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reftable

Overview

Problem statement

Some repositories contain a lot of references (e.g. android at 866k, rails at 31k). The existing packed-refs format takes up a lot of space (e.g. 62M), and does not scale with additional references. Lookup of a single reference requires linearly scanning the file.

Atomic pushes modifying multiple references require copying the entire packed-refs file, which can be a considerable amount of data moved (e.g. 62M in, 62M out) for even small transactions (2 refs modified).

Repositories with many loose references occupy a large number of disk blocks from the local file system, as each reference is its own file storing 41 bytes (and another file for the corresponding reflog). This negatively affects the number of inodes available when a large number of repositories are stored on the same filesystem. Readers can be penalized due to the larger number of syscalls required to traverse and read the $GIT_DIR/refs directory.

Objectives

  • Near constant time lookup for any single reference, even when the repository is cold and not in process or kernel cache.

  • Near constant time verification if an object name is referred to by at least one reference (for allow-tip-sha1-in-want).

  • Efficient enumeration of an entire namespace, such as refs/tags/.

  • Support atomic push with O(size_of_update) operations.

  • Combine reflog storage with ref storage for small transactions.

  • Separate reflog storage for base refs and historical logs.

Description

A reftable file is a portable binary file format customized for reference storage. References are sorted, enabling linear scans, binary search lookup, and range scans.

Storage in the file is organized into variable sized blocks. Prefix compression is used within a single block to reduce disk space. Block size and alignment is tunable by the writer.

Performance

Space used, packed-refs vs. reftable:

repository packed-refs reftable % original avg ref avg obj

android

62.2 M

36.1 M

58.0%

33 bytes

5 bytes

rails

1.8 M

1.1 M

57.7%

29 bytes

4 bytes

git

78.7 K

48.1 K

61.0%

50 bytes

4 bytes

git (heads)

332 b

269 b

81.0%

33 bytes

0 bytes

Scan (read 866k refs), by reference name lookup (single ref from 866k refs), and by SHA-1 lookup (refs with that SHA-1, from 866k refs):

format cache scan by name by SHA-1

packed-refs

cold

402 ms

409,660.1 usec

412,535.8 usec

packed-refs

hot

6,844.6 usec

20,110.1 usec

reftable

cold

112 ms

33.9 usec

323.2 usec

reftable

hot

20.2 usec

320.8 usec

Space used for 149,932 log entries for 43,061 refs, reflog vs. reftable:

format size avg entry

$GIT_DIR/logs

173 M

1209 bytes

reftable

5 M

37 bytes

Details

Peeling

References stored in a reftable are peeled, a record for an annotated (or signed) tag records both the tag object, and the object it refers to. This is analogous to storage in the packed-refs format.

Reference name encoding

Reference names are an uninterpreted sequence of bytes that must pass git-check-ref-format[1] as a valid reference name.

Key unicity

Each entry must have a unique key; repeated keys are disallowed.

Network byte order

All multi-byte, fixed width fields are in network byte order.

Varint encoding

Varint encoding is identical to the ofs-delta encoding method used within pack files.

Decoder works such as:

val = buf[ptr] & 0x7f
while (buf[ptr] & 0x80) {
  ptr++
  val = ((val + 1) << 7) | (buf[ptr] & 0x7f)
}

Ordering

Blocks are lexicographically ordered by their first reference.

Directory/file conflicts

The reftable format accepts both refs/heads/foo and refs/heads/foo/bar as distinct references.

This property is useful for retaining log records in reftable, but may confuse versions of Git using $GIT_DIR/refs directory tree to maintain references. Users of reftable may choose to continue to reject foo and foo/bar type conflicts to prevent problems for peers.

File format

Structure

A reftable file has the following high-level structure:

first_block {
  header
  first_ref_block
}
ref_block*
ref_index*
obj_block*
obj_index*
log_block*
log_index*
footer

A log-only file omits the ref_block, ref_index, obj_block and obj_index sections, containing only the file header and log block:

first_block {
  header
}
log_block*
log_index*
footer

in a log-only file the first log block immediately follows the file header, without padding to block alignment.

