM from 2017, see the Link below). So make it interruptible. Link: https://lwn.net/Articles/718805/Reviewed-by:
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zhang songyi authored
Return the sysfs_emit() and iterate_object_props() directly instead of using unnecessary variables. Reported-by:
Zeal Robot <zealci@zte.com.cn> Reviewed-by:
Anand Jain <anand.jain@oracle.com> Signed-off-by:
zhang songyi <zhang.songyi@zte.com.cn> Reviewed-by:
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Qu Wenruo authored
The problem of long mount time caused by block group item search is already known for some time, and the solution of block group tree has been proposed. There is really no need to bound this feature into extent tree v2, just introduce compat RO flag, BLOCK_GROUP_TREE, to correctly solve the problem. All the code handling block group root is already in the upstream kernel, thus this patch really only needs to introduce the new compat RO flag. This patch introduces one extra artificial limitation on block group tree feature, that free space cache v2 and no-holes feature must be enabled to use this new compat RO feature. This artificial requirement is mostly to reduce the test combinations, and can be a guideline for future features, to mostly rely on the latest default features. Signed-off-by:
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Qu Wenruo authored
The extent tree v2 needs a new root for storing all block group items, the whole feature hasn't been finished yet so we can afford to do some changes. My initial proposal years ago just added a new tree rootid, and load it from tree root, just like what we did for quota/free space tree/uuid/extent roots. But the extent tree v2 patches introduced a completely new way to store block group tree root into super block which is arguably wasteful. Currently there are only 3 trees stored in super blocks, and they all have their valid reasons: - Chunk root Needed for bootstrap. - Tree root Really the entry point for all trees. - Log root This is special as log root has to be updated out of existing transaction mechanism. There is not even any reason to put block group root into super blocks, the block group tree is updated at the same time as the old extent tree, no need for extra bootstrap/out-of-transaction update. So just move block group root from super block into tree root. Signed-off-by:
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Qu Wenruo authored
Currently there are two corner cases not handling compat RO flags correctly: - Remount We can still mount the fs RO with compat RO flags, then remount it RW. We should not allow any write into a fs with unsupported RO flags. - Still try to search block group items In fact, behavior/on-disk format change to extent tree should not need a full incompat flag. And since we can ensure fs with unsupported RO flags never got any writes (with above case fixed), then we can even skip block group items search at mount time. This patch will enhance the unsupported RO compat flags by: - Reject read-write remount if there are unsupported RO compat flags - Go dummy block group items directly for unsupported RO compat flags In fact, only changes to chunk/subvolume/root/csum trees should go incompat flags. The latter part should allow future change to extent tree to be compat RO flags. Thus this patch also needs to be backported to all stable trees. CC: stable@vger.kernel.org # 4.9+ Reviewed-by:
Nikolay Borisov <nborisov@suse.com> Signed-off-by:
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Qu Wenruo authored
We have hit some transaction abort due to -ENOSPC internally. Normally we should always reserve enough space for metadata for every transaction, thus hitting -ENOSPC should really indicate some cases we didn't expect. But unfortunately current error reporting will only give a kernel warning and stack trace, not really helpful to debug what's causing the problem. And mount option debug_enospc can only help when user can reproduce the problem, but under most cases, such transaction abort by -ENOSPC is really hard to reproduce. So this patch will dump all space infos (data, metadata, system) when we abort the first transaction with -ENOSPC. This should at least provide some clue to us. The example of a dump would look like this: BTRFS: Transaction aborted (error -28) WARNING: CPU: 8 PID: 3366 at fs/btrfs/transaction.c:2137 btrfs_commit_transaction+0xf81/0xfb0 [btrfs] <call trace skipped> ---[ end trace 0000000000000000 ]--- BTRFS info (device dm-1: state A): dumping space info: BTRFS info (device dm-1: state A): space_info DATA has 6791168 free, is not full BTRFS info (device dm-1: state A): space_info total=8388608, used=1597440, pinned=0, reserved=0, may_use=0, readonly=0 zone_unusable=0 BTRFS info (device dm-1: state A): space_info METADATA has 257114112 free, is not full BTRFS info (device dm-1: state A): space_info total=268435456, used=131072, pinned=180224, reserved=65536, may_use=10878976, readonly=65536 zone_unusable=0 BTRFS info (device dm-1: state A): space_info SYSTEM has 8372224 free, is not full BTRFS info (device dm-1: state A): space_info total=8388608, used=16384, pinned=0, reserved=0, may_use=0, readonly=0 zone_unusable=0 BTRFS info (device dm-1: state A): global_block_rsv: size 3670016 reserved 3670016 BTRFS info (device dm-1: state A): trans_block_rsv: size 0 reserved 0 BTRFS info (device dm-1: state A): chunk_block_rsv: size 0 reserved 0 BTRFS info (device dm-1: state A): delayed_block_rsv: size 4063232 reserved 4063232 BTRFS info (device dm-1: state A): delayed_refs_rsv: size 3145728 reserved 3145728 BTRFS: error (device dm-1: state A) in btrfs_commit_transaction:2137: errno=-28 No space left BTRFS info (device dm-1: state EA): forced readonly Reviewed-by:
