- 30 10月, 2008 1 次提交
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由 Chris Mason 提交于
This is a large change for adding compression on reading and writing, both for inline and regular extents. It does some fairly large surgery to the writeback paths. Compression is off by default and enabled by mount -o compress. Even when the -o compress mount option is not used, it is possible to read compressed extents off the disk. If compression for a given set of pages fails to make them smaller, the file is flagged to avoid future compression attempts later. * While finding delalloc extents, the pages are locked before being sent down to the delalloc handler. This allows the delalloc handler to do complex things such as cleaning the pages, marking them writeback and starting IO on their behalf. * Inline extents are inserted at delalloc time now. This allows us to compress the data before inserting the inline extent, and it allows us to insert an inline extent that spans multiple pages. * All of the in-memory extent representations (extent_map.c, ordered-data.c etc) are changed to record both an in-memory size and an on disk size, as well as a flag for compression. From a disk format point of view, the extent pointers in the file are changed to record the on disk size of a given extent and some encoding flags. Space in the disk format is allocated for compression encoding, as well as encryption and a generic 'other' field. Neither the encryption or the 'other' field are currently used. In order to limit the amount of data read for a single random read in the file, the size of a compressed extent is limited to 128k. This is a software only limit, the disk format supports u64 sized compressed extents. In order to limit the ram consumed while processing extents, the uncompressed size of a compressed extent is limited to 256k. This is a software only limit and will be subject to tuning later. Checksumming is still done on compressed extents, and it is done on the uncompressed version of the data. This way additional encodings can be layered on without having to figure out which encoding to checksum. Compression happens at delalloc time, which is basically singled threaded because it is usually done by a single pdflush thread. This makes it tricky to spread the compression load across all the cpus on the box. We'll have to look at parallel pdflush walks of dirty inodes at a later time. Decompression is hooked into readpages and it does spread across CPUs nicely. Signed-off-by: NChris Mason <chris.mason@oracle.com>
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- 02 10月, 2008 2 次提交
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由 Jim Meyering 提交于
Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Jim Meyering 提交于
Signed-off-by: NChris Mason <chris.mason@oracle.com>
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- 30 9月, 2008 1 次提交
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由 Chris Mason 提交于
This improves the comments at the top of many functions. It didn't dive into the guts of functions because I was trying to avoid merging problems with the new allocator and back reference work. extent-tree.c and volumes.c were both skipped, and there is definitely more work todo in cleaning and commenting the code. Signed-off-by: NChris Mason <chris.mason@oracle.com>
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- 29 9月, 2008 1 次提交
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由 Chris Mason 提交于
btrfs-vol -a /dev/xxx will zero the first and last two MB of the device. The kernel code needs to wait for this IO to finish before it adds the device. btrfs metadata IO does not happen through the block device inode. A separate address space is used, allowing the zero filled buffer heads in the block device inode to be written to disk after FS metadata starts going down to the disk via the btrfs metadata inode. The end result is zero filled metadata blocks after adding new devices into the filesystem. The fix is a simple filemap_write_and_wait on the block device inode before actually inserting it into the pool of available devices. Signed-off-by: NChris Mason <chris.mason@oracle.com>
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- 26 9月, 2008 4 次提交
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由 Zheng Yan 提交于
This patch updates the space balancing code to utilize the new backref format. Before, btrfs-vol -b would break any COW links on data blocks or metadata. This was slow and caused the amount of space used to explode if a large number of snapshots were present. The new code can keeps the sharing of all data extents and most of the tree blocks. To maintain the sharing of data extents, the space balance code uses a seperate inode hold data extent pointers, then updates the references to point to the new location. To maintain the sharing of tree blocks, the space balance code uses reloc trees to relocate tree blocks in reference counted roots. There is one reloc tree for each subvol, and all reloc trees share same root key objectid. Reloc trees are snapshots of the latest committed roots of subvols (root->commit_root). To relocate a tree block referenced by a subvol, there are two steps. COW the block through subvol's reloc tree, then update block pointer in the subvol to point to the new block. Since all reloc trees share same root key objectid, doing special handing for tree blocks owned by them is easy. Once a tree block has been COWed in one reloc tree, we can use the resulting new block directly when the same block is required to COW again through other reloc trees. In this way, relocated tree blocks are shared between reloc trees, so they are also shared between subvols. Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Zheng Yan 提交于
Btrfs has a cache of reference counts in leaves, allowing it to avoid reading tree leaves while deleting snapshots. To reduce contention with multiple subvolumes, this cache is private to each subvolume. This patch adds shared reference cache support. The new space balancing code plays with multiple subvols at the same time, So the old per-subvol reference cache is not well suited. Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Chris Mason 提交于
Btrfs metadata writeback is fairly expensive. Once a tree block is written it must be cowed before it can be changed again. The btree writepages code has a threshold based on a count of dirty btree bytes which is updated as IO is sent out. This changes btree_writepages to skip the writeout if there are less than 32MB of dirty bytes from the btrees, improving performance across many workloads. Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Chris Mason 提交于
