1.. SPDX-License-Identifier: GPL-2.0 2 3================== 4XFS Logging Design 5================== 6 7Preamble 8======== 9 10This document describes the design and algorithms that the XFS journalling 11subsystem is based on. This document describes the design and algorithms that 12the XFS journalling subsystem is based on so that readers may familiarize 13themselves with the general concepts of how transaction processing in XFS works. 14 15We begin with an overview of transactions in XFS, followed by describing how 16transaction reservations are structured and accounted, and then move into how we 17guarantee forwards progress for long running transactions with finite initial 18reservations bounds. At this point we need to explain how relogging works. With 19the basic concepts covered, the design of the delayed logging mechanism is 20documented. 21 22 23Introduction 24============ 25 26XFS uses Write Ahead Logging for ensuring changes to the filesystem metadata 27are atomic and recoverable. For reasons of space and time efficiency, the 28logging mechanisms are varied and complex, combining intents, logical and 29physical logging mechanisms to provide the necessary recovery guarantees the 30filesystem requires. 31 32Some objects, such as inodes and dquots, are logged in logical format where the 33details logged are made up of the changes to in-core structures rather than 34on-disk structures. Other objects - typically buffers - have their physical 35changes logged. Long running atomic modifications have individual changes 36chained together by intents, ensuring that journal recovery can restart and 37finish an operation that was only partially done when the system stopped 38functioning. 39 40The reason for these differences is to keep the amount of log space and CPU time 41required to process objects being modified as small as possible and hence the 42logging overhead as low as possible. Some items are very frequently modified, 43and some parts of objects are more frequently modified than others, so keeping 44the overhead of metadata logging low is of prime importance. 45 46The method used to log an item or chain modifications together isn't 47particularly important in the scope of this document. It suffices to know that 48the method used for logging a particular object or chaining modifications 49together are different and are dependent on the object and/or modification being 50performed. The logging subsystem only cares that certain specific rules are 51followed to guarantee forwards progress and prevent deadlocks. 52 53 54Transactions in XFS 55=================== 56 57XFS has two types of high level transactions, defined by the type of log space 58reservation they take. These are known as "one shot" and "permanent" 59transactions. Permanent transaction reservations can take reservations that span 60commit boundaries, whilst "one shot" transactions are for a single atomic 61modification. 62 63The type and size of reservation must be matched to the modification taking 64place. This means that permanent transactions can be used for one-shot 65modifications, but one-shot reservations cannot be used for permanent 66transactions. 67 68In the code, a one-shot transaction pattern looks somewhat like this:: 69 70 tp = xfs_trans_alloc(<reservation>) 71 <lock items> 72 <join item to transaction> 73 <do modification> 74 xfs_trans_commit(tp); 75 76As items are modified in the transaction, the dirty regions in those items are 77tracked via the transaction handle. Once the transaction is committed, all 78resources joined to it are released, along with the remaining unused reservation 79space that was taken at the transaction allocation time. 80 81In contrast, a permanent transaction is made up of multiple linked individual 82transactions, and the pattern looks like this:: 83 84 tp = xfs_trans_alloc(<reservation>) 85 xfs_ilock(ip, XFS_ILOCK_EXCL) 86 87 loop { 88 xfs_trans_ijoin(tp, 0); 89 <do modification> 90 xfs_trans_log_inode(tp, ip); 91 xfs_trans_roll(&tp); 92 } 93 94 xfs_trans_commit(tp); 95 xfs_iunlock(ip, XFS_ILOCK_EXCL); 96 97While this might look similar to a one-shot transaction, there is an important 98difference: xfs_trans_roll() performs a specific operation that links two 99transactions together:: 100 101 ntp = xfs_trans_dup(tp); 102 xfs_trans_commit(tp); 103 xfs_log_reserve(ntp); 104 105This results in a series of "rolling transactions" where the inode is locked 106across the entire chain of transactions. Hence while this series of rolling 107transactions is running, nothing else can read from or write to the inode and 108this provides a mechanism for complex changes to appear atomic from an external 109observer's point of view. 110 111It is important to note that a series of rolling transactions in a permanent 112transaction does not form an atomic change in the journal. While each 113individual modification is atomic, the chain is *not atomic*. If we crash half 114way through, then recovery will only replay up to the last transactional 115modification the loop made that was committed to the journal. 116 117This affects long running permanent transactions in that it is not possible to 118predict how much of a long running operation will actually be recovered because 119there is no guarantee of how much of the operation reached stale storage. Hence 120if a long running operation requires multiple transactions to fully complete, 121the high level operation must use intents and deferred operations to guarantee 122recovery can complete the operation once the first transactions is persisted in 123the on-disk journal. 124 125 126Transactions are Asynchronous 127============================= 128 129In XFS, all high level transactions are asynchronous by default. This means that 130xfs_trans_commit() does not guarantee that the modification has been committed 131to stable storage when it returns. Hence when a system crashes, not all the 132completed transactions will be replayed during recovery. 133 134However, the logging subsystem does provide global ordering guarantees, such 135that if a specific change is seen after recovery, all metadata modifications 136that were committed prior to that change will also be seen. 137 138For single shot operations that need to reach stable storage immediately, or 139ensuring that a long running permanent transaction is fully committed once it is 140complete, we can explicitly tag a transaction as synchronous. This will trigger 141a "log force" to flush the outstanding committed transactions to stable storage 142in the journal and wait for that to complete. 143 144Synchronous transactions are rarely used, however, because they limit logging 145throughput to the IO latency limitations of the underlying storage. Instead, we 146tend to use log forces to ensure modifications are on stable storage only when 147a user operation requires a synchronisation point to occur (e.g. fsync). 148 149 150Transaction Reservations 151======================== 152 153It has been mentioned a number of times now that the logging subsystem needs to 154provide a forwards progress guarantee so that no modification ever stalls 155because it can't be written to the journal due to a lack of space in the 156journal. This is achieved by the transaction reservations that are made when 157a transaction is first allocated. For permanent transactions, these reservations 158are maintained as part of the transaction rolling mechanism. 