Johannes Thumshirn <johannes.thumshirn@wdc.com> Signed-off-by:
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Qu Wenruo authored
For btrfs_space_info, its flags has only 4 possible values: - BTRFS_BLOCK_GROUP_SYSTEM - BTRFS_BLOCK_GROUP_METADATA | BTRFS_BLOCK_GROUP_DATA - BTRFS_BLOCK_GROUP_METADATA - BTRFS_BLOCK_GROUP_DATA Make the output more human readable, now it looks like: BTRFS info (device dm-1: state A): space_info METADATA has 251494400 free, is not full Reviewed-by:
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Qu Wenruo authored
[BACKGROUND] There is an incident report that, one user hibernated the system, with one btrfs on removable device still mounted. Then by some incident, the btrfs got mounted and modified by another system/OS, then back to the hibernated system. After resuming from the hibernation, new write happened into the victim btrfs. Now the fs is completely broken, since the underlying btrfs is no longer the same one before the hibernation, and the user lost their data due to various transid mismatch. [REPRODUCER] We can emulate the situation using the following small script: truncate -s 1G $dev mkfs.btrfs -f $dev mount $dev $mnt fsstress -w -d $mnt -n 500 sync xfs_freeze -f $mnt cp $dev $dev.backup # There is no way to mount the same cloned fs on the same system, # as the conflicting fsid will be rejected by btrfs. # Thus here we have to wipe the fs using a different btrfs. mkfs.btrfs -f $dev.backup dd if=$dev.backup of=$dev bs=1M xfs_freeze -u $mnt fsstress -w -d $mnt -n 20 umount $mnt btrfs check $dev The final fsck will fail due to some tree blocks has incorrect fsid. This is enough to emulate the problem hit by the unfortunate user. [ENHANCEMENT] Although such case should not be that common, it can still happen from time to time. From the view of btrfs, we can detect any unexpected super block change, and if there is any unexpected change, we just mark the fs read-only, and thaw the fs. By this we can limit the damage to minimal, and I hope no one would lose their data by this anymore. Suggested-by:
Goffredo Baroncelli <kreijack@libero.it> Link: https://lore.kernel.org/linux-btrfs/83bf3b4b-7f4c-387a-b286-9251e3991e34@bluemole.com/Reviewed-by:
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Christoph Hellwig authored
The I/O context structure is only used to pass the btrfs_device to the end I/O handler for I/Os that go to a single device. Stop allocating the I/O context for these cases by passing the optional btrfs_io_stripe argument to __btrfs_map_block to query the mapping information and then using a fast path submission and I/O completion handler. As the old btrfs_io_context based I/O submission path is only used for mirrored writes, rename the functions to make that clear and stop setting the btrfs_bio device and mirror_num field that is only used for reads. Reviewed-by:
Christoph Hellwig <hch@lst.de> Signed-off-by:
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Christoph Hellwig authored
There is no need for most of the btrfs_io_context when doing I/O to a single device. To support such I/O without the extra btrfs_io_context allocation, turn the mirror_num argument into a pointer so that it can be used to output the selected mirror number, and add an optional argument that points to a btrfs_io_stripe structure, which will be filled with a single extent if provided by the caller. In that case the btrfs_io_context allocation can be skipped as all information for the single device I/O is provided in the mirror_num argument and the on-stack btrfs_io_stripe. A caller that makes use of this new argument will be added in the next commit. Reviewed-by:
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Christoph Hellwig authored
Remove the orig_bio argument as it can be derived from the bioc, and the clone argument as it can be calculated from bioc and dev_nr. Reviewed-by:
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Christoph Hellwig authored
Split out a low-level btrfs_submit_dev_bio helper that just submits the bio without any cloning decisions or setting up the end I/O handler for future reuse by a different caller. Reviewed-by:
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Christoph Hellwig authored