Block size

The file’s block size is arbitrarily determined by the writer, and does not have to be a power of 2. The block size must be larger than the longest reference name or log entry used in the repository, as references cannot span blocks.

Powers of two that are friendly to the virtual memory system or filesystem (such as 4k or 8k) are recommended. Larger sizes (64k) can yield better compression, with a possible increased cost incurred by readers during access.

The largest block size is 16777215 bytes (15.99 MiB).

Block alignment

Writers may choose to align blocks at multiples of the block size by including padding filled with NUL bytes at the end of a block to round out to the chosen alignment. When alignment is used, writers must specify the alignment with the file header’s block_size field.

Block alignment is not required by the file format. Unaligned files must set block_size = 0 in the file header, and omit padding. Unaligned files with more than one ref block must include the ref index to support fast lookup. Readers must be able to read both aligned and non-aligned files.

Very small files (e.g. a single ref block) may omit padding and the ref index to reduce total file size.

Header (version 1)

A 24-byte header appears at the beginning of the file:

'REFT'
uint8( version_number = 1 )
uint24( block_size )
uint64( min_update_index )
uint64( max_update_index )

Aligned files must specify block_size to configure readers with the expected block alignment. Unaligned files must set block_size = 0.

The min_update_index and max_update_index describe bounds for the update_index field of all log records in this file. When reftables are used in a stack for transactions, these fields can order the files such that the prior file’s max_update_index + 1 is the next file’s min_update_index.

Header (version 2)

A 28-byte header appears at the beginning of the file:

'REFT'
uint8( version_number = 2 )
uint24( block_size )
uint64( min_update_index )
uint64( max_update_index )
uint32( hash_id )

The header is identical to version_number=1, with the 4-byte hash ID ("sha1" for SHA1 and "s256" for SHA-256) append to the header.

For maximum backward compatibility, it is recommended to use version 1 when writing SHA1 reftables.

First ref block

The first ref block shares the same block as the file header, and is 24 bytes smaller than all other blocks in the file. The first block immediately begins after the file header, at position 24.

If the first block is a log block (a log-only file), its block header begins immediately at position 24.

Ref block format

A ref block is written as:

'r'
uint24( block_len )
ref_record+
uint24( restart_offset )+
uint16( restart_count )

padding?

Blocks begin with block_type = 'r' and a 3-byte block_len which encodes the number of bytes in the block up to, but not including the optional padding. This is always less than or equal to the file’s block size. In the first ref block, block_len includes 24 bytes for the file header.

The 2-byte restart_count stores the number of entries in the restart_offset list, which must not be empty. Readers can use restart_count to binary search between restarts before starting a linear scan.

Exactly restart_count 3-byte restart_offset values precedes the restart_count. Offsets are relative to the start of the block and refer to the first byte of any ref_record whose name has not been prefix compressed. Entries in the restart_offset list must be sorted, ascending. Readers can start linear scans from any of these records.

A variable number of ref_record fill the middle of the block, describing reference names and values. The format is described below.

As the first ref block shares the first file block with the file header, all restart_offset in the first block are relative to the start of the file (position 0), and include the file header. This forces the first restart_offset to be 28.

ref record

A ref_record describes a single reference, storing both the name and its value(s). Records are formatted as:

varint( prefix_length )
varint( (suffix_length << 3) | value_type )
suffix
varint( update_index_delta )
value?

The prefix_length field specifies how many leading bytes of the prior reference record’s name should be copied to obtain this reference’s name. This must be 0 for the first reference in any block, and also must be 0 for any ref_record whose offset is listed in the restart_offset table at the end of the block.

Recovering a reference name from any ref_record is a simple concat:

this_name = prior_name[0..prefix_length] + suffix

The suffix_length value provides the number of bytes available in suffix to copy from suffix to complete the reference name.

The update_index that last modified the reference can be obtained by adding update_index_delta to the min_update_index from the file header: min_update_index + update_index_delta.

The value follows. Its format is determined by value_type, one of the following:

  • 0x0: deletion; no value data (see transactions, below)

  • 0x1: one object name; value of the ref

  • 0x2: two object names; value of the ref, peeled target

  • 0x3: symbolic reference: varint( target_len ) target

Symbolic references use 0x3, followed by the complete name of the reference target. No compression is applied to the target name.