159 160A transaction reservation provides a guarantee that there is physical log space 161available to write the modification into the journal before we start making 162modifications to objects and items. As such, the reservation needs to be large 163enough to take into account the amount of metadata that the change might need to 164log in the worst case. This means that if we are modifying a btree in the 165transaction, we have to reserve enough space to record a full leaf-to-root split 166of the btree. As such, the reservations are quite complex because we have to 167take into account all the hidden changes that might occur. 168 169For example, a user data extent allocation involves allocating an extent from 170free space, which modifies the free space trees. That's two btrees. Inserting 171the extent into the inode's extent map might require a split of the extent map 172btree, which requires another allocation that can modify the free space trees 173again. Then we might have to update reverse mappings, which modifies yet 174another btree which might require more space. And so on. Hence the amount of 175metadata that a "simple" operation can modify can be quite large. 176 177This "worst case" calculation provides us with the static "unit reservation" 178for the transaction that is calculated at mount time. We must guarantee that the 179log has this much space available before the transaction is allowed to proceed 180so that when we come to write the dirty metadata into the log we don't run out 181of log space half way through the write. 182 183For one-shot transactions, a single unit space reservation is all that is 184required for the transaction to proceed. For permanent transactions, however, we 185also have a "log count" that affects the size of the reservation that is to be 186made. 187 188While a permanent transaction can get by with a single unit of space 189reservation, it is somewhat inefficient to do this as it requires the 190transaction rolling mechanism to re-reserve space on every transaction roll. We 191know from the implementation of the permanent transactions how many transaction 192rolls are likely for the common modifications that need to be made. 193 194For example, and inode allocation is typically two transactions - one to 195physically allocate a free inode chunk on disk, and another to allocate an inode 196from an inode chunk that has free inodes in it. Hence for an inode allocation 197transaction, we might set the reservation log count to a value of 2 to indicate 198that the common/fast path transaction will commit two linked transactions in a 199chain. Each time a permanent transaction rolls, it consumes an entire unit 200reservation. 201 202Hence when the permanent transaction is first allocated, the log space 203reservation is increases from a single unit reservation to multiple unit 204reservations. That multiple is defined by the reservation log count, and this 205means we can roll the transaction multiple times before we have to re-reserve 206log space when we roll the transaction. This ensures that the common 207modifications we make only need to reserve log space once. 208 209If the log count for a permanent transaction reaches zero, then it needs to 210re-reserve physical space in the log. This is somewhat complex, and requires 211an understanding of how the log accounts for space that has been reserved. 212 213 214Log Space Accounting 215==================== 216 217The position in the log is typically referred to as a Log Sequence Number (LSN). 218The log is circular, so the positions in the log are defined by the combination 219of a cycle number - the number of times the log has been overwritten - and the 220offset into the log. A LSN carries the cycle in the upper 32 bits and the 221offset in the lower 32 bits. The offset is in units of "basic blocks" (512 222bytes). Hence we can do realtively simple LSN based math to keep track of 223available space in the log. 224 225Log space accounting is done via a pair of constructs called "grant heads". The 226position of the grant heads is an absolute value, so the amount of space 227available in the log is defined by the distance between the position of the 228grant head and the current log tail. That is, how much space can be 229reserved/consumed before the grant heads would fully wrap the log and overtake 230the tail position. 231 232The first grant head is the "reserve" head. This tracks the byte count of the 233reservations currently held by active transactions. It is a purely in-memory 234accounting of the space reservation and, as such, actually tracks byte offsets 235into the log rather than basic blocks. Hence it technically isn't using LSNs to 236represent the log position, but it is still treated like a split {cycle,offset} 237tuple for the purposes of tracking reservation space. 238 239The reserve grant head is used to accurately account for exact transaction 240reservations amounts and the exact byte count that modifications actually make 241and need to write into the log. The reserve head is used to prevent new 242transactions from taking new reservations when the head reaches the current 243tail. It will block new reservations in a FIFO queue and as the log tail moves 244forward it will wake them in order once sufficient space is available. This FIFO 245mechanism ensures no transaction is starved of resources when log space 246shortages occur. 247 248The other grant head is the "write" head. Unlike the reserve head, this grant 249head contains an LSN and it tracks the physical space usage in the log. While 250this might sound like it is accounting the same state as the reserve grant head 251- and it mostly does track exactly the same location as the reserve grant head - 252there are critical differences in behaviour between them that provides the 253forwards progress guarantees that rolling permanent transactions require. 254 255These differences when a permanent transaction is rolled and the internal "log 256count" reaches zero and the initial set of unit reservations have been 257exhausted. At this point, we still require a log space reservation to continue 258the next transaction in the sequeunce, but we have none remaining. We cannot 259sleep during the transaction commit process waiting for new log space to become 260available, as we may end up on the end of the FIFO queue and the items we have 261locked while we sleep could end up pinning the tail of the log before there is 262enough free space in the log to fulfil all of the pending reservations and 263then wake up transaction commit in progress. 264 265To take a new reservation without sleeping requires us to be able to take a 266reservation even if there is no reservation space currently available. That is, 267we need to be able to *overcommit* the log reservation space. As has already 268been detailed, we cannot overcommit physical log space. However, the reserve 269grant head does not track physical space - it only accounts for the amount of 270reservations we currently have outstanding. Hence if the reserve head passes 271over the tail of the log all it means is that new reservations will be throttled 272immediately and remain throttled until the log tail is moved forward far enough 273to remove the overcommit and start taking new reservations. In other words, we 274can overcommit the reserve head without violating the physical log head and tail 275rules. 