Review: FFS background. Review: FFS [McKusic] basics. Basic FFS data structures. FFS disk layout. FFS superblock. Cylinder groups

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Review: FFS background 1980s improvement to original Unix FS, which had: byte blocks - Free blocks in linked list - All inodes at beginning of disk - Low throughput: 512 bytes per average seek time
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Review: FFS background 1980s improvement to original Unix FS, which had: byte blocks - Free blocks in linked list - All inodes at beginning of disk - Low throughput: 512 bytes per average seek time Unix FS performance problems: - Transfers only 512 bytes per disk access - Eventually random allocation 512 bytes / disk seek - Inodes far from directory and file - Within directory, inodes far from each other Also had some usability problems: - 14-character file names a pain - Can t atomically update file in crash-proof way Review: FFS [McKusic] basics Change block size to at least 4K - To avoid wasting space, use fragments for ends of files Cylinder groups spread inodes around disk Bitmaps replace free list FS reserves space to improve allocation - Tunable parameter, default 10% - Only superuser can use space when over 90% full Usability improvements: - File names up to 255 characters - Atomic rename system call - Symbolic links assign one file name to another 1/36 2/36 cylinder groups bookkeeping information FFS disk layout inodes superblocks blocks Each cylinder group has its own: - Superblock - Bookkeeping information - Set of inodes - Data/directory blocks 3/36 directory name i-number Basic FFS structures inode meta indirect ptr double indir indirect block Inode is key structure for each file contents - Has permissions and access/modification/inode-change times - Has link count (# directories containing file); file deleted when 0 - Points to blocks of file (and indirect blocks) By convention, inode #2 always root directory 4/36 FFS superblock Cylinder groups Superblock contains file system parameters - characteristics, block size, CG info - Information necessary to get inode given i-number Replicated once per cylinder group - At shifting offsets, so as to span multiple platters - Contains magic number to find replicas if 1st superblock dies Contains non-replicated summary info - # blocks, fragments, inodes, directories in FS - Flag stating if FS was cleanly unmounted Groups related inodes and their Contains a number of inodes (set when FS created) - Default one inode per 2K Contains file and directory blocks Contains bookkeeping information - Block map bit map of available fragments - Summary info within CG # free inodes, blocks/frags, files, directories - # free blocks by rotational position (8 positions) [Was reasonable in 1980s when disks weren t commonly zoned] 5/36 6/36 Inode allocation Fragment allocation Allocate inodes in same CG as directory if possible New directories put in new cylinder groups - Consider CGs with greater than average # free inodes - Chose CG with smallest # directories Within CG, inodes allocated randomly (next free) - Would like related inodes as close as possible - OK, because one CG doesn t have that many inodes Allocate space when user writes beyond end of file Want last block to be a fragment if not full-size - If already a fragment, may contain space for write done - Else, must deallocate any existing fragment, allocate new If no appropriate free fragments, break full block Problem: Slow for many small writes - (Partial) soution: newstat struct fieldst blksize - Tells applications file system block size - stdio library can buffer this much 7/36 8/36 Block allocation Directories Try to optimize for sequential access - If available, use rotationally close block in same cylinder - Otherwise, use block in same CG - If CG totally full, find other CG with quadratic hashing ie, if CG #n is full, try n + 1 2, n + 2 2, n + 3 2, (mod #CGs) - Otherwise, search all CGs for some free space Problem: Don t want one file filling up whole CG - Otherwise other inodes will have far away Solution: Break big files over many CGs - But large extents in each CGs, so sequential access doesn t require many seeks Inodes like files, but with different type bits Contents considered as 512-byte chunks Each chunk has direct structure(s) with: - 32-bit inumber - 16-bit size of directory entry - 8-bit file type (NEW) - 8-bit length of file name Coalesce when deleting - If firstdirect in chunk deleted, set inumber = 0 Periodically compact directory chunks 9/36 10/36 Updating FFS for the 90s No longer wanted to assume rotational delay - With disk caches, want