Currently btrfs_bio end I/O handling is a bit of a mess. The bi_end_io handler and bi_private pointer of the embedded struct bio are both used to handle the completion of the high-level btrfs_bio and for the I/O completion for the low-level device that the embedded bio ends up being sent to. To support this bi_end_io and bi_private are saved into the btrfs_io_context structure and then restored after the bio sent to the underlying device has completed the actual I/O. Untangle this by adding an end I/O handler and private data to struct btrfs_bio for the high-level btrfs_bio based completions, and leave the actual bio bi_end_io handler and bi_private pointer entirely to the low-level device I/O. Reviewed-by:
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Christoph Hellwig authored
The parity raid write/recover functionality is currently not very well abstracted from the bio submission and completion handling in volumes.c: - the raid56 code directly completes the original btrfs_bio fed into btrfs_submit_bio instead of dispatching back to volumes.c - the raid56 code consumes the bioc and bio_counter references taken by volumes.c, which also leads to special casing of the calls from the scrub code into the raid56 code To fix this up supply a bi_end_io handler that calls back into the volumes.c machinery, which then puts the bioc, decrements the bio_counter and completes the original bio, and updates the scrub code to also take ownership of the bioc and bio_counter in all cases. Reviewed-by:
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Christoph Hellwig authored
The stripes_pending in the btrfs_io_context counts number of inflight low-level bios for an upper btrfs_bio. For reads this is generally one as reads are never cloned, while for writes we can trivially use the bio remaining mechanisms that is used for chained bios. To be able to make use of that mechanism, split out a separate trivial end_io handler for the cloned bios that does a minimal amount of error tracking and which then calls bio_endio on the original bio to transfer control to that, with the remaining counter making sure it is completed last. This then allows to merge btrfs_end_bioc into the original bio bi_end_io handler. To make this all work all error handling needs to happen through the bi_end_io handler, which requires a small amount of reshuffling in submit_stripe_bio so that the bio is cloned already by the time the suitability of the device is checked. This reduces the size of the btrfs_io_context and prepares splitting the btrfs_bio at the stripe boundary. Reviewed-by:
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Christoph Hellwig authored
Stop grabbing an extra bio_counter reference for each clone bio in a mirrored write and instead just release the one original reference in btrfs_end_bioc once all the bios for a single btrfs_bio have completed instead of at the end of btrfs_submit_bio once all bios have been submitted. This means the reference is now carried by the "upper" btrfs_bio only instead of each lower bio. Also remove the now unused btrfs_bio_counter_inc_noblocked helper. Reviewed-by:
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Christoph Hellwig authored
Pass the operation to btrfs_bio_alloc, matching what bio_alloc_bioset set does. Reviewed-by:
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Christoph Hellwig authored
volumes.c is the place that implements the storage layer using the btrfs_bio structure, so move the bio_set and allocation helpers there as well. To make up for the new initialization boilerplate, merge the two init/exit helpers in extent_io.c into a single one. Reviewed-by:
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Christoph Hellwig authored
btrfs never uses bio integrity data itself, so don't allocate the integrity pools for btrfs_bioset. This patch is a revert of the commit b208c2f7 ("btrfs: Fix crash due to not allocating integrity data for a set"). The integrity data pool is not needed, the bio-integrity code now handles allocating the integrity payload without that. Reviewed-by:
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Josef Bacik authored
We are calling __btrfs_remove_free_space_cache everywhere to cleanup the block group free space, however we can just use btrfs_remove_free_space_cache and pass in the block group in all of these places. Then we can remove __btrfs_remove_free_space_cache and rename __btrfs_remove_free_space_cache_locked to __btrfs_remove_free_space_cache. Signed-off-by:
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Josef Bacik authored