Types 0x4..0x7 are reserved for future use.

Ref index

The ref index stores the name of the last reference from every ref block in the file, enabling reduced disk seeks for lookups. Any reference can be found by searching the index, identifying the containing block, and searching within that block.

The index may be organized into a multi-level index, where the 1st level index block points to additional ref index blocks (2nd level), which may in turn point to either additional index blocks (e.g. 3rd level) or ref blocks (leaf level). Disk reads required to access a ref go up with higher index levels. Multi-level indexes may be required to ensure no single index block exceeds the file format’s max block size of 16777215 bytes (15.99 MiB). To achieve constant O(1) disk seeks for lookups the index must be a single level, which is permitted to exceed the file’s configured block size, but not the format’s max block size of 15.99 MiB.

If present, the ref index block(s) appears after the last ref block.

If there are at least 4 ref blocks, a ref index block should be written to improve lookup times. Cold reads using the index require 2 disk reads (read index, read block), and binary searching < 4 blocks also requires ⇐ 2 reads. Omitting the index block from smaller files saves space.

If the file is unaligned and contains more than one ref block, the ref index must be written.

Index block format:

'i'
uint24( block_len )
index_record+
uint24( restart_offset )+
uint16( restart_count )

padding?

The index blocks begin with block_type = 'i' and a 3-byte block_len which encodes the number of bytes in the block, up to but not including the optional padding.

The restart_offset and restart_count fields are identical in format, meaning and usage as in ref blocks.

To reduce the number of reads required for random access in very large files the index block may be larger than other blocks. However, readers must hold the entire index in memory to benefit from this, so it’s a time-space tradeoff in both file size and reader memory.

Increasing the file’s block size decreases the index size. Alternatively a multi-level index may be used, keeping index blocks within the file’s block size, but increasing the number of blocks that need to be accessed.

index record

An index record describes the last entry in another block. Index records are written as:

varint( prefix_length )
varint( (suffix_length << 3) | 0 )
suffix
varint( block_position )

Index records use prefix compression exactly like ref_record.

Index records store block_position after the suffix, specifying the absolute position in bytes (from the start of the file) of the block that ends with this reference. Readers can seek to block_position to begin reading the block header.

Readers must examine the block header at block_position to determine if the next block is another level index block, or the leaf-level ref block.

Reading the index

Readers loading the ref index must first read the footer (below) to obtain ref_index_position. If not present, the position will be 0. The ref_index_position is for the 1st level root of the ref index.

Obj block format

Object blocks are optional. Writers may choose to omit object blocks, especially if readers will not use the object name to ref mapping.

Object blocks use unique, abbreviated 2-32 object name keys, mapping to ref blocks containing references pointing to that object directly, or as the peeled value of an annotated tag. Like ref blocks, object blocks use the file’s standard block size. The abbreviation length is available in the footer as obj_id_len.

To save space in small files, object blocks may be omitted if the ref index is not present, as brute force search will only need to read a few ref blocks. When missing, readers should brute force a linear search of all references to lookup by object name.

An object block is written as:

'o'
uint24( block_len )
obj_record+
uint24( restart_offset )+
uint16( restart_count )

padding?

Fields are identical to ref block. Binary search using the restart table works the same as in reference blocks.

Because object names are abbreviated by writers to the shortest unique abbreviation within the reftable, obj key lengths have a variable length. Their length must be at least 2 bytes. Readers must compare only for common prefix match within an obj block or obj index.

obj record

An obj_record describes a single object abbreviation, and the blocks containing references using that unique abbreviation:

varint( prefix_length )
varint( (suffix_length << 3) | cnt_3 )
suffix
varint( cnt_large )?
varint( position_delta )*

Like in reference blocks, abbreviations are prefix compressed within an obj block. On large reftables with many unique objects, higher block sizes (64k), and higher restart interval (128), a prefix_length of 2 or 3 and suffix_length of 3 may be common in obj records (unique abbreviation of 5-6 raw bytes, 10-12 hex digits).

Each record contains position_count number of positions for matching ref blocks. For 1-7 positions the count is stored in cnt_3. When cnt_3 = 0 the actual count follows in a varint, cnt_large.