276 277As a result, permanent transactions only "regrant" reservation space during 278xfs_trans_commit() calls, while the physical log space reservation - tracked by 279the write head - is then reserved separately by a call to xfs_log_reserve() 280after the commit completes. Once the commit completes, we can sleep waiting for 281physical log space to be reserved from the write grant head, but only if one 282critical rule has been observed:: 283 284 Code using permanent reservations must always log the items they hold 285 locked across each transaction they roll in the chain. 286 287"Re-logging" the locked items on every transaction roll ensures that the items 288attached to the transaction chain being rolled are always relocated to the 289physical head of the log and so do not pin the tail of the log. If a locked item 290pins the tail of the log when we sleep on the write reservation, then we will 291deadlock the log as we cannot take the locks needed to write back that item and 292move the tail of the log forwards to free up write grant space. Re-logging the 293locked items avoids this deadlock and guarantees that the log reservation we are 294making cannot self-deadlock. 295 296If all rolling transactions obey this rule, then they can all make forwards 297progress independently because nothing will block the progress of the log 298tail moving forwards and hence ensuring that write grant space is always 299(eventually) made available to permanent transactions no matter how many times 300they roll. 301 302 303Re-logging Explained 304==================== 305 306XFS allows multiple separate modifications to a single object to be carried in 307the log at any given time. This allows the log to avoid needing to flush each 308change to disk before recording a new change to the object. XFS does this via a 309method called "re-logging". Conceptually, this is quite simple - all it requires 310is that any new change to the object is recorded with a *new copy* of all the 311existing changes in the new transaction that is written to the log. 312 313That is, if we have a sequence of changes A through to F, and the object was 314written to disk after change D, we would see in the log the following series 315of transactions, their contents and the log sequence number (LSN) of the 316transaction:: 317 318 Transaction Contents LSN 319 A A X 320 B A+B X+n 321 C A+B+C X+n+m 322 D A+B+C+D X+n+m+o 323 <object written to disk> 324 E E Y (> X+n+m+o) 325 F E+F Y+p 326 327In other words, each time an object is relogged, the new transaction contains 328the aggregation of all the previous changes currently held only in the log. 329 330This relogging technique allows objects to be moved forward in the log so that 331an object being relogged does not prevent the tail of the log from ever moving 332forward. This can be seen in the table above by the changing (increasing) LSN 333of each subsequent transaction, and it's the technique that allows us to 334implement long-running, multiple-commit permanent transactions. 335 336A typical example of a rolling transaction is the removal of extents from an 337inode which can only be done at a rate of two extents per transaction because 338of reservation size limitations. Hence a rolling extent removal transaction 339keeps relogging the inode and btree buffers as they get modified in each 340removal operation. This keeps them moving forward in the log as the operation 341progresses, ensuring that current operation never gets blocked by itself if the 342log wraps around. 343 344Hence it can be seen that the relogging operation is fundamental to the correct 345working of the XFS journalling subsystem. From the above description, most 346people should be able to see why the XFS metadata operations writes so much to 347the log - repeated operations to the same objects write the same changes to 348the log over and over again. Worse is the fact that objects tend to get 349dirtier as they get relogged, so each subsequent transaction is writing more 350metadata into the log. 351 352It should now also be obvious how relogging and asynchronous transactions go 353hand in hand. That is, transactions don't get written to the physical journal 354until either a log buffer is filled (a log buffer can hold multiple 355transactions) or a synchronous operation forces the log buffers holding the 356transactions to disk. This means that XFS is doing aggregation of transactions 357in memory - batching them, if you like - to minimise the impact of the log IO on 358transaction throughput. 359 360The limitation on asynchronous transaction throughput is the number and size of 361log buffers made available by the log manager. By default there are 8 log 362buffers available and the size of each is 32kB - the size can be increased up 363to 256kB by use of a mount option. 364 365Effectively, this gives us the maximum bound of outstanding metadata changes 366that can be made to the filesystem at any point in time - if all the log 367buffers are full and under IO, then no more transactions can be committed until 368the current batch completes. It is now common for a single current CPU core to 369be to able to issue enough transactions to keep the log buffers full and under 370IO permanently. Hence the XFS journalling subsystem can be considered to be IO 371bound. 372 373Delayed Logging: Concepts 374========================= 375 376The key thing to note about the asynchronous logging combined with the 377relogging technique XFS uses is that we can be relogging changed objects 378multiple times before they are committed to disk in the log buffers. If we 379return to the previous relogging example, it is entirely possible that 380transactions A through D are committed to disk in the same log buffer. 381 382That is, a single log buffer may contain multiple copies of the same object, 383but only one of those copies needs to be there - the last one "D", as it 384contains all the changes from the previous changes. In other words, we have one 385necessary copy in the log buffer, and three stale copies that are simply 386wasting space. When we are doing repeated operations on the same set of 387objects, these "stale objects" can be over 90% of the space used in the log 388buffers. It is clear that reducing the number of stale objects written to the 389log would greatly reduce the amount of metadata we write to the log, and this 390is the fundamental goal of delayed logging. 391 392From a conceptual point of view, XFS is already doing relogging in memory (where 393memory == log buffer), only it is doing it extremely inefficiently. It is using 394logical to physical formatting to do the relogging because there is no 395infrastructure to keep track of logical changes in memory prior to physically 396formatting the changes in a transaction to the log buffer. Hence we cannot avoid 397accumulating stale objects in the log buffers. 398 399Delayed logging is the name we've given to keeping and tracking transactional 400changes to objects in memory outside the log buffer infrastructure. Because of 401the relogging concept fundamental to the XFS journalling subsystem, this is 402actually relatively easy to do - all the changes to logged items are already 403tracked in the current infrastructure. The big problem is how to accumulate 404them and get them to the log in a consistent, recoverable manner. 