diff options
Diffstat (limited to 'Documentation/admin-guide/device-mapper/vdo-design.rst')
| -rw-r--r-- | Documentation/admin-guide/device-mapper/vdo-design.rst | 633 |
1 files changed, 633 insertions, 0 deletions
diff --git a/Documentation/admin-guide/device-mapper/vdo-design.rst b/Documentation/admin-guide/device-mapper/vdo-design.rst new file mode 100644 index 000000000000..3cd59decbec0 --- /dev/null +++ b/Documentation/admin-guide/device-mapper/vdo-design.rst @@ -0,0 +1,633 @@ +.. SPDX-License-Identifier: GPL-2.0-only + +================ +Design of dm-vdo +================ + +The dm-vdo (virtual data optimizer) target provides inline deduplication, +compression, zero-block elimination, and thin provisioning. A dm-vdo target +can be backed by up to 256TB of storage, and can present a logical size of +up to 4PB. This target was originally developed at Permabit Technology +Corp. starting in 2009. It was first released in 2013 and has been used in +production environments ever since. It was made open-source in 2017 after +Permabit was acquired by Red Hat. This document describes the design of +dm-vdo. For usage, see vdo.rst in the same directory as this file. + +Because deduplication rates fall drastically as the block size increases, a +vdo target has a maximum block size of 4K. However, it can achieve +deduplication rates of 254:1, i.e. up to 254 copies of a given 4K block can +reference a single 4K of actual storage. It can achieve compression rates +of 14:1. All zero blocks consume no storage at all. + +Theory of Operation +=================== + +The design of dm-vdo is based on the idea that deduplication is a two-part +problem. The first is to recognize duplicate data. The second is to avoid +storing multiple copies of those duplicates. Therefore, dm-vdo has two main +parts: a deduplication index (called UDS) that is used to discover +duplicate data, and a data store with a reference counted block map that +maps from logical block addresses to the actual storage location of the +data. + +Zones and Threading +------------------- + +Due to the complexity of data optimization, the number of metadata +structures involved in a single write operation to a vdo target is larger +than most other targets. Furthermore, because vdo must operate on small +block sizes in order to achieve good deduplication rates, acceptable +performance can only be achieved through parallelism. Therefore, vdo's +design attempts to be lock-free. + +Most of a vdo's main data structures are designed to be easily divided into +"zones" such that any given bio must only access a single zone of any zoned +structure. Safety with minimal locking is achieved by ensuring that during +normal operation, each zone is assigned to a specific thread, and only that +thread will access the portion of the data structure in that zone. +Associated with each thread is a work queue. Each bio is associated with a +request object (the "data_vio") which will be added to a work queue when +the next phase of its operation requires access to the structures in the +zone associated with that queue. + +Another way of thinking about this arrangement is that the work queue for +each zone has an implicit lock on the structures it manages for all its +operations, because vdo guarantees that no other thread will alter those +structures. + +Although each structure is divided into zones, this division is not +reflected in the on-disk representation of each data structure. Therefore, +the number of zones for each structure, and hence the number of threads, +can be reconfigured each time a vdo target is started. + +The Deduplication Index +----------------------- + +In order to identify duplicate data efficiently, vdo was designed to +leverage some common characteristics of duplicate data. From empirical +observations, we gathered two key insights. The first is that in most data +sets with significant amounts of duplicate data, the duplicates tend to +have temporal locality. When a duplicate appears, it is more likely that +other duplicates will be detected, and that those duplicates will have been +written at about the same time. This is why the index keeps records in +temporal order. The second insight is that new data is more likely to +duplicate recent data than it is to duplicate older data and in general, +there are diminishing returns to looking further back in time. Therefore, +when the index is full, it should cull its oldest records to make space for +new ones. Another important idea behind the design of the index is that the +ultimate goal of deduplication is to reduce storage costs. Since there is a +trade-off between the storage saved and the resources expended to achieve +those savings, vdo does not attempt to find every last duplicate block. It +is sufficient to find and eliminate most of the redundancy. + +Each block of data is hashed to produce a 16-byte block name. An index +record consists of this block name paired with the presumed location of +that data on the underlying storage. However, it is not possible to +guarantee that the index is accurate. In the most common case, this occurs +because it is too costly to update the index when a block is over-written +or discarded. Doing