contiguously allocated Solution: Cluster writes - FS delays writing a block back to get more blocks - Accumulates blocks into 64K clusters, written at once Allocation of clusters similar to fragments/blocks - Summary info - Cluster map has one bit for each 64K if all free Also read in 64K chunks when doing read ahead Fixing corruption fsck Must run FS check (fsck) program after crash Summary info usually bad after crash - Scan to check free block map, block/inode counts System may have corrupt inodes (not simple crash) - Bad block numbers, cross-allocation, etc - Do sanity check, clear inodes with garbage Fields in inodes may be wrong - Count number of directory entries to verify link count, if no entries but count = 0, move tolost+found - Make sure size and used counts match blocks Directories may be bad - Holes illegal, and must be valid, - All directories must be reachable 11/36 12/36 Crash recovery permeates FS code Ordering of updates Have to ensure fsck can recover file system Example: Suppose all written asynchronously Delete/truncate a file, append to other file, crash - New file may reuse block from old - Old inode may not be updated - Cross-allocation! - Often inode with older mtime wrong, but can t be sure Append to file, allocate indirect block, crash - Inode points to indirect block - But indirect block may contain garbage Must be careful about order of updates - Write new inode to disk before directory entry - Remove directory name before deallocating inode - Write cleared inode to disk before updating CG free map Solution: Many meta updates syncrhonous - Doing one write at a time ensures ordering - Of course, this hurts performance - Eg, untar much slower than disk bandwidth Note: Cannot update buffers on the disk queue - Eg, say you make two updates to same directory block - But crash recovery requires first to be synchronous - Must wait for first write to complete before doing second 13/36 14/36 Performance vs consistency First attempt: Ordered updates FFS crash recoverability comes at huge cost - Makes tasks such as untar easily times slower - All because you might lose power or reboot at any time Even while slowing ordinary usage, recovery slow - If fsck takes one minute, then disks get 10 bigger One solution: battery-backed RAM - Expensive (requires specialized hardware) - Often don t learn battery has died until too late - A pain if computer dies (can t just move disk) - If OS bug causes crash, RAM might be garbage Better solution: Advanced file system techniques - Topic of rest of lecture Must follow three rules in ordering updates: 1 Never write pointer before initializing the structure it points to 2 Never reuse a resource before nullifying all pointers to it 3 Never clear last pointer to live resource before setting new one If you do this, file system will be recoverable Moreover, can recover quickly - Might leak free disk space, but otherwise correct - So start running after reboot, scavenge for space in background How to achieve? - Keep a partial order on buffered blocks 15/36 16/36 Ordered updates (continued) Problem: Cyclic dependencies Example: Create file A - Block X contains an inode - Block Y contains a directory block - Create file A in inode block X, dir block Y We say Y X, pronounced Y depends on X - Means Y cannot be written before X is written - X is called the dependee, Y the depender Can delay both writes, so long as order preserved - Say you create a second file B in blocks X and Y - Only have to write each out once for both creates Suppose you create file A, unlink file B - Both files in same directory block & inode block Can t write directory until inode A initialized - Otherwise, after crash directory will point to bogus inode - Worse yet, same inode # might be re-allocated - So could end up with file name A being an unrelated file Can t write inode block until dir entry B cleared - Otherwise, B could end up with too small a link count - File could be deleted while links to it still exist Otherwise, fsck has to be very slow - Check every directory entry and inode link count 17/36 18/36 Cyclic dependencies illustrated = in use Original B,#5 = free Modified (a) Original Organization B,#5 More problems Crash might occur between ordered but related writes - Eg, summary information wrong after block freed Block aging - Block that always has dependency will never get written back Solution: Soft updates [Ganger] - Write blocks in any order - But keep track of dependencies - When writing a block, temporarily roll back any changes you can t yet commit to disk (b) Create File A (c) Remove file B 19/36 20/36 Breaking dependencies w rollback Breaking