The use of cnt_3 bets most objects are pointed to by only a single reference, some may be pointed to by a couple of references, and very few (if any) are pointed to by more than 7 references.

A special case exists when cnt_3 = 0 and cnt_large = 0: there are no position_delta, but at least one reference starts with this abbreviation. A reader that needs exact reference names must scan all references to find which specific references have the desired object. Writers should use this format when the position_delta list would have overflowed the file’s block size due to a high number of references pointing to the same object.

The first position_delta is the position from the start of the file. Additional position_delta entries are sorted ascending and relative to the prior entry, e.g. a reader would perform:

pos = position_delta[0]
prior = pos
for (j = 1; j < position_count; j++) {
  pos = prior + position_delta[j]
  prior = pos
}

With a position in hand, a reader must linearly scan the ref block, starting from the first ref_record, testing each reference’s object names (for value_type = 0x1 or 0x2) for full equality. Faster searching by object name within a single ref block is not supported by the reftable format. Smaller block sizes reduce the number of candidates this step must consider.

Obj index

The obj index stores the abbreviation from the last entry for every obj block in the file, enabling reduced disk seeks for all lookups. It is formatted exactly the same as the ref index, but refers to obj blocks.

The obj index should be present if obj blocks are present, as obj blocks should only be written in larger files.

Readers loading the obj index must first read the footer (below) to obtain obj_index_position. If not present, the position will be 0.

Log block format

Unlike ref and obj blocks, log blocks are always unaligned.

Log blocks are variable in size, and do not match the block_size specified in the file header or footer. Writers should choose an appropriate buffer size to prepare a log block for deflation, such as 2 * block_size.

A log block is written as:

'g'
uint24( block_len )
zlib_deflate {
  log_record+
  uint24( restart_offset )+
  uint16( restart_count )
}

Log blocks look similar to ref blocks, except block_type = 'g'.

The 4-byte block header is followed by the deflated block contents using zlib deflate. The block_len in the header is the inflated size (including 4-byte block header), and should be used by readers to preallocate the inflation output buffer. A log block’s block_len may exceed the file’s block size.

Offsets within the log block (e.g. restart_offset) still include the 4-byte header. Readers may prefer prefixing the inflation output buffer with the 4-byte header.

Within the deflate container, a variable number of log_record describe reference changes. The log record format is described below. See ref block format (above) for a description of restart_offset and restart_count.

Because log blocks have no alignment or padding between blocks, readers must keep track of the bytes consumed by the inflater to know where the next log block begins.

log record

Log record keys are structured as:

ref_name '\0' reverse_int64( update_index )

where update_index is the unique transaction identifier. The update_index field must be unique within the scope of a ref_name. See the update transactions section below for further details.

The reverse_int64 function inverses the value so lexicographical ordering the network byte order encoding sorts the more recent records with higher update_index values first:

reverse_int64(int64 t) {
  return 0xffffffffffffffff - t;
}

Log records have a similar starting structure to ref and index records, utilizing the same prefix compression scheme applied to the log record key described above.

    varint( prefix_length )
    varint( (suffix_length << 3) | log_type )
    suffix
    log_data {
      old_id
      new_id
      varint( name_length    )  name
      varint( email_length   )  email
      varint( time_seconds )
      sint16( tz_offset )
      varint( message_length )  message
    }?

Log record entries use log_type to indicate what follows:

  • 0x0: deletion; no log data.

  • 0x1: standard git reflog data using log_data above.

The log_type = 0x0 is mostly useful for git stash drop, removing an entry from the reflog of refs/stash in a transaction file (below), without needing to rewrite larger files. Readers reading a stack of reflogs must treat this as a deletion.

For log_type = 0x1, the log_data section follows git-update-ref[1] logging and includes:

  • two object names (old id, new id)

  • varint string of committer’s name

  • varint string of committer’s email

  • varint time in seconds since epoch (Jan 1, 1970)

  • 2-byte timezone offset in minutes (signed)

  • varint string of message

tz_offset is the absolute number of minutes from GMT the committer was at the time of the update. For example GMT-0800 is encoded in reftable as sint16(-480) and GMT+0230 is sint16(150).