405Describing the problems and how they have been solved is the focus of this 406document. 407 408One of the key changes that delayed logging makes to the operation of the 409journalling subsystem is that it disassociates the amount of outstanding 410metadata changes from the size and number of log buffers available. In other 411words, instead of there only being a maximum of 2MB of transaction changes not 412written to the log at any point in time, there may be a much greater amount 413being accumulated in memory. Hence the potential for loss of metadata on a 414crash is much greater than for the existing logging mechanism. 415 416It should be noted that this does not change the guarantee that log recovery 417will result in a consistent filesystem. What it does mean is that as far as the 418recovered filesystem is concerned, there may be many thousands of transactions 419that simply did not occur as a result of the crash. This makes it even more 420important that applications that care about their data use fsync() where they 421need to ensure application level data integrity is maintained. 422 423It should be noted that delayed logging is not an innovative new concept that 424warrants rigorous proofs to determine whether it is correct or not. The method 425of accumulating changes in memory for some period before writing them to the 426log is used effectively in many filesystems including ext3 and ext4. Hence 427no time is spent in this document trying to convince the reader that the 428concept is sound. Instead it is simply considered a "solved problem" and as 429such implementing it in XFS is purely an exercise in software engineering. 430 431The fundamental requirements for delayed logging in XFS are simple: 432 433 1. Reduce the amount of metadata written to the log by at least 434 an order of magnitude. 435 2. Supply sufficient statistics to validate Requirement #1. 436 3. Supply sufficient new tracing infrastructure to be able to debug 437 problems with the new code. 438 4. No on-disk format change (metadata or log format). 439 5. Enable and disable with a mount option. 440 6. No performance regressions for synchronous transaction workloads. 441 442Delayed Logging: Design 443======================= 444 445Storing Changes 446--------------- 447 448The problem with accumulating changes at a logical level (i.e. just using the 449existing log item dirty region tracking) is that when it comes to writing the 450changes to the log buffers, we need to ensure that the object we are formatting 451is not changing while we do this. This requires locking the object to prevent 452concurrent modification. Hence flushing the logical changes to the log would 453require us to lock every object, format them, and then unlock them again. 454 455This introduces lots of scope for deadlocks with transactions that are already 456running. For example, a transaction has object A locked and modified, but needs 457the delayed logging tracking lock to commit the transaction. However, the 458flushing thread has the delayed logging tracking lock already held, and is 459trying to get the lock on object A to flush it to the log buffer. This appears 460to be an unsolvable deadlock condition, and it was solving this problem that 461was the barrier to implementing delayed logging for so long. 462 463The solution is relatively simple - it just took a long time to recognise it. 464Put simply, the current logging code formats the changes to each item into an 465vector array that points to the changed regions in the item. The log write code 466simply copies the memory these vectors point to into the log buffer during 467transaction commit while the item is locked in the transaction. Instead of 468using the log buffer as the destination of the formatting code, we can use an 469allocated memory buffer big enough to fit the formatted vector. 470 471If we then copy the vector into the memory buffer and rewrite the vector to 472point to the memory buffer rather than the object itself, we now have a copy of 473the changes in a format that is compatible with the log buffer writing code. 474that does not require us to lock the item to access. This formatting and 475rewriting can all be done while the object is locked during transaction commit, 476resulting in a vector that is transactionally consistent and can be accessed 477without needing to lock the owning item. 478 479Hence we avoid the need to lock items when we need to flush outstanding 480asynchronous transactions to the log. The differences between the existing 481formatting method and the delayed logging formatting can be seen in the 482diagram below. 483 484Current format log vector:: 485 486 Object +---------------------------------------------+ 487 Vector 1 +----+ 488 Vector 2 +----+ 489 Vector 3 +----------+ 490 491After formatting:: 492 493 Log Buffer +-V1-+-V2-+----V3----+ 494 495Delayed logging vector:: 496 497 Object +---------------------------------------------+ 498 Vector 1 +----+ 499 Vector 2 +----+ 500 Vector 3 +----------+ 501 502After formatting:: 503 504 Memory Buffer +-V1-+-V2-+----V3----+ 505 Vector 1 +----+ 506 Vector 2 +----+ 507 Vector 3 +----------+ 508 509The memory buffer and associated vector need to be passed as a single object, 510but still need to be associated with the parent object so if the object is 511relogged we can replace the current memory buffer with a new memory buffer that 512contains the latest changes. 513 514The reason for keeping the vector around after we've formatted the memory 515buffer is to support splitting vectors across log buffer boundaries correctly. 516If we don't keep the vector around, we do not know where the region boundaries 517are in the item, so we'd need a new encapsulation method for regions in the log 518buffer writing (i.e. double encapsulation). This would be an on-disk format 519change and as such is not desirable. It also means we'd have to write the log 520region headers in the formatting stage, which is problematic as there is per 521region state that needs to be placed into the headers during the log write. 522 523Hence we need to keep the vector, but by attaching the memory buffer to it and 524rewriting the vector addresses to point at the memory buffer we end up with a 525self-describing object that can be passed to the log buffer write code to be 526handled in exactly the same manner as the existing log vectors are handled. 527Hence we avoid needing a new on-disk format to handle items that have been 528relogged in memory. 529 530 531Tracking Changes 532---------------- 533 534Now that we can record transactional changes in memory in a form that allows 535them to be used without limitations, we need to be able to track and accumulate 536them so that they can be written to the log at some later point in time. The 537log item is the natural place to store this vector and buffer, and also makes sense 538to be the object that is used to track committed objects as it will always 539exist once the object has been included in a transaction. 