so would require either storing the block name along +with the blocks, which is difficult to do efficiently in block-based +storage, or reading and rehashing each block before overwriting it. +Inaccuracy can also result from a hash collision where two different blocks +have the same name. In practice, this is extremely unlikely, but because +vdo does not use a cryptographic hash, a malicious workload could be +constructed. Because of these inaccuracies, vdo treats the locations in the +index as hints, and reads each indicated block to verify that it is indeed +a duplicate before sharing the existing block with a new one. + +Records are collected into groups called chapters. New records are added to +the newest chapter, called the open chapter. This chapter is stored in a +format optimized for adding and modifying records, and the content of the +open chapter is not finalized until it runs out of space for new records. +When the open chapter fills up, it is closed and a new open chapter is +created to collect new records. + +Closing a chapter converts it to a different format which is optimized for +reading. The records are written to a series of record pages based on the +order in which they were received. This means that records with temporal +locality should be on a small number of pages, reducing the I/O required to +retrieve them. The chapter also compiles an index that indicates which +record page contains any given name. This index means that a request for a +name can determine exactly which record page may contain that record, +without having to load the entire chapter from storage. This index uses +only a subset of the block name as its key, so it cannot guarantee that an +index entry refers to the desired block name. It can only guarantee that if +there is a record for this name, it will be on the indicated page. Closed +chapters are read-only structures and their contents are never altered in +any way. + +Once enough records have been written to fill up all the available index +space, the oldest chapter is removed to make space for new chapters. Any +time a request finds a matching record in the index, that record is copied +into the open chapter. This ensures that useful block names remain available +in the index, while unreferenced block names are forgotten over time. + +In order to find records in older chapters, the index also maintains a +higher level structure called the volume index, which contains entries +mapping each block name to the chapter containing its newest record. This +mapping is updated as records for the block name are copied or updated, +ensuring that only the newest record for a given block name can be found. +An older record for a block name will no longer be found even though it has +not been deleted from its chapter. Like the chapter index, the volume index +uses only a subset of the block name as its key and can not definitively +say that a record exists for a name. It can only say which chapter would +contain the record if a record exists. The volume index is stored entirely +in memory and is saved to storage only when the vdo target is shut down. + +From the viewpoint of a request for a particular block name, it will first +look up the name in the volume index. This search will either indicate that +the name is new, or which chapter to search. If it returns a chapter, the +request looks up its name in the chapter index. This will indicate either +that the name is new, or which record page to search. Finally, if it is not +new, the request will look for its name in the indicated record page. +This process may require up to two page reads per request (one for the +chapter index page and one for the request page). However, recently +accessed pages are cached so that these page reads can be amortized across +many block name requests. + +The volume index and the chapter indexes are implemented using a +memory-efficient structure called a delta index. Instead of storing the +entire block name (the key) for each entry, the entries are sorted by name +and only the difference between adjacent keys (the delta) is stored. +Because we expect the hashes to be randomly distributed, the size of the +deltas follows an exponential distribution. Because of this distribution, +the deltas are expressed using a Huffman code to take up even less space. +The entire sorted list of keys is called a delta list. This structure +allows the index to use many fewer bytes per entry than a traditional hash +table, but it is slightly more expensive to look up entries, because a +request must read every entry in a delta list to add up the deltas in order +to find the record it needs. The delta index reduces this lookup cost by +splitting its key space into many sub-lists, each starting at a fixed key +value, so that each individual list is short. + +The default index size can hold 64 million records, corresponding to about +256GB of data. This means that the index can identify duplicate data if the +original data was written within the last 256GB of writes. This range is +called the deduplication window. If new writes duplicate data that is older +than that, the index will not be able to find it because the records