Now that lockdep is staying enabled through our entire CI runs I started seeing the following stack in generic/475 ------------[ cut here ]------------ WARNING: CPU: 1 PID: 2171864 at fs/btrfs/discard.c:604 btrfs_discard_update_discardable+0x98/0xb0 CPU: 1 PID: 2171864 Comm: kworker/u4:0 Not tainted 5.19.0-rc8+ #789 Hardware name: QEMU Standard PC (Q35 + ICH9, 2009), BIOS 1.13.0-2.fc32 04/01/2014 Workqueue: btrfs-cache btrfs_work_helper RIP: 0010:btrfs_discard_update_discardable+0x98/0xb0 RSP: 0018:ffffb857c2f7bad0 EFLAGS: 00010246 RAX: 0000000000000000 RBX: ffff8c85c605c200 RCX: 0000000000000001 RDX: 0000000000000000 RSI: ffffffff86807c5b RDI: ffffffff868a831e RBP: ffff8c85c4c54000 R08: 0000000000000000 R09: 0000000000000000 R10: ffff8c85c66932f0 R11: 0000000000000001 R12: ffff8c85c3899010 R13: ffff8c85d5be4f40 R14: ffff8c85c4c54000 R15: ffff8c86114bfa80 FS: 0000000000000000(0000) GS:ffff8c863bd00000(0000) knlGS:0000000000000000 CS: 0010 DS: 0000 ES: 0000 CR0: 0000000080050033 CR2: 00007f2e7f168160 CR3: 000000010289a004 CR4: 0000000000370ee0 Call Trace: __btrfs_remove_free_space_cache+0x27/0x30 load_free_space_cache+0xad2/0xaf0 caching_thread+0x40b/0x650 ? lock_release+0x137/0x2d0 btrfs_work_helper+0xf2/0x3e0 ? lock_is_held_type+0xe2/0x140 process_one_work+0x271/0x590 ? process_one_work+0x590/0x590 worker_thread+0x52/0x3b0 ? process_one_work+0x590/0x590 kthread+0xf0/0x120 ? kthread_complete_and_exit+0x20/0x20 ret_from_fork+0x1f/0x30 This is the code ctl = block_group->free_space_ctl; discard_ctl = &block_group->fs_info->discard_ctl; lockdep_assert_held(&ctl->tree_lock); We have a temporary free space ctl for loading the free space cache in order to avoid having allocations happening while we're loading the cache. When we hit an error we free it all up, however this also calls btrfs_discard_update_discardable, which requires block_group->free_space_ctl->tree_lock to be held. However this is our temporary ctl so this lock isn't held. Fix this by calling __btrfs_remove_free_space_cache_locked instead so that we only clean up the entries and do not mess with the discardable stats. Signed-off-by: -
Filipe Manana authored
When enabling quotas, at btrfs_quota_enable(), after committing the transaction, we change fs_info->quota_root to point to the quota root we created and set BTRFS_FS_QUOTA_ENABLED at fs_info->flags. Then we try to start the qgroup rescan worker, first by initializing it with a call to qgroup_rescan_init() - however if that fails we end up freeing the quota root but we leave fs_info->quota_root still pointing to it, this can later result in a use-after-free somewhere else. We have previously set the flags BTRFS_FS_QUOTA_ENABLED and BTRFS_QGROUP_STATUS_FLAG_ON, so we can only fail with -EINPROGRESS at btrfs_quota_enable(), which is possible if someone already called the quota rescan ioctl, and therefore started the rescan worker. So fix this by ignoring an -EINPROGRESS and asserting we can't get any other error. Reported-by:
Ye Bin <yebin10@huawei.com> Link: https://lore.kernel.org/linux-btrfs/20220823015931.421355-1-yebin10@huawei.com/ CC: stable@vger.kernel.org # 4.19+ Reviewed-by:
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Maciej S. Szmigiero authored
btrfs currently prints information about space cache or free space tree being in use on every remount, regardless whether such remount actually enabled or disabled one of these features. This is actually unnecessary since providing remount options changing the state of these features will explicitly print the appropriate notice. Let's instead print such unconditional information just on an initial mount to avoid filling the kernel log when, for example, laptop-mode-tools remount the fs on some events. Signed-off-by:
Maciej S. Szmigiero <maciej.szmigiero@oracle.com> Reviewed-by:
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Filipe Manana authored
At btrfs_del_root_ref() we are using two return variables, named 'ret' and 'err'. This makes it harder to follow and easier to return the wrong value in case an error happens - the previous patch in the series, which has the subject "btrfs: fix silent failure when deleting root reference", fixed a bug due to confusion created by these two variables. So change the function to use a single variable for tracking the return value of the function, using only 'ret', which is consistent with most of the codebase. Reviewed-by:
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Omar Sandoval authored
struct btrfs_caching_ctl::progress and struct btrfs_block_group::last_byte_to_unpin were previously needed to ensure that unpin_extent_range() didn't return a range to the free space cache before the caching thread had a chance to cache that range. However, the commit "btrfs: fix space cache corruption and potential double allocations" made it so that we always synchronously cache the block group at the time that we pin the extent, so this machinery is no longer necessary. Reviewed-by:
Omar Sandoval <osandov@fb.com> Signed-off-by:
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BingJing Chang authored