540 541The log item is already used to track the log items that have been written to 542the log but not yet written to disk. Such log items are considered "active" 543and as such are stored in the Active Item List (AIL) which is a LSN-ordered 544double linked list. Items are inserted into this list during log buffer IO 545completion, after which they are unpinned and can be written to disk. An object 546that is in the AIL can be relogged, which causes the object to be pinned again 547and then moved forward in the AIL when the log buffer IO completes for that 548transaction. 549 550Essentially, this shows that an item that is in the AIL can still be modified 551and relogged, so any tracking must be separate to the AIL infrastructure. As 552such, we cannot reuse the AIL list pointers for tracking committed items, nor 553can we store state in any field that is protected by the AIL lock. Hence the 554committed item tracking needs it's own locks, lists and state fields in the log 555item. 556 557Similar to the AIL, tracking of committed items is done through a new list 558called the Committed Item List (CIL). The list tracks log items that have been 559committed and have formatted memory buffers attached to them. It tracks objects 560in transaction commit order, so when an object is relogged it is removed from 561it's place in the list and re-inserted at the tail. This is entirely arbitrary 562and done to make it easy for debugging - the last items in the list are the 563ones that are most recently modified. Ordering of the CIL is not necessary for 564transactional integrity (as discussed in the next section) so the ordering is 565done for convenience/sanity of the developers. 566 567 568Delayed Logging: Checkpoints 569---------------------------- 570 571When we have a log synchronisation event, commonly known as a "log force", 572all the items in the CIL must be written into the log via the log buffers. 573We need to write these items in the order that they exist in the CIL, and they 574need to be written as an atomic transaction. The need for all the objects to be 575written as an atomic transaction comes from the requirements of relogging and 576log replay - all the changes in all the objects in a given transaction must 577either be completely replayed during log recovery, or not replayed at all. If 578a transaction is not replayed because it is not complete in the log, then 579no later transactions should be replayed, either. 580 581To fulfill this requirement, we need to write the entire CIL in a single log 582transaction. Fortunately, the XFS log code has no fixed limit on the size of a 583transaction, nor does the log replay code. The only fundamental limit is that 584the transaction cannot be larger than just under half the size of the log. The 585reason for this limit is that to find the head and tail of the log, there must 586be at least one complete transaction in the log at any given time. If a 587transaction is larger than half the log, then there is the possibility that a 588crash during the write of a such a transaction could partially overwrite the 589only complete previous transaction in the log. This will result in a recovery 590failure and an inconsistent filesystem and hence we must enforce the maximum 591size of a checkpoint to be slightly less than a half the log. 592 593Apart from this size requirement, a checkpoint transaction looks no different 594to any other transaction - it contains a transaction header, a series of 595formatted log items and a commit record at the tail. From a recovery 596perspective, the checkpoint transaction is also no different - just a lot 597bigger with a lot more items in it. The worst case effect of this is that we 598might need to tune the recovery transaction object hash size. 599 600Because the checkpoint is just another transaction and all the changes to log 601items are stored as log vectors, we can use the existing log buffer writing 602code to write the changes into the log. To do this efficiently, we need to 603minimise the time we hold the CIL locked while writing the checkpoint 604transaction. The current log write code enables us to do this easily with the 605way it separates the writing of the transaction contents (the log vectors) from 606the transaction commit record, but tracking this requires us to have a 607per-checkpoint context that travels through the log write process through to 608checkpoint completion. 609 610Hence a checkpoint has a context that tracks the state of the current 611checkpoint from initiation to checkpoint completion. A new context is initiated 612at the same time a checkpoint transaction is started. That is, when we remove 613all the current items from the CIL during a checkpoint operation, we move all 614those changes into the current checkpoint context. We then initialise a new 615context and attach that to the CIL for aggregation of new transactions. 616 617This allows us to unlock the CIL immediately after transfer of all the 618committed items and effectively allow new transactions to be issued while we 619are formatting the checkpoint into the log. It also allows concurrent 620checkpoints to be written into the log buffers in the case of log force heavy 621workloads, just like the existing transaction commit code does. This, however, 622requires that we strictly order the commit records in the log so that 623checkpoint sequence order is maintained during log replay. 624 625To ensure that we can be writing an item into a checkpoint transaction at 626the same time another transaction modifies the item and inserts the log item 627into the new CIL, then checkpoint transaction commit code cannot use log items 628to store the list of log vectors that need to be written into the transaction. 629Hence log vectors need to be able to be chained together to allow them to be 630detached from the log items. That is, when the CIL is flushed the memory 631buffer and log vector attached to each log item needs to be attached to the 632checkpoint context so that the log item can be released. In diagrammatic form, 633the CIL would look like this before the flush:: 634 635 CIL Head 636 | 637 V 638 Log Item <-> log vector 1 -> memory buffer 639 | -> vector array 640 V 641 Log Item <-> log vector 2 -> memory buffer 642 | -> vector array 643 V 644 ...... 645 | 646 V 647 Log Item <-> log vector N-1 -> memory buffer 648 | -> vector array 649 V 650 Log Item <-> log vector N -> memory buffer 651 -> vector array 652 653And after the flush the CIL head is empty, and the checkpoint context log 654vector list would look like:: 655 656 Checkpoint Context 657 | 658 V 659 log vector 1 -> memory buffer 660 | -> vector array 661 | -> Log Item 662 V 663 log vector 2 -> memory buffer 664 | -> vector array 665 | -> Log Item 666 V 667 ...... 668 | 669 V 670 log vector N-1 -> memory buffer 671 | -> vector array 672 | -> Log Item 673 V 674 log vector N -> memory buffer 675 -> vector array 676 -> Log Item 677 678Once this transfer is done, the CIL can be unlocked and new transactions can 679start, while the checkpoint flush code works over the log vector chain to 680commit the checkpoint. 681 682Once the checkpoint is written into the log buffers, the checkpoint context is 683attached to the log buffer that the commit record was written to along with a 684completion callback. Log IO completion will call that callback, which can then 685run transaction committed processing for the log items (i.e. insert into AIL 686and unpin) in the log vector chain and then free the log vector chain and 687checkpoint context. 