of the +older data have been removed. This means that if an application writes a +200 GB file to a vdo target and then immediately writes it again, the two +copies will deduplicate perfectly. Doing the same with a 500 GB file will +result in no deduplication, because the beginning of the file will no +longer be in the index by the time the second write begins (assuming there +is no duplication within the file itself). + +If an application anticipates a data workload that will see useful +deduplication beyond the 256GB threshold, vdo can be configured to use a +larger index with a correspondingly larger deduplication window. (This +configuration can only be set when the target is created, not altered +later. It is important to consider the expected workload for a vdo target +before configuring it.) There are two ways to do this. + +One way is to increase the memory size of the index, which also increases +the amount of backing storage required. Doubling the size of the index will +double the length of the deduplication window at the expense of doubling +the storage size and the memory requirements. + +The other option is to enable sparse indexing. Sparse indexing increases +the deduplication window by a factor of 10, at the expense of also +increasing the storage size by a factor of 10. However with sparse +indexing, the memory requirements do not increase. The trade-off is +slightly more computation per request and a slight decrease in the amount +of deduplication detected. For most workloads with significant amounts of +duplicate data, sparse indexing will detect 97-99% of the deduplication +that a standard index will detect. + +The vio and data_vio Structures +------------------------------- + +A vio (short for Vdo I/O) is conceptually similar to a bio, with additional +fields and data to track vdo-specific information. A struct vio maintains a +pointer to a bio but also tracks other fields specific to the operation of +vdo. The vio is kept separate from its related bio because there are many +circumstances where vdo completes the bio but must continue to do work +related to deduplication or compression. + +Metadata reads and writes, and other writes that originate within vdo, use +a struct vio directly. Application reads and writes use a larger structure +called a data_vio to track information about their progress. A struct +data_vio contain a struct vio and also includes several other fields +related to deduplication and other vdo features. The data_vio is the +primary unit of application work in vdo. Each data_vio proceeds through a +set of steps to handle the application data, after which it is reset and +returned to a pool of data_vios for reuse. + +There is a fixed pool of 2048 data_vios. This number was chosen to bound +the amount of work that is required to recover from a crash. In addition, +benchmarks have indicated that increasing the size of the pool does not +significantly improve performance. + +The Data Store +-------------- + +The data store is implemented by three main data structures, all of which +work in concert to reduce or amortize metadata updates across as many data +writes as possible. + +*The Slab Depot* + +Most of the vdo volume belongs to the slab depot. The depot contains a +collection of slabs. The slabs can be up to 32GB, and are divided into +three sections. Most of a slab consists of a linear sequence of 4K blocks. +These blocks are used either to store data, or to hold portions of the +block map (see below). In addition to the data blocks, each slab has a set +of reference counters, using 1 byte for each data block. Finally each slab +has a journal. + +Reference updates are written to the slab journal. Slab journal blocks are +written out either when they are full, or when the recovery journal +requests they do so in order to allow the main recovery journal (see below) +to free up space. The slab journal is used both to ensure that the main +recovery journal can regularly free up space, and also to amortize the cost +of updating individual reference blocks. The reference counters are kept in +memory and are written out, a block at a time in oldest-dirtied-order, only +when there is a need to reclaim slab journal space. The write operations +are performed in the background as needed so they do not add latency to +particular I/O operations. + +Each slab is independent of every other. They are assigned to "physical +zones" in round-robin fashion. If there are P physical zones, then slab n +is assigned to zone n mod P. + +The slab depot maintains an additional small data structure, the "slab +summary," which is used to reduce the amount of work needed to come back +online after a crash. The slab summary maintains an entry for each slab +indicating whether or not the slab has ever been used, whether all of its +reference count updates have been persisted to storage, and approximately +how full it is. During recovery, each physical zone will attempt to recover +at least one slab, stopping whenever it has recovered a slab which has some +free blocks. Once each zone has some space, or has determined that none is +available, the target can resume normal operation in a degraded mode. Read +and write