There is a bug causing send failures when processing an orphan directory with no links. In commit 46b2f459 ("Btrfs: fix send failure when root has deleted files still open")', the orphan inode issue was addressed. The send operation fails with a ENOENT error because of any attempts to generate a path for the inode with a link count of zero. Therefore, in that patch, sctx->ignore_cur_inode was introduced to be set if the current inode has a link count of zero for bypassing some unnecessary steps. And a helper function btrfs_unlink_all_paths() was introduced and called to clean up old paths found in the parent snapshot. However, not only regular files but also directories can be orphan inodes. So if the send operation meets an orphan directory, it will issue a wrong unlink command for that directory now. Soon the receive operation fails with a EISDIR error. Besides, the send operation also fails with a ENOENT error later when it tries to generate a path of it. Similar example but making an orphan dir for an incremental send: $ btrfs subvolume create vol $ mkdir vol/dir $ touch vol/dir/foo $ btrfs subvolume snapshot -r vol snap1 $ btrfs subvolume snapshot -r vol snap2 # Turn the second snapshot to RW mode and delete the whole dir while # holding an open file descriptor on it. $ btrfs property set snap2 ro false $ exec 73<snap2/dir $ rm -rf snap2/dir # Set the second snapshot back to RO mode and do an incremental send. $ btrfs property set snap2 ro true $ mkdir receive_dir $ btrfs send snap2 -p snap1 | btrfs receive receive_dir/ At subvol snap2 At snapshot snap2 ERROR: send ioctl failed with -2: No such file or directory ERROR: unlink dir failed. Is a directory Actually, orphan inodes are more common use cases in cascading backups. (Please see the illustration below.) In a cascading backup, a user wants to replicate a couple of snapshots from Machine A to Machine B and from Machine B to Machine C. Machine B doesn't take any RO snapshots for sending. All a receiver does is create an RW snapshot of its parent snapshot, apply the send stream and turn it into RO mode at the end. Even if all paths of some inodes are deleted in applying the send stream, these inodes would not be deleted and become orphans after changing the subvolume from RW to RO. Moreover, orphan inodes can occur not only in send snapshots but also in parent snapshots because Machine B may do a batch replication of a couple of snapshots. An illustration for cascading backups: Machine A (snapshot {1..n}) --> Machine B --> Machine C The idea to solve the problem is to delete all the items of orphan inodes before using these snapshots for sending. I used to think that the reasonable timing for doing that is during the ioctl of changing the subvolume from RW to RO because it sounds good that we will not modify the fs tree of a RO snapshot anymore. However, attempting to do the orphan cleanup in the ioctl would be pointless. Because if someone is holding an open file descriptor on the inode, the reference count of the inode will never drop to 0. Then iput() cannot trigger eviction, which finally deletes all the items of it. So we try to extend the original patch to handle orphans in send/parent snapshots. Here are several cases that need to be considered: Case 1: BTRFS_COMPARE_TREE_NEW | send snapshot | action -------------------------------- nlink | 0 | ignore In case 1, when we get a BTRFS_COMPARE_TREE_NEW tree comparison result, it means that a new inode is found in the send snapshot and it doesn't appear in the parent snapshot. Since this inode has a link count of zero (It's an orphan and there're no paths for it.), we can leverage sctx->ignore_cur_inode in the original patch to prevent it from being created. Case 2: BTRFS_COMPARE_TREE_DELETED | parent snapshot | action ---------------------------------- nlink | 0 | as usual In case 2, when we get a BTRFS_COMPARE_TREE_DELETED tree comparison result, it means that the inode only appears in the parent snapshot. As usual, the send operation will try to delete all its paths. However, this inode has a link count of zero, so no paths of it will be found. No deletion operations will be issued. We don't need to change any logic. Case 3: BTRFS_COMPARE_TREE_CHANGED | | parent snapshot | send snapshot | action ----------------------------------------------------------------------- subcase 1 | nlink | 0 | 0 | ignore subcase 2 | nlink | >0 | 0 | new_gen(deletion) subcase 3 | nlink | 0 | >0 | new_gen(creation) In case 3, when we get a BTRFS_COMPARE_TREE_CHANGED tree comparison result, it means that the inode appears in both snapshots. Here are 3 subcases. First, when the inode has link counts of zero in both snapshots. Since there are no paths for this inode in (source/destination) parent snapshots and we don't care about whether there is also an orphan inode in destination or not, we can set sctx->ignore_cur_inode on to prevent it from being created. For the second and the third subcases, if there are paths in one snapshot and there're no paths in the other snapshot for this inode. We can treat this inode as a new generation. We can also leverage the logic handling a new generation of an inode with small adjustments. Then it will delete all old paths and create a new inode with new attributes and paths only when there's a positive link count in the send snapshot. In subcase 2, the send operation only needs to delete all old paths as in the parent snapshot. But it may require more operations for a directory to remove its old paths. If a not-empty directory is going to be deleted (because it has a link count of zero in the send snapshot) but there are files/directories with bigger inode numbers under it, the send operation will need to rename it to its orphan name first. After processing and deleting the last item under