688 689Discussion Point: I am uncertain as to whether the log item is the most 690efficient way to track vectors, even though it seems like the natural way to do 691it. The fact that we walk the log items (in the CIL) just to chain the log 692vectors and break the link between the log item and the log vector means that 693we take a cache line hit for the log item list modification, then another for 694the log vector chaining. If we track by the log vectors, then we only need to 695break the link between the log item and the log vector, which means we should 696dirty only the log item cachelines. Normally I wouldn't be concerned about one 697vs two dirty cachelines except for the fact I've seen upwards of 80,000 log 698vectors in one checkpoint transaction. I'd guess this is a "measure and 699compare" situation that can be done after a working and reviewed implementation 700is in the dev tree.... 701 702Delayed Logging: Checkpoint Sequencing 703-------------------------------------- 704 705One of the key aspects of the XFS transaction subsystem is that it tags 706committed transactions with the log sequence number of the transaction commit. 707This allows transactions to be issued asynchronously even though there may be 708future operations that cannot be completed until that transaction is fully 709committed to the log. In the rare case that a dependent operation occurs (e.g. 710re-using a freed metadata extent for a data extent), a special, optimised log 711force can be issued to force the dependent transaction to disk immediately. 712 713To do this, transactions need to record the LSN of the commit record of the 714transaction. This LSN comes directly from the log buffer the transaction is 715written into. While this works just fine for the existing transaction 716mechanism, it does not work for delayed logging because transactions are not 717written directly into the log buffers. Hence some other method of sequencing 718transactions is required. 719 720As discussed in the checkpoint section, delayed logging uses per-checkpoint 721contexts, and as such it is simple to assign a sequence number to each 722checkpoint. Because the switching of checkpoint contexts must be done 723atomically, it is simple to ensure that each new context has a monotonically 724increasing sequence number assigned to it without the need for an external 725atomic counter - we can just take the current context sequence number and add 726one to it for the new context. 727 728Then, instead of assigning a log buffer LSN to the transaction commit LSN 729during the commit, we can assign the current checkpoint sequence. This allows 730operations that track transactions that have not yet completed know what 731checkpoint sequence needs to be committed before they can continue. As a 732result, the code that forces the log to a specific LSN now needs to ensure that 733the log forces to a specific checkpoint. 734 735To ensure that we can do this, we need to track all the checkpoint contexts 736that are currently committing to the log. When we flush a checkpoint, the 737context gets added to a "committing" list which can be searched. When a 738checkpoint commit completes, it is removed from the committing list. Because 739the checkpoint context records the LSN of the commit record for the checkpoint, 740we can also wait on the log buffer that contains the commit record, thereby 741using the existing log force mechanisms to execute synchronous forces. 742 743It should be noted that the synchronous forces may need to be extended with 744mitigation algorithms similar to the current log buffer code to allow 745aggregation of multiple synchronous transactions if there are already 746synchronous transactions being flushed. Investigation of the performance of the 747current design is needed before making any decisions here. 748 749The main concern with log forces is to ensure that all the previous checkpoints 750are also committed to disk before the one we need to wait for. Therefore we 751need to check that all the prior contexts in the committing list are also 752complete before waiting on the one we need to complete. We do this 753synchronisation in the log force code so that we don't need to wait anywhere 754else for such serialisation - it only matters when we do a log force. 755 756The only remaining complexity is that a log force now also has to handle the 757case where the forcing sequence number is the same as the current context. That 758is, we need to flush the CIL and potentially wait for it to complete. This is a 759simple addition to the existing log forcing code to check the sequence numbers 760and push if required. Indeed, placing the current sequence checkpoint flush in 761the log force code enables the current mechanism for issuing synchronous 762transactions to remain untouched (i.e. commit an asynchronous transaction, then 763force the log at the LSN of that transaction) and so the higher level code 764behaves the same regardless of whether delayed logging is being used or not. 765 766Delayed Logging: Checkpoint Log Space Accounting 767------------------------------------------------ 768 769The big issue for a checkpoint transaction is the log space reservation for the 770transaction. We don't know how big a checkpoint transaction is going to be 771ahead of time, nor how many log buffers it will take to write out, nor the 772number of split log vector regions are going to be used. We can track the 773amount of log space required as we add items to the commit item list, but we 774still need to reserve the space in the log for the checkpoint. 775 776A typical transaction reserves enough space in the log for the worst case space 777usage of the transaction. The reservation accounts for log record headers, 778transaction and region headers, headers for split regions, buffer tail padding, 779etc. as well as the actual space for all the changed metadata in the 780transaction. While some of this is fixed overhead, much of it is dependent on 781the size of the transaction and the number of regions being logged (the number 782of log vectors in the transaction). 783 784An example of the differences would be logging directory changes versus logging 785inode changes. If you modify lots of inode cores (e.g. ``chmod -R g+w *``), then 786there are lots of transactions that only contain an inode core and an inode log 787format structure. That is, two vectors totaling roughly 150 bytes. If we modify 78810,000 inodes, we have about 1.5MB of metadata to write in 20,000 vectors. Each 789vector is 12 bytes, so the total to be logged is approximately 1.75MB. In 790comparison, if we are logging full directory buffers, they are typically 4KB 791each, so we in 1.5MB of directory buffers we'd have roughly 400 buffers and a 792buffer format structure for each buffer - roughly 800 vectors or 1.51MB total 793space. From this, it should be obvious that a static log space reservation is 794not particularly flexible and is difficult to select the "optimal value" for 795all workloads. 796 797Further, if we are going to use a static reservation, which bit of the entire 798reservation does it cover? We account for space used by the transaction 799reservation by tracking the space currently used by the object in the CIL and 800then calculating the increase or decrease in space used as the object is 801relogged. This allows for a checkpoint reservation to only have to account for 802log buffer metadata used such as log header records. 