The committer email does not contain < or >, it’s the value normally found between the <> in a git commit object header.

The message_length may be 0, in which case there was no message supplied for the update.

Contrary to traditional reflog (which is a file), renames are encoded as a combination of ref deletion and ref creation. A deletion is a log record with a zero new_id, and a creation is a log record with a zero old_id.

Reading the log

Readers accessing the log must first read the footer (below) to determine the log_position. The first block of the log begins at log_position bytes since the start of the file. The log_position is not block aligned.

Importing logs

When importing from $GIT_DIR/logs writers should globally order all log records roughly by timestamp while preserving file order, and assign unique, increasing update_index values for each log line. Newer log records get higher update_index values.

Although an import may write only a single reftable file, the reftable file must span many unique update_index, as each log line requires its own update_index to preserve semantics.

Log index

The log index stores the log key (refname \0 reverse_int64(update_index)) for the last log record of every log block in the file, supporting bounded-time lookup.

A log index block must be written if 2 or more log blocks are written to the file. If present, the log index appears after the last log block. There is no padding used to align the log index to block alignment.

Log index format is identical to ref index, except the keys are 9 bytes longer to include '\0' and the 8-byte reverse_int64(update_index). Records use block_position to refer to the start of a log block.

Reading the index

Readers loading the log index must first read the footer (below) to obtain log_index_position. If not present, the position will be 0.

After the last block of the file, a file footer is written. It begins like the file header, but is extended with additional data.

    HEADER

    uint64( ref_index_position )
    uint64( (obj_position << 5) | obj_id_len )
    uint64( obj_index_position )

    uint64( log_position )
    uint64( log_index_position )

    uint32( CRC-32 of above )

If a section is missing (e.g. ref index) the corresponding position field (e.g. ref_index_position) will be 0.

  • obj_position: byte position for the first obj block.

  • obj_id_len: number of bytes used to abbreviate object names in obj blocks.

  • log_position: byte position for the first log block.

  • ref_index_position: byte position for the start of the ref index.

  • obj_index_position: byte position for the start of the obj index.

  • log_index_position: byte position for the start of the log index.

The size of the footer is 68 bytes for version 1, and 72 bytes for version 2.

Readers must first read the file start to determine the version number. Then they seek to file_length - FOOTER_LENGTH to access the footer. A trusted external source (such as stat(2)) is necessary to obtain file_length. When reading the footer, readers must verify:

  • 4-byte magic is correct

  • 1-byte version number is recognized

  • 4-byte CRC-32 matches the other 64 bytes (including magic, and version)

Once verified, the other fields of the footer can be accessed.

Empty tables

A reftable may be empty. In this case, the file starts with a header and is immediately followed by a footer.

Binary search within a block is supported by the restart_offset fields at the end of the block. Readers can binary search through the restart table to locate between which two restart points the sought reference or key should appear.

Each record identified by a restart_offset stores the complete key in the suffix field of the record, making the compare operation during binary search straightforward.

Once a restart point lexicographically before the sought reference has been identified, readers can linearly scan through the following record entries to locate the sought record, terminating if the current record sorts after (and therefore the sought key is not present).

Restart point selection

Writers determine the restart points at file creation. The process is arbitrary, but every 16 or 64 records is recommended. Every 16 may be more suitable for smaller block sizes (4k or 8k), every 64 for larger block sizes (64k).

More frequent restart points reduces prefix compression and increases space consumed by the restart table, both of which increase file size.

Less frequent restart points makes prefix compression more effective, decreasing overall file size, with increased penalties for readers walking through more records after the binary search step.

A maximum of 65535 restart points per block is supported.

Considerations

Lightweight refs dominate

The reftable format assumes the vast majority of references are single object names valued with common prefixes, such as Gerrit Code Review’s refs/changes/ namespace, GitHub’s refs/pulls/ namespace, or many lightweight tags in the refs/tags/ namespace.

Annotated tags storing the peeled object cost an additional object name per reference.

Low overhead

A reftable with very few references (e.g. git.git with 5 heads) is 269 bytes for reftable, vs. 332 bytes for packed-refs. This supports reftable scaling down for transaction logs (below).

