Project: CockroachDB (App Definition and Development / Database)
File: 20170522_external_storage.md

Feature Name: External Storage for DistSQL processors Status: completed Start Date: 2017-05-11 Authors: Arjun Narayan and Alfonso Subiotto Marques RFC PR: Cockroach Issue: Add support for DistSQL processors to use external storage in addition to memory to store intermediate data when processing large DistSQL queries. Currently, DistSQL processors have a hard limit on the amount of memory they can use before they get OOM-killed. This limits the size of certain queries, particularly when efficient secondary indices do not exist, so some amount of computation must be done holding a large amount of intermediate state at a processor. For instance, a JOIN on two very large tables, where there is no ordering on either table's JOIN predicate columns, must either Use a HashJoiner, which stores one table in a single large hashtable, and stream the other table, looking up for matches on the hashtable, emitting rows when appropriate. Sort both tables by the JOIN predicate columns, and then stream them to a MergeJoiner. Either of these two solutions runs out of memory at some query size, and requires external storage to process larger queries. This problem was first encountered in running TPC-H queries on moderately sized scale factors (5 and above). It was exacerbated by the fact that DistSQL currently does not plan MergeJoins and resorts to HashJoins in all cases. TPC-H itself should not be bottlenecked on the lack of external storage, as relevant secondary indices exist for all queries that we are aware of, such that optimal planning keeps tuples flowing smoothly through the dataflow graph without requiring any processor to accumulate a disproportional number of tuples. However, users might still want to run the occasional analytics query without indexes, and we should support that use case. At a minimum, external storage should provide enough swap space that it stops all DistSQL processors from being OOM-killed on queries on the TPC-H scalefactor 300 dataset. Whatever storage format is chosen, this should be compatible with all DistSQL processors. None of this is new. Existing databases use both paths for query execution, and typically use a cost-based query planner to decide which of the two paths to use. So for us, it's more a case of prioritization and beginning with an implementation that is extensible. External sorting is typically the first implemented in OLTP databases, since it comes in handy for building new indices, but we have a different pathway for executing those. On-disk storage for query execution only comes into play for OLAP-style queries, For example, in the case of running a JOIN: They have implementations for the GRACE hash join, as well as the external sorter + merge join. They have extensive microbenchmarks for the underlying costs of each operation (e.g. total time cost of writing a tuple into a serialized stream, or the expected time cost of doing a random insert into a sorted on-disk BTree). They have extensive table statistics. At query planning time, with the table statistics, the database can compute the cost of using either plan

estimate the number of passes the GRACE hash join would take and multiply it out with the average cost of writing each tuple, etc.) and choose the better one. We don't have (2), and thus we don't have a cost-based query planner capable of doing (3). Right now, our implementations only do purely logical query planning (i.e. do only those wins that are asymptotic wins), so the RocksDB path is along those lines as well. Examining all DistSQL processors, we have the following data structures used: | Processor name | On Disk Data structure| |--|-:| | NoopCore | n/a | | TableReader | n/a | | JoinReader | n/a | | Sorter | Sorted Map| | Aggregator | Map | | Distinct | Map | | MergeJoiner | n/a | | HashJoiner | Map | | Values | n/a | | AlgebraicSetOp | Map | Essentially, we have two data types to store: a sorted map and a hash map. The processors that use a hash map (distinct, hashjoin, and algebraic set op) do internally benefit from keys being sorted, but when they are sorted, the emit rows eagerly, and don't need to build expensive in-memory state! They thus only accumulate state when keys are not sorted. Thus, sorting is in direct opposition to accumulating state large enough for external storage. The first and simplest solution, is to create an on-disk method exclusively for the Sorter processor. All other processors use less memory when the input data is appropriately sorted, so the planner would leverage the on-disk capable Sorter to sort data upstream. The simplest on-disk sorter is to use an external Unix `sort` process, or some other on-disk-capable sorter, feeding it rows, and reading back the sorted stream of tuples. More seriously, we could use an external sort implementation that's more battle tested, or even roll our own native Go implementation (e.g. see Postgresql's ). The biggest downside of only supporting external sorting is that we might have to do multiple re-sorts for a complex query plan (for instance, if a query plan has multiple sequential joins on multiple different join conditions, that would require a re-join after the first join and before the second), since each processor's ordering requirements must be treated as a stricter precondition than it currently is (at least until we have sophisticated table statistics and join planning that can take that into account). This is a quick-and-dirty fix. For processors that require an unsorted map, we could use temporary files to page out partial output streams, a la GRACE hash joins, merging them back later. Walking through two examples for how flat files would be used by processors: Consider the following algorithm for processing a HashJoin: Read n rows from the right input stream, where n is the upper limit on the number of rows we can hold in-memory. Read all rows from the left input stream, and where we see a match, serialize the row and emit it to a flat