requests can be serviced, perhaps with degraded performance, +while the remainder of the dirty slabs are recovered. + +*The Block Map* + +The block map contains the logical to physical mapping. It can be thought +of as an array with one entry per logical address. Each entry is 5 bytes, +36 bits of which contain the physical block number which holds the data for +the given logical address. The other 4 bits are used to indicate the nature +of the mapping. Of the 16 possible states, one represents a logical address +which is unmapped (i.e. it has never been written, or has been discarded), +one represents an uncompressed block, and the other 14 states are used to +indicate that the mapped data is compressed, and which of the compression +slots in the compressed block contains the data for this logical address. + +In practice, the array of mapping entries is divided into "block map +pages," each of which fits in a single 4K block. Each block map page +consists of a header and 812 mapping entries. Each mapping page is actually +a leaf of a radix tree which consists of block map pages at each level. +There are 60 radix trees which are assigned to "logical zones" in round +robin fashion. (If there are L logical zones, tree n will belong to zone n +mod L.) At each level, the trees are interleaved, so logical addresses +0-811 belong to tree 0, logical addresses 812-1623 belong to tree 1, and so +on. The interleaving is maintained all the way up to the 60 root nodes. +Choosing 60 trees results in an evenly distributed number of trees per zone +for a large number of possible logical zone counts. The storage for the 60 +tree roots is allocated at format time. All other block map pages are +allocated out of the slabs as needed. This flexible allocation avoids the +need to pre-allocate space for the entire set of logical mappings and also +makes growing the logical size of a vdo relatively easy. + +In operation, the block map maintains two caches. It is prohibitive to keep +the entire leaf level of the trees in memory, so each logical zone +maintains its own cache of leaf pages. The size of this cache is +configurable at target start time. The second cache is allocated at start +time, and is large enough to hold all the non-leaf pages of the entire +block map. This cache is populated as pages are needed. + +*The Recovery Journal* + +The recovery journal is used to amortize updates across the block map and +slab depot. Each write request causes an entry to be made in the journal. +Entries are either "data remappings" or "block map remappings." For a data +remapping, the journal records the logical address affected and its old and +new physical mappings. For a block map remapping, the journal records the +block map page number and the physical block allocated for it. Block map +pages are never reclaimed or repurposed, so the old mapping is always 0. + +Each journal entry is an intent record summarizing the metadata updates +that are required for a data_vio. The recovery journal issues a flush +before each journal block write to ensure that the physical data for the +new block mappings in that block are stable on storage, and journal block +writes are all issued with the FUA bit set to ensure the recovery journal +entries themselves are stable. The journal entry and the data write it +represents must be stable on disk before the other metadata structures may +be updated to reflect the operation. These entries allow the vdo device to +reconstruct the logical to physical mappings after an unexpected +interruption such as a loss of power. + +*Write Path* + +All write I/O to vdo is asynchronous. Each bio will be acknowledged as soon +as vdo has done enough work to guarantee that it can complete the write +eventually. Generally, the data for acknowledged but unflushed write I/O +can be treated as though it is cached in memory. If an application +requires data to be stable on storage, it must issue a flush or write the +data with the FUA bit set like any other asynchronous I/O. Shutting down +the vdo target will also flush any remaining I/O. + +Application write bios follow the steps outlined below. + +1. A data_vio is obtained from the data_vio pool and associated with the + application bio. If there are no data_vios available, the incoming bio + will block until a data_vio is available. This provides back pressure + to the application. The data_vio pool is protected by a spin lock. + + The newly acquired data_vio is reset and the bio's data is copied into + the data_vio if it is a write and the data is not all zeroes. The data + must be copied because the application bio can be acknowledged before + the data_vio processing is complete, which means later processing steps + will no longer have access to the application bio. The application bio + may also be smaller than 4K, in which case the data_vio will have + already read the underlying block and the data is instead copied over + the relevant portion of the larger block. + +2. The data_vio places a claim (the "logical lock") on the logical address + of the bio. It is vital to prevent simultaneous modifications of the + same logical address, because deduplication involves sharing blocks. + This claim is