Block size

For a Gerrit Code Review type repository with many change refs, larger block sizes (64 KiB) and less frequent restart points (every 64) yield better compression due to more references within the block compressing against the prior reference.

Larger block sizes reduce the index size, as the reftable will require fewer blocks to store the same number of references.

Minimal disk seeks

Assuming the index block has been loaded into memory, binary searching for any single reference requires exactly 1 disk seek to load the containing block.

Scans and lookups dominate

Scanning all references and lookup by name (or namespace such as refs/heads/) are the most common activities performed on repositories. Object names are stored directly with references to optimize this use case.

Logs are infrequently read

Logs are infrequently accessed, but can be large. Deflating log blocks saves disk space, with some increased penalty at read time.

Logs are stored in an isolated section from refs, reducing the burden on reference readers that want to ignore logs. Further, historical logs can be isolated into log-only files.

Logs are read backwards

Logs are frequently accessed backwards (most recent N records for master to answer master@{4}), so log records are grouped by reference, and sorted descending by update index.

Repository format

Version 1

A repository must set its $GIT_DIR/config to configure reftable:

[core]
    repositoryformatversion = 1
[extensions]
    refStorage = reftable

Layout

A collection of reftable files are stored in the $GIT_DIR/reftable/ directory. Their names should have a random element, such that each filename is globally unique; this helps avoid spurious failures on Windows, where open files cannot be removed or overwritten. It suggested to use ${min_update_index}-${max_update_index}-${random}.ref as a naming convention.

Log-only files use the .log extension, while ref-only and mixed ref and log files use .ref. extension.

The stack ordering file is $GIT_DIR/reftable/tables.list and lists the current files, one per line, in order, from oldest (base) to newest (most recent):

$ cat .git/reftable/tables.list
00000001-00000001-RANDOM1.log
00000002-00000002-RANDOM2.ref
00000003-00000003-RANDOM3.ref

Readers must read $GIT_DIR/reftable/tables.list to determine which files are relevant right now, and search through the stack in reverse order (last reftable is examined first).

Reftable files not listed in tables.list may be new (and about to be added to the stack by the active writer), or ancient and ready to be pruned.

Backward compatibility

Older clients should continue to recognize the directory as a git repository so they don’t look for an enclosing repository in parent directories. To this end, a reftable-enabled repository must contain the following dummy files

  • .git/HEAD, a regular file containing ref: refs/heads/.invalid.

  • .git/refs/, a directory

  • .git/refs/heads, a regular file

Readers

Readers can obtain a consistent snapshot of the reference space by following:

  1. Open and read the tables.list file.

  2. Open each of the reftable files that it mentions.

  3. If any of the files is missing, goto 1.

  4. Read from the now-open files as long as necessary.

Update transactions

Although reftables are immutable, mutations are supported by writing a new reftable and atomically appending it to the stack:

  1. Acquire tables.list.lock.

  2. Read tables.list to determine current reftables.

  3. Select update_index to be most recent file’s max_update_index + 1.

  4. Prepare temp reftable tmp_XXXXXX, including log entries.

  5. Rename tmp_XXXXXX to ${update_index}-${update_index}-${random}.ref.

  6. Copy tables.list to tables.list.lock, appending file from (5).

  7. Rename tables.list.lock to tables.list.

During step 4 the new file’s min_update_index and max_update_index are both set to the update_index selected by step 3. All log records for the transaction use the same update_index in their keys. This enables later correlation of which references were updated by the same transaction.

Because a single tables.list.lock file is used to manage locking, the repository is single-threaded for writers. Writers may have to busy-spin (with backoff) around creating tables.list.lock, for up to an acceptable wait period, aborting if the repository is too busy to mutate. Application servers wrapped around repositories (e.g. Gerrit Code Review) can layer their own lock/wait queue to improve fairness to writers.

Reference deletions

Deletion of any reference can be explicitly stored by setting the type to 0x0 and omitting the value field of the ref_record. This serves as a tombstone, overriding any assertions about the existence of the reference from earlier files in the stack.

Compaction

A partial stack of reftables can be compacted by merging references using a straightforward merge join across reftables, selecting the most recent value for output, and omitting deleted references that do not appear in remaining, lower reftables.