Read the next n rows from the right input stream, and read all rows from the start from the left input stream, and emit it to a second flat file (note that DistSQL does not currently allow for rereading streams, so that capability would have to be added as well to make this work). Repeat until we exhaust the right input stream. Read all the flat files back, merging them in (if we need to preserve some ordering), and emitting tuples. Read n rows from the input stream, inserting it into an in-memory heap. Flush the heap to disk. Repeat until the input stream is exhausted. Merge all the files. The first downside is in the runtime: GRACE hash joins are not particularly pretty in worst case algorithmic analysis, and this is evident in that the hash join reads complete streams multiple times. The second downside that we need to build a new processor that can provide this buffering service. However, this is mitigated by the fact that a BufferedStream processor would reuse the same code, since all it does it write a stream to disk, and reread it as many times as necessary. Finally, this requires a bunch of processor rewriting work: each processor would have to be rewritten to deal with multiple passes, and this also makes the physical planner's job more critical: runtimes can get pathological if we have to resort to paging out streams, so getting ordering requirements right becomes even more important. While writing serialized streams to disk has the advantage of having no write amplification, it has very high read amplification since streams reread multiple times. Unfortunately, GRACE hash joins can become very expensive for very large datasets, where the number of passes over the data exceeds `log(n)`. Thus, this is not the most scalable implementation. In combination with a specialized sorter node, which does recover the `n log(n)` complexity, however, and some syntax for explicitly hinting when to sort data, we can recover reasonable performance at scale. For processors that require (or could benefit from) a sorted map, RocksDB offers the obvious straightforward on-disk representation. A way to elegantly augment existing processors is for them to use a new map library, which stores a map in-memory, but flushes it to a temporary RocksDB keyspace when it reaches some threshold. Instead of then using a second in-memory map and having to merge streams on the return path, the processors continue using the on-disk map, inserting directly into RocksDB, and performing a single scan at the end. This version also gives us random access reads and writes, which makes using external storage for aggregations where there are a large number of buckets simple: we simply use the KV interface to store aggregations by bucket. As a bonus, when we finally read off the aggregations, they will be in sorted order. Do note that if configured correctly, RocksDB recovers the algorithm in Option 2 Example 2: if we insert sequential sorted `WriteBatches`, and turn compaction off. This leaves `cgo` overhead as the only remaining downside of this path. We propose to use RocksDB for external storage, using a dedicated additional RocksDB instance solely for DistSQL

This is the most general implementation (in terms of scaling to very complex queries), is completely opaque to the user, requires no special syntax at large scales, and requires lower implementation work in modifying processors one-by-one (their internal maps would be changed to use an opaque KV library that switches from in-memory maps to RocksDB at a configurable threshold). The internal knowledge base for tuning and working with RocksDB is available at CockroachDB, so this reduces future implementation risk as well. As a sanity check to ensure that this isn't wildly off-base, we will first benchmark using RocksDB with minimal compaction versus a standard external sort library, and make sure it isn't egregious. We need to isolate the following costs: External sort vs RocksDB: We will sort a billion rows using a standard external sort, and insert (in random order) and scan it from RocksDB. RocksDB vs Pure-go: We will also try, as a stretch goal, to measure the RocksDB overhead against a pure-Go implementation, such as . The primary concern in utilizing disk is contention with the existing transactional RocksDB storage engine for Disk IO. We want to allow for the use of an explicit flag to denote the temporary storage location to be used, if the node is deployed with multiple disks. The second concern is that the existing RocksDB instance is tuned to be as efficient as possible for reads, and thus trades off much higher write amplification (via aggressive compactions) in exchange for much lower read amplification. This is not the same tradeoff that we want. Finally, RocksDB is very complex, and for this use case we don't need the majority of its database features, so it's potentially overkill. We have to be careful that the overhead is not too high. These concerns lead us to a few rules for implementing the RocksDB options: Use a separate RocksDB instance explicitly for DistSQL processors. By default this instance uses a `tmp_storage` location under the first store's directory, overridable by a command line flag. It is deleted and recreated from scratch whenever the node starts. Disable write-ahead logging (`write_option.disableWAL`) to lower write amplification. Processors carve out temporary keyspaces in this RocksDB instance. Use range deletion tombstones for rapidly deleting a processor's entire keyspace when it is done. Experiment with compaction settings (including the extreme of turning it off completely), since, fundamentally, we need to write data once and read it back once. This is an ephemeral RocksDB instance, and some cloud providers provide non-persistent disks. We want to eventually support configuring nodes to use non-persistent disks for this use case. `storage/engine` currently is tailored almost exclusively for use by the `storage` package. This implementation should reuse as much of the existing `engine` code as possible, modifying (and generalizing the package) where appropriate. Eventually, other users who also need a non-persistent node-local key-value store can use it as well. Using RocksDB to store on-disk data will result in on-disk data being sorted by key. Thus, for example, if a HashJoin evicts to disk, we can recover an extra ordering that was not expected in planning. Currently, there is no way for a DistSQL stream to indicate that tuples are being delivered with additional orderings not specified in the physical plan, but if we could send this information in the producer metadata, downstream processors could potentially take advantage of the additional sorting.