implemented as an entry in a hashtable where the key is + the logical address and the value is a pointer to the data_vio + currently handling that address. + + If a data_vio looks in the hashtable and finds that another data_vio is + already operating on that logical address, it waits until the previous + operation finishes. It also sends a message to inform the current + lock holder that it is waiting. Most notably, a new data_vio waiting + for a logical lock will flush the previous lock holder out of the + compression packer (step 8d) rather than allowing it to continue + waiting to be packed. + + This stage requires the data_vio to get an implicit lock on the + appropriate logical zone to prevent concurrent modifications of the + hashtable. This implicit locking is handled by the zone divisions + described above. + +3. The data_vio traverses the block map tree to ensure that all the + necessary internal tree nodes have been allocated, by trying to find + the leaf page for its logical address. If any interior tree page is + missing, it is allocated at this time out of the same physical storage + pool used to store application data. + + a. If any page-node in the tree has not yet been allocated, it must be + allocated before the write can continue. This step requires the + data_vio to lock the page-node that needs to be allocated. This + lock, like the logical block lock in step 2, is a hashtable entry + that causes other data_vios to wait for the allocation process to + complete. + + The implicit logical zone lock is released while the allocation is + happening, in order to allow other operations in the same logical + zone to proceed. The details of allocation are the same as in + step 4. Once a new node has been allocated, that node is added to + the tree using a similar process to adding a new data block mapping. + The data_vio journals the intent to add the new node to the block + map tree (step 10), updates the reference count of the new block + (step 11), and reacquires the implicit logical zone lock to add the + new mapping to the parent tree node (step 12). Once the tree is + updated, the data_vio proceeds down the tree. Any other data_vios + waiting on this allocation also proceed. + + b. In the steady-state case, the block map tree nodes will already be + allocated, so the data_vio just traverses the tree until it finds + the required leaf node. The location of the mapping (the "block map + slot") is recorded in the data_vio so that later steps do not need + to traverse the tree again. The data_vio then releases the implicit + logical zone lock. + +4. If the block is a zero block, skip to step 9. Otherwise, an attempt is + made to allocate a free data block. This allocation ensures that the + data_vio can write its data somewhere even if deduplication and + compression are not possible. This stage gets an implicit lock on a + physical zone to search for free space within that zone. + + The data_vio will search each slab in a zone until it finds a free + block or decides there are none. If the first zone has no free space, + it will proceed to search the next physical zone by taking the implicit + lock for that zone and releasing the previous one until it finds a + free block or runs out of zones to search. The data_vio will acquire a + struct pbn_lock (the "physical block lock") on the free block. The + struct pbn_lock also has several fields to record the various kinds of + claims that data_vios can have on physical blocks. The pbn_lock is + added to a hashtable like the logical block locks in step 2. This + hashtable is also covered by the implicit physical zone lock. The + reference count of the free block is updated to prevent any other + data_vio from considering it free. The reference counters are a + sub-component of the slab and are thus also covered by the implicit + physical zone lock. + +5. If an allocation was obtained, the data_vio has all the resources it + needs to complete the write. The application bio can safely be + acknowledged at this point. The acknowledgment happens on a separate + thread to prevent the application callback from blocking other data_vio + operations. + + If an allocation could not be obtained, the data_vio continues to + attempt to deduplicate or compress the data, but the bio is not + acknowledged because the vdo device may be out of space. + +6. At this point vdo must determine where to store the application data. + The data_vio's data is hashed and the hash (the "record name") is + recorded in the data_vio. + +7. The data_vio reserves or joins a struct hash_lock, which manages all of + the data_vios currently writing the same data. Active hash locks are + tracked in a hashtable similar to the way logical block locks are + tracked in step 2. This hashtable is covered by the implicit lock on + the hash zone. + + If there is no existing hash lock for this data_vio's record_name, the + data_vio obtains a hash lock from the pool, adds it to the hashtable, + and sets itself as the new hash lock's "agent." The hash_lock pool is + also covered by the implicit hash zone lock. The hash lock agent will + do all the work to decide where the application data will be + written. If a hash