this directory, the send operation will check this directory, aka the parent directory of the last item, again and issue a rmdir operation to remove it finally. Therefore, we also need to treat inodes with a link count of zero as if they didn't exist in get_cur_inode_state(), which is used in process_recorded_refs(). By doing this, when checking a directory with orphan names after the last item under it has been deleted, the send operation now can properly issue a rmdir operation. Otherwise, without doing this, the orphan directory with an orphan name would be kept here at the end due to the existing inode with a link count of zero being found. In subcase 3, as in case 2, no old paths would be found, so no deletion operations will be issued. The send operation will only create a new one for that inode. Note that subcase 3 is not common. That's because it's easy to reduce the hard links of an inode, but once all valid paths are removed, there are no valid paths for creating other hard links. The only way to do that is trying to send an older snapshot after a newer snapshot has been sent. Reviewed-by:
Robbie Ko <robbieko@synology.com> Reviewed-by:
BingJing Chang <bingjingc@synology.com> Signed-off-by:
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BingJing Chang authored
Refactor get_inode_info() to populate all wanted fields on an output structure. Besides, also introduce a helper function called get_inode_gen(), which is commonly used. Reviewed-by:
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Ethan Lien authored
After we copied data to page cache in buffered I/O, we 1. Insert a EXTENT_UPTODATE state into inode's io_tree, by endio_readpage_release_extent(), set_extent_delalloc() or set_extent_defrag(). 2. Set page uptodate before we unlock the page. But the only place we check io_tree's EXTENT_UPTODATE state is in btrfs_do_readpage(). We know we enter btrfs_do_readpage() only when we have a non-uptodate page, so it is unnecessary to set EXTENT_UPTODATE. For example, when performing a buffered random read: fio --rw=randread --ioengine=libaio --direct=0 --numjobs=4 \ --filesize=32G --size=4G --bs=4k --name=job \ --filename=/mnt/file --name=job Then check how many extent_state in io_tree: cat /proc/slabinfo | grep btrfs_extent_state | awk '{print $2}' w/o this patch, we got 640567 btrfs_extent_state. w/ this patch, we got 204 btrfs_extent_state. Maintaining such a big tree brings overhead since every I/O needs to insert EXTENT_LOCKED, insert EXTENT_UPTODATE, then remove EXTENT_LOCKED. And in every insert or remove, we need to lock io_tree, do tree search, alloc or dealloc extent states. By removing unnecessary EXTENT_UPTODATE, we keep io_tree in a minimal size and reduce overhead when performing buffered I/O. Reviewed-by:Ethan Lien <ethanlien@synology.com> Signed-off-by:
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Filipe Manana authored
During log replay, when adding/replacing inode references, there are two special cases that have special code for them: 1) When we have an inode with two or more hardlinks in the same directory, therefore two or more names encoded in the same inode reference item, and one of the hard links gets renamed to the old name of another hard link - that is, the index number for a name changes. This was added in commit 0d836392 ("Btrfs: fix mount failure after fsync due to hard link recreation"), and is covered by test case generic/502 from fstests; 2) When we have several inodes that got renamed to an old name of some other inode, in a cascading style. The code to deal with this special case was added in commit 6b5fc433 ("Btrfs: fix fsync after succession of renames of different files"), and is covered by test cases generic/526 and generic/527 from fstests. Both cases can be deal with by making sure __add_inode_ref() is always called by add_inode_ref() for every name encoded in the inode reference item, and not just for the first name that has a conflict. With such change we no longer need that special casing for the two cases mentioned before. So do those changes. Signed-off-by:
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David Sterba authored
When discard=async was introduced there were also sysfs knobs and stats for debugging and tuning, hidden under CONFIG_BTRFS_DEBUG. The defaults have been set and so far seem to satisfy all users on a range of workloads. As there are not only tunables (like iops or kbps) but also stats tracking amount of discardable bytes, that should be available when the async discard is on (otherwise it's not). The stats are moved from the per-fs debug directory, so it's under /sys/fs/btrfs/FSID/discard - discard_bitmap_bytes - amount of discarded bytes from data tracked as bitmaps - discard_extent_bytes - dtto but as extents - discard_bytes_saved - - discardable_bytes - amount of bytes that can be discarded - discardable_extents - number of extents to be discarded - iops_limit - tunable limit of number of discard IOs to be issued - kbps_limit - tunable limit of kilobytes per second issued as discard IO - max_discard_size - tunable limit for size of one IO discard request Reviewed-by: -