803 804However, even using a static reservation for just the log metadata is 805problematic. Typically log record headers use at least 16KB of log space per 8061MB of log space consumed (512 bytes per 32k) and the reservation needs to be 807large enough to handle arbitrary sized checkpoint transactions. This 808reservation needs to be made before the checkpoint is started, and we need to 809be able to reserve the space without sleeping. For a 8MB checkpoint, we need a 810reservation of around 150KB, which is a non-trivial amount of space. 811 812A static reservation needs to manipulate the log grant counters - we can take a 813permanent reservation on the space, but we still need to make sure we refresh 814the write reservation (the actual space available to the transaction) after 815every checkpoint transaction completion. Unfortunately, if this space is not 816available when required, then the regrant code will sleep waiting for it. 817 818The problem with this is that it can lead to deadlocks as we may need to commit 819checkpoints to be able to free up log space (refer back to the description of 820rolling transactions for an example of this). Hence we *must* always have 821space available in the log if we are to use static reservations, and that is 822very difficult and complex to arrange. It is possible to do, but there is a 823simpler way. 824 825The simpler way of doing this is tracking the entire log space used by the 826items in the CIL and using this to dynamically calculate the amount of log 827space required by the log metadata. If this log metadata space changes as a 828result of a transaction commit inserting a new memory buffer into the CIL, then 829the difference in space required is removed from the transaction that causes 830the change. Transactions at this level will *always* have enough space 831available in their reservation for this as they have already reserved the 832maximal amount of log metadata space they require, and such a delta reservation 833will always be less than or equal to the maximal amount in the reservation. 834 835Hence we can grow the checkpoint transaction reservation dynamically as items 836are added to the CIL and avoid the need for reserving and regranting log space 837up front. This avoids deadlocks and removes a blocking point from the 838checkpoint flush code. 839 840As mentioned early, transactions can't grow to more than half the size of the 841log. Hence as part of the reservation growing, we need to also check the size 842of the reservation against the maximum allowed transaction size. If we reach 843the maximum threshold, we need to push the CIL to the log. This is effectively 844a "background flush" and is done on demand. This is identical to 845a CIL push triggered by a log force, only that there is no waiting for the 846checkpoint commit to complete. This background push is checked and executed by 847transaction commit code. 848 849If the transaction subsystem goes idle while we still have items in the CIL, 850they will be flushed by the periodic log force issued by the xfssyncd. This log 851force will push the CIL to disk, and if the transaction subsystem stays idle, 852allow the idle log to be covered (effectively marked clean) in exactly the same 853manner that is done for the existing logging method. A discussion point is 854whether this log force needs to be done more frequently than the current rate 855which is once every 30s. 856 857 858Delayed Logging: Log Item Pinning 859--------------------------------- 860 861Currently log items are pinned during transaction commit while the items are 862still locked. This happens just after the items are formatted, though it could 863be done any time before the items are unlocked. The result of this mechanism is 864that items get pinned once for every transaction that is committed to the log 865buffers. Hence items that are relogged in the log buffers will have a pin count 866for every outstanding transaction they were dirtied in. When each of these 867transactions is completed, they will unpin the item once. As a result, the item 868only becomes unpinned when all the transactions complete and there are no 869pending transactions. Thus the pinning and unpinning of a log item is symmetric 870as there is a 1:1 relationship with transaction commit and log item completion. 871 872For delayed logging, however, we have an asymmetric transaction commit to 873completion relationship. Every time an object is relogged in the CIL it goes 874through the commit process without a corresponding completion being registered. 875That is, we now have a many-to-one relationship between transaction commit and 876log item completion. The result of this is that pinning and unpinning of the 877log items becomes unbalanced if we retain the "pin on transaction commit, unpin 878on transaction completion" model. 879 880To keep pin/unpin symmetry, the algorithm needs to change to a "pin on 881insertion into the CIL, unpin on checkpoint completion". In other words, the 882pinning and unpinning becomes symmetric around a checkpoint context. We have to 883pin the object the first time it is inserted into the CIL - if it is already in 884the CIL during a transaction commit, then we do not pin it again. Because there 885can be multiple outstanding checkpoint contexts, we can still see elevated pin 886counts, but as each checkpoint completes the pin count will retain the correct 887value according to it's context. 888 889Just to make matters more slightly more complex, this checkpoint level context 890for the pin count means that the pinning of an item must take place under the 891CIL commit/flush lock. If we pin the object outside this lock, we cannot 892guarantee which context the pin count is associated with. This is because of 893the fact pinning the item is dependent on whether the item is present in the 894current CIL or not. If we don't pin the CIL first before we check and pin the 895object, we have a race with CIL being flushed between the check and the pin 896(or not pinning, as the case may be). Hence we must hold the CIL flush/commit 897lock to guarantee that we pin the items correctly. 898 899Delayed Logging: Concurrent Scalability 900--------------------------------------- 901 902A fundamental requirement for the CIL is that accesses through transaction 903commits must scale to many concurrent commits. The current transaction commit 904code does not break down even when there are transactions coming from 2048 905processors at once. The current transaction code does not go any faster than if 906there was only one CPU using it, but it does not slow down either. 907 908As a result, the delayed logging transaction commit code needs to be designed 909for concurrency from the ground up. It is obvious that there are serialisation 910points in the design - the three important ones are: 911 912 1. Locking out new transaction commits while flushing the CIL 913 2. Adding items to the CIL and updating item space accounting 914 3. Checkpoint commit ordering 915 916Looking at the transaction commit and CIL flushing interactions, it is clear 917that we have a many-to-one interaction here. That is, the only restriction on 918the number of concurrent transactions that can be trying to commit at once is 919the amount of space available in the log for their reservations. The practical 920limit here is in the order of several hundred concurrent transactions for a 921128MB log, which means that it is generally one per CPU in a machine. 