lock for the data_vio's record_name already exists, + and the data_vio's data is the same as the agent's data, the new + data_vio will wait for the agent to complete its work and then share + its result. + + In the rare case that a hash lock exists for the data_vio's hash but + the data does not match the hash lock's agent, the data_vio skips to + step 8h and attempts to write its data directly. This can happen if two + different data blocks produce the same hash, for example. + +8. The hash lock agent attempts to deduplicate or compress its data with + the following steps. + + a. The agent initializes and sends its embedded deduplication request + (struct uds_request) to the deduplication index. This does not + require the data_vio to get any locks because the index components + manage their own locking. The data_vio waits until it either gets a + response from the index or times out. + + b. If the deduplication index returns advice, the data_vio attempts to + obtain a physical block lock on the indicated physical address, in + order to read the data and verify that it is the same as the + data_vio's data, and that it can accept more references. If the + physical address is already locked by another data_vio, the data at + that address may soon be overwritten so it is not safe to use the + address for deduplication. + + c. If the data matches and the physical block can add references, the + agent and any other data_vios waiting on it will record this + physical block as their new physical address and proceed to step 9 + to record their new mapping. If there are more data_vios in the hash + lock than there are references available, one of the remaining + data_vios becomes the new agent and continues to step 8d as if no + valid advice was returned. + + d. If no usable duplicate block was found, the agent first checks that + it has an allocated physical block (from step 3) that it can write + to. If the agent does not have an allocation, some other data_vio in + the hash lock that does have an allocation takes over as agent. If + none of the data_vios have an allocated physical block, these writes + are out of space, so they proceed to step 13 for cleanup. + + e. The agent attempts to compress its data. If the data does not + compress, the data_vio will continue to step 8h to write its data + directly. + + If the compressed size is small enough, the agent will release the + implicit hash zone lock and go to the packer (struct packer) where + it will be placed in a bin (struct packer_bin) along with other + data_vios. All compression operations require the implicit lock on + the packer zone. + + The packer can combine up to 14 compressed blocks in a single 4k + data block. Compression is only helpful if vdo can pack at least 2 + data_vios into a single data block. This means that a data_vio may + wait in the packer for an arbitrarily long time for other data_vios + to fill out the compressed block. There is a mechanism for vdo to + evict waiting data_vios when continuing to wait would cause + problems. Circumstances causing an eviction include an application + flush, device shutdown, or a subsequent data_vio trying to overwrite + the same logical block address. A data_vio may also be evicted from + the packer if it cannot be paired with any other compressed block + before more compressible blocks need to use its bin. An evicted + data_vio will proceed to step 8h to write its data directly. + + f. If the agent fills a packer bin, either because all 14 of its slots + are used or because it has no remaining space, it is written out + using the allocated physical block from one of its data_vios. Step + 8d has already ensured that an allocation is available. + + g. Each data_vio sets the compressed block as its new physical address. + The data_vio obtains an implicit lock on the physical zone and + acquires the struct pbn_lock for the compressed block, which is + modified to be a shared lock. Then it releases the implicit physical + zone lock and proceeds to step 8i. + + h. Any data_vio evicted from the packer will have an allocation from + step 3. It will write its data to that allocated physical block. + + i. After the data is written, if the data_vio is the agent of a hash + lock, it will reacquire the implicit hash zone lock and share its + physical address with as many other data_vios in the hash lock as + possible. Each data_vio will then proceed to step 9 to record its + new mapping. + + j. If the agent actually wrote new data (whether compressed or not), + the deduplication index is updated to reflect the location of the + new data. The agent then releases the implicit hash zone lock. + +9. The data_vio determines the previous mapping of the logical address. + There is a cache for block map leaf pages (the "block map cache"), + because there are usually too many block map leaf nodes to store + entirely in memory. If the desired leaf page is not in the cache, the + data_vio will reserve a slot in the cache and load the desired page + into it, possibly evicting an older cached page. The data_vio then + finds the current physical address for this logical address (the "old + physical mapping"), if any, and records it. This step requires