Filipe Manana authored
When logging a directory we start by flushing all its delayed items. That results in adding dir index items to the subvolume btree, for new dentries, and removing dir index items from the subvolume btree for any dentries that were deleted. This makes it straightforward to log a directory simply by iterating over all the modified subvolume btree leaves, especially when we used to log both dir index keys and dir item keys (before commit 339d0354 ("btrfs: only copy dir index keys when logging a directory") and when we used to copy old dir index entries for leaves modified in the current transaction (before commit 732d591a ("btrfs: stop copying old dir items when logging a directory")). From an efficiency point of view this has a couple of drawbacks: 1) Adds extra latency, due to copying delayed items to the subvolume btree and deleting dir index items from the btree. Further if there are other tasks accessing the btree, which is common (syscalls like creat, mkdir, rename, link, unlink, truncate, reflinks, etc, finishing an ordered extent, etc), lock contention can cause further delays, both to the task logging a directory and to the other tasks accessing the btree; 2) More time spent overall flushing delayed items, if after logging the directory further changes are done to the directory in the same transaction. For example, if we add 10 dentries to a directory, fsync it, add more 10 dentries, fsync it again, then add more 10 dentries and fsync it again, then we end up inserting 3 batches of 10 items to the subvolume btree. With the changes from this patch, we flush all the delayed items to the btree only once - a single batch of 30 items, and outside the logging code (transaction commit or when delayed items are flushed asynchronously). This change simply skips the flushing of delayed items every time we log a directory. Instead we copy the delayed insertion items directly to the log tree and delete delayed deletion items directly from the log tree. Therefore avoiding changing first the subvolume btree and then scanning it for new items to copy from it to the log tree and detecting deletions by observing gaps in consecutive dir index keys in subvolume btree leaves. Running the following tests on a non-debug kernel (Debian's default kernel config), on a box with a NVMe device, a 12 cores Intel CPU and 64G of ram, produced the results below. The results compare a branch without this patch and all the other patches it depends on versus the same branch with the patchset applied. The patchset is comprised of the following patches: btrfs: don't drop dir index range items when logging a directory btrfs: remove the root argument from log_new_dir_dentries() btrfs: update stale comment for log_new_dir_dentries() btrfs: free list element sooner at log_new_dir_dentries() btrfs: avoid memory allocation at log_new_dir_dentries() for common case btrfs: remove root argument from btrfs_delayed_item_reserve_metadata() btrfs: store index number instead of key in struct btrfs_delayed_item btrfs: remove unused logic when looking up delayed items btrfs: shrink the size of struct btrfs_delayed_item btrfs: search for last logged dir index if it's not cached in the inode btrfs: move need_log_inode() to above log_conflicting_inodes() btrfs: move log_new_dir_dentries() above btrfs_log_inode() btrfs: log conflicting inodes without holding log mutex of the initial inode btrfs: skip logging parent dir when conflicting inode is not a dir btrfs: use delayed items when logging a directory Custom test script for testing time spent at btrfs_log_inode(): #!/bin/bash DEV=/dev/nvme0n1 MNT=/mnt/nvme0n1 # Total number of files to create in the test directory. NUM_FILES=10000 # Fsync after creating or renaming N files. FSYNC_AFTER=100 umount $DEV &> /dev/null mkfs.btrfs -f $DEV mount -o ssd $DEV $MNT TEST_DIR=$MNT/testdir mkdir $TEST_DIR echo "Creating files..." for ((i = 1; i <= $NUM_FILES; i++)); do echo -n > $TEST_DIR/file_$i if (( ($i % $FSYNC_AFTER) == 0 )); then xfs_io -c "fsync" $TEST_DIR fi done sync echo "Renaming files..." for ((i = 1; i <= $NUM_FILES; i++)); do mv $TEST_DIR/file_$i $TEST_DIR/file_$i.renamed if (( ($i % $FSYNC_AFTER) == 0 )); then xfs_io -c "fsync" $TEST_DIR fi done umount $MNT And using the following bpftrace script to capture the total time that is spent at btrfs_log_inode(): #!/usr/bin/bpftrace k:btrfs_log_inode { @start_log_inode[tid] = nsecs; } kr:btrfs_log_inode /@start_log_inode[tid]/ { $dur = (nsecs - @start_log_inode[tid]) / 1000; @btrfs_log_inode_total_time = sum($dur); delete(@start_log_inode[tid]); } END { clear(@start_log_inode); } Result before applying patchset: @btrfs_log_inode_total_time: 622642 Result after applying patchset: @btrfs_log_inode_total_time: 354134 (-43.1% time spent) The following dbench script was also used for testing: #!