A compacted reftable should set its min_update_index to the smallest of the input files' min_update_index, and its max_update_index likewise to the largest input max_update_index.

For sake of illustration, assume the stack currently consists of reftable files (from oldest to newest): A, B, C, and D. The compactor is going to compact B and C, leaving A and D alone.

  1. Obtain lock tables.list.lock and read the tables.list file.

  2. Obtain locks B.lock and C.lock. Ownership of these locks prevents other processes from trying to compact these files.

  3. Release tables.list.lock.

  4. Compact B and C into a temp file ${min_update_index}-${max_update_index}_XXXXXX.

  5. Reacquire lock tables.list.lock.

  6. Verify that B and C are still in the stack, in that order. This should always be the case, assuming that other processes are adhering to the locking protocol.

  7. Rename ${min_update_index}-${max_update_index}_XXXXXX to ${min_update_index}-${max_update_index}-${random}.ref.

  8. Write the new stack to tables.list.lock, replacing B and C with the file from (4).

  9. Rename tables.list.lock to tables.list.

  10. Delete B and C, perhaps after a short sleep to avoid forcing readers to backtrack.

This strategy permits compactions to proceed independently of updates.

Each reftable (compacted or not) is uniquely identified by its name, so open reftables can be cached by their name.

Windows

On windows, and other systems that do not allow deleting or renaming to open files, compaction may succeed, but other readers may prevent obsolete tables from being deleted.

On these platforms, the following strategy can be followed: on closing a reftable stack, reload tables.list, and delete any tables no longer mentioned in tables.list.

Irregular program exit may still leave about unused files. In this case, a cleanup operation should proceed as follows:

  • take a lock tables.list.lock to prevent concurrent modifications

  • refresh the reftable stack, by reading tables.list

  • for each *.ref file, remove it if

    • it is not mentioned in tables.list, and

    • its max update_index is not beyond the max update_index of the stack

Alternatives considered

bzip packed-refs

bzip2 can significantly shrink a large packed-refs file (e.g. 62 MiB compresses to 23 MiB, 37%). However the bzip format does not support random access to a single reference. Readers must inflate and discard while performing a linear scan.

Breaking packed-refs into chunks (individually compressing each chunk) would reduce the amount of data a reader must inflate, but still leaves the problem of indexing chunks to support readers efficiently locating the correct chunk.

Given the compression achieved by reftable’s encoding, it does not seem necessary to add the complexity of bzip/gzip/zlib.

Michael Haggerty’s alternate format

Michael Haggerty proposed an alternate format to reftable on the Git mailing list. This format uses smaller chunks, without the restart table, and avoids block alignment with padding. Reflog entries immediately follow each ref, and are thus interleaved between refs.

Performance testing indicates reftable is faster for lookups (51% faster, 11.2 usec vs. 5.4 usec), although reftable produces a slightly larger file (+ ~3.2%, 28.3M vs 29.2M):

format size seek cold seek hot

mh-alt

28.3 M

23.4 usec

11.2 usec

reftable

29.2 M

19.9 usec

5.4 usec

JGit Ketch RefTree

JGit Ketch proposed RefTree, an encoding of references inside Git tree objects stored as part of the repository’s object database.

The RefTree format adds additional load on the object database storage layer (more loose objects, more objects in packs), and relies heavily on the packer’s delta compression to save space. Namespaces which are flat (e.g. thousands of tags in refs/tags) initially create very large loose objects, and so RefTree does not address the problem of copying many references to modify a handful.

Flat namespaces are not efficiently searchable in RefTree, as tree objects in canonical formatting cannot be binary searched. This fails the need to handle a large number of references in a single namespace, such as GitHub’s refs/pulls, or a project with many tags.

LMDB

David Turner proposed using LMDB, as LMDB is lightweight (64k of runtime code) and GPL-compatible license.

A downside of LMDB is its reliance on a single C implementation. This makes embedding inside JGit (a popular reimplementation of Git) difficult, and hoisting onto virtual storage (for JGit DFS) virtually impossible.

A common format that can be supported by all major Git implementations (git-core, JGit, libgit2) is strongly preferred.

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