Btrfs had compatibility code for kernels back to 2.6.18. These have been removed, and will be maintained in a separate backport git tree from now on. Signed-off-by: NChris Mason <chris.mason@oracle.com>
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- 25 9月, 2008 31 次提交
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由 Zheng Yan 提交于
This patch makes the back reference system to explicit record the location of parent node for all types of extents. The location of parent node is placed into the offset field of backref key. Every time a tree block is balanced, the back references for the affected lower level extents are updated. Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Chris Mason 提交于
Tree blocks were using async bio submission, but the sum was still being done directly during writepage. This moves the checksumming into the worker thread. Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Josef Bacik 提交于
1) replace the per fs_info extent_io_tree that tracked free space with two rb-trees per block group to track free space areas via offset and size. The reason to do this is because most allocations come with a hint byte where to start, so we can usually find a chunk of free space at that hint byte to satisfy the allocation and get good space packing. If we cannot find free space at or after the given offset we fall back on looking for a chunk of the given size as close to that given offset as possible. When we fall back on the size search we also try to find a slot as close to the size we want as possible, to avoid breaking small chunks off of huge areas if possible. 2) remove the extent_io_tree that tracked the block group cache from fs_info and replaced it with an rb-tree thats tracks block group cache via offset. also added a per space_info list that tracks the block group cache for the particular space so we can lookup related block groups easily. 3) cleaned up the allocation code to make it a little easier to read and a little less complicated. Basically there are 3 steps, first look from our provided hint. If we couldn't find from that given hint, start back at our original search start and look for space from there. If that fails try to allocate space if we can and start looking again. If not we're screwed and need to start over again. 4) small fixes. there were some issues in volumes.c where we wouldn't allocate the rest of the disk. fixed cow_file_range to actually pass the alloc_hint, which has helped a good bit in making the fs_mark test I run have semi-normal results as we run out of space. Generally with data allocations we don't track where we last allocated from, so everytime we did a data allocation we'd search through every block group that we have looking for free space. Now searching a block group with no free space isn't terribly time consuming, it was causing a slight degradation as we got more data block groups. The alloc_hint has fixed this slight degredation and made things semi-normal. There is still one nagging problem I'm working on where we will get ENOSPC when there is definitely plenty of space. This only happens with metadata allocations, and only when we are almost full. So you generally hit the 85% mark first, but sometimes you'll hit the BUG before you hit the 85% wall. I'm still tracking it down, but until then this seems to be pretty stable and make a significant performance gain. Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Chris Mason 提交于
I had incorrectly disabled the check for the block number being correct in the header block. Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Chris Mason 提交于
This is the same way the transaction code makes sure that all the other tree blocks are safely on disk. There's an extent_io tree for each root, and any blocks allocated to the tree logs are recorded in that tree. At tree-log sync, the extent_io tree is walked to flush down the dirty pages and wait for them. The main benefit is less time spent walking the tree log and skipping clean pages, and getting sequential IO down to the drive. Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Chris Mason 提交于
Since tree log blocks get freed every transaction, they never really need to be written to disk. This skips the step where we update metadata to record they were allocated. Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Chris Mason 提交于
Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Chris Mason 提交于
Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Chris Mason 提交于
* Pin down data blocks to prevent them from being reallocated like so: trans 1: allocate file extent trans 2: free file extent trans 3: free file extent during old snapshot deletion trans 3: allocate file extent to new file trans 3: fsync new file Before the tree logging code, this was legal because the fsync would commit the transation that did the final data extent free and the transaction that allocated the extent to the new file at the same time. With the tree logging code, the tree log subtransaction can commit before the transaction that freed the extent. If we crash, we're left with two different files using the extent. * Don't wait in start_transaction if log replay is going on. This avoids deadlocks from iput while we're cleaning up link counts in the replay code. * Don't deadlock in replay_one_name by trying to read an inode off the disk while holding paths for the directory * Hold the buffer lock while we mark a buffer as written. This closes a race where someone is changing a buffer while we write it. They are supposed to mark it dirty again after they change it, but this violates the cow rules. Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Chris Mason 提交于
File syncs and directory syncs are optimized by copying their items into a special (copy-on-write) log tree. There is one log tree per subvolume and the btrfs super block points to a tree of log tree roots. After a crash, items are copied out of the log tree and back into the subvolume. See tree-log.c for all the details. Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Chris Mason 提交于
Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Chris Mason 提交于