a lock + on the block map cache structures, covered by the implicit logical zone + lock. + +10. The data_vio makes an entry in the recovery journal containing the + logical block address, the old physical mapping, and the new physical + mapping. Making this journal entry requires holding the implicit + recovery journal lock. The data_vio will wait in the journal until all + recovery blocks up to the one containing its entry have been written + and flushed to ensure the transaction is stable on storage. + +11. Once the recovery journal entry is stable, the data_vio makes two slab + journal entries: an increment entry for the new mapping, and a + decrement entry for the old mapping. These two operations each require + holding a lock on the affected physical slab, covered by its implicit + physical zone lock. For correctness during recovery, the slab journal + entries in any given slab journal must be in the same order as the + corresponding recovery journal entries. Therefore, if the two entries + are in different zones, they are made concurrently, and if they are in + the same zone, the increment is always made before the decrement in + order to avoid underflow. After each slab journal entry is made in + memory, the associated reference count is also updated in memory. + +12. Once both of the reference count updates are done, the data_vio + acquires the implicit logical zone lock and updates the + logical-to-physical mapping in the block map to point to the new + physical block. At this point the write operation is complete. + +13. If the data_vio has a hash lock, it acquires the implicit hash zone + lock and releases its hash lock to the pool. + + The data_vio then acquires the implicit physical zone lock and releases + the struct pbn_lock it holds for its allocated block. If it had an + allocation that it did not use, it also sets the reference count for + that block back to zero to free it for use by subsequent data_vios. + + The data_vio then acquires the implicit logical zone lock and releases + the logical block lock acquired in step 2. + + The application bio is then acknowledged if it has not previously been + acknowledged, and the data_vio is returned to the pool. + +*Read Path* + +An application read bio follows a much simpler set of steps. It does steps +1 and 2 in the write path to obtain a data_vio and lock its logical +address. If there is already a write data_vio in progress for that logical +address that is guaranteed to complete, the read data_vio will copy the +data from the write data_vio and return it. Otherwise, it will look up the +logical-to-physical mapping by traversing the block map tree as in step 3, +and then read and possibly decompress the indicated data at the indicated +physical block address. A read data_vio will not allocate block map tree +nodes if they are missing. If the interior block map nodes do not exist +yet, the logical block map address must still be unmapped and the read +data_vio will return all zeroes. A read data_vio handles cleanup and +acknowledgment as in step 13, although it only needs to release the logical +lock and return itself to the pool. + +*Small Writes* + +All storage within vdo is managed as 4KB blocks, but it can accept writes +as small as 512 bytes. Processing a write that is smaller than 4K requires +a read-modify-write operation that reads the relevant 4K block, copies the +new data over the approriate sectors of the block, and then launches a +write operation for the modified data block. The read and write stages of +this operation are nearly identical to the normal read and write +operations, and a single data_vio is used throughout this operation. + +*Recovery* + +When a vdo is restarted after a crash, it will attempt to recover from the +recovery journal. During the pre-resume phase of the next start, the +recovery journal is read. The increment portion of valid entries are played +into the block map. Next, valid entries are played, in order as required, +into the slab journals. Finally, each physical zone attempts to replay at +least one slab journal to reconstruct the reference counts of one slab. +Once each zone has some free space (or has determined that it has none), +the vdo comes back online, while the remainder of the slab journals are +used to reconstruct the rest of the reference counts in the background. + +*Read-only Rebuild* + +If a vdo encounters an unrecoverable error, it will enter read-only mode. +This mode indicates that some previously acknowledged data may have been +lost. The vdo may be instructed to rebuild as best it can in order to +return to a writable state. However, this is never done automatically due +to the possibility that data has been lost. During a read-only rebuild, the +block map is recovered from the recovery journal as before. However, the +reference counts are not rebuilt from the slab journals. Instead, the +reference counts are zeroed, the entire block map is traversed, and the +reference counts are updated from the block mappings. While this may lose +some data, it ensures that the block map and reference counts are +consistent with each other. This allows vdo to resume normal operation and +accept further writes. |