922 923The amount of time a transaction commit needs to hold out a flush is a 924relatively long period of time - the pinning of log items needs to be done 925while we are holding out a CIL flush, so at the moment that means it is held 926across the formatting of the objects into memory buffers (i.e. while memcpy()s 927are in progress). Ultimately a two pass algorithm where the formatting is done 928separately to the pinning of objects could be used to reduce the hold time of 929the transaction commit side. 930 931Because of the number of potential transaction commit side holders, the lock 932really needs to be a sleeping lock - if the CIL flush takes the lock, we do not 933want every other CPU in the machine spinning on the CIL lock. Given that 934flushing the CIL could involve walking a list of tens of thousands of log 935items, it will get held for a significant time and so spin contention is a 936significant concern. Preventing lots of CPUs spinning doing nothing is the 937main reason for choosing a sleeping lock even though nothing in either the 938transaction commit or CIL flush side sleeps with the lock held. 939 940It should also be noted that CIL flushing is also a relatively rare operation 941compared to transaction commit for asynchronous transaction workloads - only 942time will tell if using a read-write semaphore for exclusion will limit 943transaction commit concurrency due to cache line bouncing of the lock on the 944read side. 945 946The second serialisation point is on the transaction commit side where items 947are inserted into the CIL. Because transactions can enter this code 948concurrently, the CIL needs to be protected separately from the above 949commit/flush exclusion. It also needs to be an exclusive lock but it is only 950held for a very short time and so a spin lock is appropriate here. It is 951possible that this lock will become a contention point, but given the short 952hold time once per transaction I think that contention is unlikely. 953 954The final serialisation point is the checkpoint commit record ordering code 955that is run as part of the checkpoint commit and log force sequencing. The code 956path that triggers a CIL flush (i.e. whatever triggers the log force) will enter 957an ordering loop after writing all the log vectors into the log buffers but 958before writing the commit record. This loop walks the list of committing 959checkpoints and needs to block waiting for checkpoints to complete their commit 960record write. As a result it needs a lock and a wait variable. Log force 961sequencing also requires the same lock, list walk, and blocking mechanism to 962ensure completion of checkpoints. 963 964These two sequencing operations can use the mechanism even though the 965events they are waiting for are different. The checkpoint commit record 966sequencing needs to wait until checkpoint contexts contain a commit LSN 967(obtained through completion of a commit record write) while log force 968sequencing needs to wait until previous checkpoint contexts are removed from 969the committing list (i.e. they've completed). A simple wait variable and 970broadcast wakeups (thundering herds) has been used to implement these two 971serialisation queues. They use the same lock as the CIL, too. If we see too 972much contention on the CIL lock, or too many context switches as a result of 973the broadcast wakeups these operations can be put under a new spinlock and 974given separate wait lists to reduce lock contention and the number of processes 975woken by the wrong event. 976 977 978Lifecycle Changes 979----------------- 980 981The existing log item life cycle is as follows:: 982 983 1. Transaction allocate 984 2. Transaction reserve 985 3. Lock item 986 4. Join item to transaction 987 If not already attached, 988 Allocate log item 989 Attach log item to owner item 990 Attach log item to transaction 991 5. Modify item 992 Record modifications in log item 993 6. Transaction commit 994 Pin item in memory 995 Format item into log buffer 996 Write commit LSN into transaction 997 Unlock item 998 Attach transaction to log buffer 999 1000 <log buffer IO dispatched> 1001 <log buffer IO completes> 1002 1003 7. Transaction completion 1004 Mark log item committed 1005 Insert log item into AIL 1006 Write commit LSN into log item 1007 Unpin log item 1008 8. AIL traversal 1009 Lock item 1010 Mark log item clean 1011 Flush item to disk 1012 1013 <item IO completion> 1014 1015 9. Log item removed from AIL 1016 Moves log tail 1017 Item unlocked 1018 1019Essentially, steps 1-6 operate independently from step 7, which is also 1020independent of steps 8-9. An item can be locked in steps 1-6 or steps 8-9 1021at the same time step 7 is occurring, but only steps 1-6 or 8-9 can occur 1022at the same time. If the log item is in the AIL or between steps 6 and 7 1023and steps 1-6 are re-entered, then the item is relogged. Only when steps 8-9 1024are entered and completed is the object considered clean. 1025 1026With delayed logging, there are new steps inserted into the life cycle:: 1027 1028 1. Transaction allocate 1029 2. Transaction reserve 1030 3. Lock item 1031 4. Join item to transaction 1032 If not already attached, 1033 Allocate log item 1034 Attach log item to owner item 1035 Attach log item to transaction 1036 5. Modify item 1037 Record modifications in log item 1038 6. Transaction commit 1039 Pin item in memory if not pinned in CIL 1040 Format item into log vector + buffer 1041 Attach log vector and buffer to log item 1042 Insert log item into CIL 1043 Write CIL context sequence into transaction 1044 Unlock item 1045 1046 <next log force> 1047 1048 7. CIL push 1049 lock CIL flush 1050 Chain log vectors and buffers together 1051 Remove items from CIL 1052 unlock CIL flush 1053 write log vectors into log 1054 sequence commit records 1055 attach checkpoint context to log buffer 1056 1057 <log buffer IO dispatched> 1058 <log buffer IO completes> 1059 1060 8. Checkpoint completion 1061 Mark log item committed 1062 Insert item into AIL 1063 Write commit LSN into log item 1064 Unpin log item 1065 9. AIL traversal 1066 Lock item 1067 Mark log item clean 1068 Flush item to disk 1069 <item IO completion> 1070 10. Log item removed from AIL 1071 Moves log tail 1072 Item unlocked 1073 1074From this, it can be seen that the only life cycle differences between the two 1075logging methods are in the middle of the life cycle - they still have the same 1076beginning and end and execution constraints. The only differences are in the 1077committing of the log items to the log itself and the completion processing. 1078Hence delayed logging should not introduce any constraints on log item 1079behaviour, allocation or freeing that don't already exist. 1080 1081As a result of this zero-impact "insertion" of delayed logging infrastructure 1082and the design of the internal structures to avoid on disk format changes, we 1083can basically switch between delayed logging and the existing mechanism with a 1084mount option. Fundamentally, there is no reason why the log manager would not 1085be able to swap methods automatically and transparently depending on load 1086characteristics, but this should not be necessary if delayed logging works as 1087designed. 1088