The current code waits for the count of async bio submits to get below a given threshold if it is too high right after adding the latest bio to the work queue. This isn't optimal because the caller may have sequential adjacent bios pending they are waiting to send down the pipe. This changeset requires the caller to wait on the async bio count, and changes the async checksumming submits to wait for async bios any time they self throttle. The end result is much higher sequential throughput. Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Chris Mason 提交于
Before, the btrfs bdi congestion function was used to test for too many async bios. This keeps that check to throttle pdflush, but also adds a check while queuing bios. Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Chris Mason 提交于
This takes the csum mutex deeper in the call chain and releases it more often. Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Chris Mason 提交于
Before this change, btrfs would use a bdi congestion function to make sure there weren't too many pending async checksum work items. This change makes the process creating async work items wait instead, leading to fewer congestion returns from the bdi. This improves pdflush background_writeout scanning. Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Chris Mason 提交于
Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Chris Mason 提交于
After writing out all the remaining btree blocks in the transaction, the commit code would use filemap_fdatawait to make sure it was all on disk. This means it would wait for blocks written by other procs as well. The new code walks the list of blocks for this transaction again and waits only for those required by this transaction. Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Chris Mason 提交于
Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Chris Mason 提交于
Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Chris Mason 提交于
Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Chris Mason 提交于
Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Chris Mason 提交于
* Make walk_down_tree wake up throttled tasks more often * Make walk_down_tree call cond_resched during long loops * As the size of the ref cache grows, wait longer in throttle * Get rid of the reada code in walk_down_tree, the leaves don't get read anymore, thanks to the ref cache. Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Chris Mason 提交于
Large streaming reads make for large bios, which means each entry on the list async work queues represents a large amount of data. IO congestion throttling on the device was kicking in before the async worker threads decided a single thread was busy and needed some help. The end result was that a streaming read would result in a single CPU running at 100% instead of balancing the work off to other CPUs. This patch also changes the pre-IO checksum lookup done by reads to work on a per-bio basis instead of a per-page. This results in many extra btree lookups on large streaming reads. Doing the checksum lookup right before bio submit allows us to reuse searches while processing adjacent offsets. Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Yan 提交于
The memory reclaiming issue happens when snapshot exists. In that case, some cache entries may not be used during old snapshot dropping, so they will remain in the cache until umount. The patch adds a field to struct btrfs_leaf_ref to record create time. Besides, the patch makes all dead roots of a given snapshot linked together in order of create time. After a old snapshot was completely dropped, we check the dead root list and remove all cache entries created before the oldest dead root in the list. Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Chris Mason 提交于
It was incorrectly clearing the up to date flag on the buffer even when the buffer properly verified. Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Chris Mason 提交于
A large reference cache is directly related to a lot of work pending for the cleaner thread. This throttles back new operations based on the size of the reference cache so the cleaner thread will be able to keep up. Overall, this actually makes the FS faster because the cleaner thread will be more likely to find things in cache. Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Chris Mason 提交于
This changes the reference cache to make a single cache per root instead of one cache per transaction, and to key by the byte number of the disk block instead of the keys inside. This makes it much less likely to have cache misses if a snapshot or something has an extra reference on a higher node or a leaf while the first transaction that added the leaf into the cache is dropping. Some throttling is added to functions that free blocks heavily so they wait for old transactions to drop. Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Yan Zheng 提交于
Much of the IO done while dropping snapshots is done looking up leaves in the filesystem trees to see if they point to any extents and to drop the references on any extents found. This creates a cache so that IO isn't required. Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Josef Bacik 提交于
Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Chris Mason 提交于
Before setting an extent to delalloc, the code needs to wait for pending ordered extents. Also, the relocation code needs to wait for ordered IO before scanning the block group again. This is because the extents are not removed until the IO for the new extents is finished Signed-off-by: NChris Mason <chris.mason@oracle.com>
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由 Chris Mason 提交于
Checksum items are not inserted into the tree until all of the io from a given extent is complete. This means one dirty page from an extent may be written, freed, and then read again before the entire extent is on disk and the checksum item is inserted. The checksums themselves are stored in the ordered extent so they can be inserted in bulk when IO is complete. On read, if a checksum item isn't found, the ordered extents were being searched for a checksum record. This all worked most of the time, but the checksum insertion code tries to reduce the number of tree operations by pre-inserting checksum items based on i_size and a few other factors. This means the read code might find a checksum item that hasn't yet really been filled in. This commit changes things to check the ordered extents first and only dive into the btree if nothing was found. This removes the need for extra locking and is more reliable. Signed-off-by: NChris Mason <chris.mason@oracle.com>
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