dependencies w rollback B,#5 (a) After Meta Updates Now say we decide to write directory block Can t write file name A to disk has dependee (b) Safe Version of Written Undo file A before writing dir block to disk - Even though we just wrote it, directory block still dirty But now inode block has no dependees - Can safely write inode block to disk as-is Breaking dependencies w rollback Breaking dependencies w rollback (c) Written Now inode block clean (same in memory as on disk) But have to write directory block a second time All stably on disk (d) Written Crash at any point would have been safe Soft updates Structure for each updated field or pointer, contains: - old value - new value - list of updates on which this update depends (dependees) Can write blocks in any order - But must temporarily undo updates with pending dependencies - Must lock rolled-back version so applications don t see it - Choose ordering based on disk arm scheduling Some dependencies better handled by postponing in-memory updates - Eg, when freeing block (eg, because file truncated), just mark block free in bitmap after block pointer cleared on disk Simple example Say you create a zero-length file A Depender: Directory entry for A - Can t be written untill dependees on disk Dependees: - Inode must be initialized before dir entry written - Bitmap must mark inode allocated before dir entry written Old value: empty directory entry New value: filename A, inode # Can write directory block to disk any time - Must substitute old value until inode & bitmap updated on disk - Once dir block on disk contains A, file fully created - Crash before A on disk, worst case might leak the inode 22/36 23/36 Operations requiring soft updates (1) Operations requiring soft updates (2) 1 Block allocation - Must write the disk block, the free map, & a pointer - block & free map must be written before pointer - Use Undo/redo on pointer (& possibly file size) 2 Block deallocation - Must write the cleared pointer & free map - Just update free map after pointer written to disk - Or just immediately update free map if pointer not on disk Say you quickly append block to file then truncate - You will know pointer to block not written because of the allocated dependency structure - So both operations together require no disk I/O! 3 Link addition (see simple example) - Must write the directory entry, inode, & free map (if new inode) - Inode and free map must be written before dir entry - Use undo/redo on i# in dir entry (ignore entries w i# 0) 4 Link removal - Must write directory entry, inode & free map (if nlinks==0) - Must decrement nlinks only after pointer cleared - Clear directory entry immediately - Decrement in-memory nlinks once pointer written - If directory entry was never written, decrement immediately (again will know by presence of dependency structure) Note: Quick create/delete requires no disk I/O 24/36 25/36 Soft update issues fsync sycall to flush file changes to disk - Must also flush directory entries, parent directories, etc unmount flush all changes to disk on shutdown - Some buffers must be flushed multiple times to get clean Deleting large directory trees frighteningly fast - unlink syscall returns even if inode/indir block not cached! - Dependencies allocated faster than blocks written - Cap # dependencies allocated to avoid exhausting memory Useless write-backs - Syncer flushes dirty buffers to disk every 30 seconds - Writing all at once means many dependencies unsatisfied - Fix syncer to write blocks one at a time - Fix LRU buffer eviction to know about dependencies Soft updates fsck Split into foreground and background parts Foreground must be done before remounting FS - Need to make sure per-cylinder summary info makes sense - Recompute free block/inode counts from bitmaps very fast - Will leave FS consistent, but might leak disk space Background does traditional fsck operations - Do after mounting to recuperate free space - Can be using the file system while this is happening - Must be done in forground after a media failure Difference from traditional FFS fsck: - May have many, many inodes with non-zero link counts - Don t stick them all in lost+found (unless media failure) 26/36 27/36 An alternative: Journaling Reserve a portion of disk for write-ahead log - Write any meta operation first to log, then to disk - After crash/reboot, re-play the log (efficient) - May re-do already committed change, but won t miss anything Performance advantage: - Log is consecutive portion of disk - Multiple log writes very fast (at disk b/w) - Consider updates committed when written to log Example: delete directory tree - Record all