/bin/bash NUM_JOBS=$(nproc --all) DEV=/dev/nvme0n1 MNT=/mnt/nvme0n1 MOUNT_OPTIONS="-o ssd" MKFS_OPTIONS="-O no-holes -R free-space-tree" echo "performance" | \ tee /sys/devices/system/cpu/cpu*/cpufreq/scaling_governor umount $DEV &> /dev/null mkfs.btrfs -f $MKFS_OPTIONS $DEV mount $MOUNT_OPTIONS $DEV $MNT dbench -D $MNT --skip-cleanup -t 120 -S $NUM_JOBS umount $MNT Before patchset: Operation Count AvgLat MaxLat ---------------------------------------- NTCreateX 3322265 0.034 21.032 Close 2440562 0.002 0.994 Rename 140664 1.150 269.633 Unlink 670796 1.093 269.678 Deltree 96 5.481 15.510 Mkdir 48 0.004 0.052 Qpathinfo 3010924 0.014 8.127 Qfileinfo 528055 0.001 0.518 Qfsinfo 552113 0.003 0.372 Sfileinfo 270575 0.005 0.688 Find 1164176 0.052 13.931 WriteX 1658537 0.019 5.918 ReadX 5207412 0.003 1.034 LockX 10818 0.003 0.079 UnlockX 10818 0.002 0.313 Flush 232811 1.027 269.735 Throughput 869.867 MB/sec (sync dirs) 12 clients 12 procs max_latency=269.741 ms After patchset: Operation Count AvgLat MaxLat ---------------------------------------- NTCreateX 4152738 0.029 20.863 Close 3050770 0.002 1.119 Rename 175829 0.871 211.741 Unlink 838447 0.845 211.724 Deltree 120 4.798 14.162 Mkdir 60 0.003 0.005 Qpathinfo 3763807 0.011 4.673 Qfileinfo 660111 0.001 0.400 Qfsinfo 690141 0.003 0.429 Sfileinfo 338260 0.005 0.725 Find 1455273 0.046 6.787 WriteX 2073307 0.017 5.690 ReadX 6509193 0.003 1.171 LockX 13522 0.003 0.077 UnlockX 13522 0.002 0.125 Flush 291044 0.811 211.631 Throughput 1089.27 MB/sec (sync dirs) 12 clients 12 procs max_latency=211.750 ms (+25.2% throughput, -21.5% max latency) Signed-off-by:
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Filipe Manana authored
When we find a conflicting inode (an inode that had the same name and parent directory as the inode we are logging now) that was deleted in the current transaction, we always end up logging its parent directory. This is to deal with the case where the conflicting inode corresponds to a deleted subvolume/snapshot or a directory that had subvolumes/snapshots (or some subdirectory inside it had subvolumes/snapshots, etc), because we can't deal with dropping subvolumes/snapshots during log replay. So if we log the parent directory, and if we are dealing with these special cases, then we fallback to a transaction commit when logging the parent, because its last_unlink_trans will match the current transaction (which gets set and propagated when a subvolume/snapshot is deleted). This change skips the logging of the parent directory when the conflicting inode is not a directory (or a subvolume/snapshot). This is ok because in this case logging the current inode is enough to trigger an unlink of the conflicting inode during log replay. So for a case like this: $ mkdir /mnt/dir $ echo -n "first foo data" > /mnt/dir/foo $ sync $ rm -f /mnt/dir/foo $ echo -n "second foo data" > /mnt/dir/foo $ xfs_io -c "fsync" /mnt/dir/foo We avoid logging parent directory "dir" when logging the new file "foo". In other cases it avoids falling back to a transaction commit, when the parent directory has a last_unlink_trans value that matches the current transaction, due to moving a file from it to some other directory. This is a case that happens frequently with dbench for example, where a new file that has the name/parent of another file that was deleted in the current transaction, is fsynced. Signed-off-by:
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Filipe Manana authored
When logging an inode, if we detect the inode has a reference that conflicts with some other inode that got renamed, we log that other inode while holding the log mutex of the current inode. We then find out if there are other inodes that conflict with the first conflicting inode, and log them while under the log mutex of the original inode. This is fine because the recursion can only happen once. For the upcoming work where we directly log delayed items without flushing them first to the subvolume tree, this recursion adds a lot of complexity and it's hard to keep lockdep happy about it. So collect a list of conflicting inodes and then log the inodes after unlocking the log mutex of the inode we started with. Also limit the maximum number of conflict inodes we log to 10, to avoid spending too much time logging (and maybe allocating too many list elements too), as typically we don't have more than 1 or 2 conflicting inodes - if we go over the limit, simply fallback to a transaction commit. It is possible to have a very long list of conflicting inodes to be intentionally created by a user if he/she creates a very long succession of renames like this: (...) rename E to F rename D to E rename C to D rename B to C rename A to B touch A (create a new file named A) fsync A If that happened for a sequence of hundreds or thousands of renames, it could massively slow down the logging and cause other secondary effects like for example blocking other fsync operations and transaction commits for a very long time (assuming it wouldn't run into -ENOSPC or -ENOMEM first). However such cases are very uncommon to happen in practice, nevertheless it's better to be prepared for them and avoid chaos. Such long sequence of conflicting inodes could be created before this change. Signed-off-by:
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