freed blocks, changed directory entries in log - Return control to user - Write out changed directories, bitmaps, etc in background (sort for good disk arm scheduling) Journaling details Must find oldest relevant log entry - Otherwise, redundant and slow to replay whole log Use checkpoints - Once all records up to log entry N have been processed and affected blocks stably committed to disk - Record N to disk either in reserved checkpoint location, or in checkpoint log record - Never need to go back before most recent checkpointed N Must also find end of log - Typically circular buffer; don t play old records out of order - Can include begin transaction/end transaction records - Also typically have checksum in case some sectors bad 28/36 29/36 Case study: XFS [Sweeney] Main idea: Think big - Big disks, files, large # of files, 64-bit everything - Yet maintain very good performance Break disk up into Allocation Groups (AGs) GB regions of disk - New directories go in new AGs - Within directory, inodes of files go in same AG - Unlike cylinder groups, AGs too large to minimize seek times - Unlike cylinder groups, no fixed # of inodes per AG Advantages of AGs: - Parallelize allocation of blocks/inodes on multiprocessor (independent locking of different free space structures) - Can use 32-bit block pointers within AGs (keeps structures smaller) 30/36 B+-trees p p p p tr K tr K tr K tr pkpkpkp pkpkpkp tr tr tr tr tr tr tr tr V V V V V V XFS makes extensive use of B+-trees - Indexed structure stores ordered Keys & Values - Keys must have an ordering defined on them - Stored in blocks for efficient disk access For B+-tree w n items, all operations O(log n): - Retrieve closest key, value to target key k - Insert a new key, value pair - Delete key, value pair 31/36 B+-trees continued See any algorithms book for details (eg, [Cormen]) Some operations on B-tree are complex: - Eg, insert item into completely full B+-tree - May require splitting nodes, adding new level to tree - Would be bad to crash & leave B+tree in inconsistent state Journal enables atomic complex operations - First write all changes to the log - If crash while writing log, incomplete log record will be discarded, and no change made - Otherwise, if crash while updating B+-tree, will replay entire log record and write everything B+-trees in XFS B+-trees are complex to implement - But once you ve done it, might as well use everywhere Use B+-trees for directories (keyed on filename hash) - Makes large directories efficient Use B+-trees for inodes - No more FFS-style fixed block pointers - Instead, B+-tree maps: file offset start block, # blocks - Ideally file is one or small number of contiguous extents - Allows small inodes & no indirect blocks even for huge files Use to find inode based on inumber - High bits of inumber specify AG - B+-tree in AG maps: starting i# block #, free-map - So free inodes tracked right in leaf of B+-tree 32/36 33/36 More B+-trees in XFS Free extents tracked by two B+-trees 1 start block # # free blocks 2 # free blocks start block # Use journal to update both atomically & consistently #1 allows you to coalesce adjacent free regions #1 allows you to allocate near some target - Eg, when extending file, put next block near previous one - When first writing to file, but near inode #2 allows you to do best fit allocation - Leave large free extents for large files Contiguous allocation Ideally want each file contiguous on disk - Sequential file I/O should be as fast as sequential disk I/O But how do you know how large a file will be? Idea: delayed allocation - write syscall only affects the buffer cache - Allow write into buffers before deciding where to place on disk - Assign disk space only when buffers are flushed Other advantages: - Short-lived files never need disk space allocated - mmaped files often written in random order in memory, but will be written to disk mostly contiguously - Write clustering: find other nearby stuff to write to disk 34/36 35/36 Journaling vs soft updates Both much better than FFS alone Some limitations of soft updates - Very specific to FFS structures (Eg, couldn t easily add B-trees like XFS even directory rename not quite right) - Meta updates may proceed out of order (Eg, create A, create B, crash maybe only B exists after reboot) - Still need slow background fsck to reclaim space Some limitations of journaling - write required for every meta operation (whereas create-then-delete might require no I/O w soft updates) - Possible contention for end of log on multi-processor - fsync must sync other operations meta to log, too 36/36
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