Sunday, November 24, 2013

The History of RocksDB

The Past
It was mid 2011, I had been developing HDFS/HBase for five years and I was in love with the Hadoop ecosystem. Here is a system that can store and query hundreds of petabytes of data without blinking an eye. Hadoop has always focussed on scalability first and Hadoop's scalability was increasing by leaps and bounds. It was apparent that we will be able to store an exabyte of data in a single Hadoop cluster in a few years. I wanted to take on a new see if we can extend HDFS's success story from Data Analytics to Query Serving workloads as well.

A Query Serving workload mostly consists of point lookups, random short range scans and point updates. The primary requirement of this workload is low-latency queries. On the other hand, a Big Data Analytics query typically involves large sequential scans and joins of two or more data sets with a very low volume of updates if any. Thus, the following year I spent comparing HBase/HDFS and MySQL for a Query Serving workload. The advantage of using HBase is that we can have multiple copies of data within a single data center.  I wanted to determine what is needed to migrate a very large Query Serving workload from a cluster of MySQL servers to an HBase/HDFS cluster. This multi-petabyte dataset was stored on spinning disks. After multiple rounds of enhancements to HBase and were able to make it such that HBase latencies were only twice as slow as a MySQL server and HBase was using only three times more IOPs to serve the same workload on the same hardware. We were making steady progress towards our goal... but then something changed!

Flash storage became a reality. The data set migrated from spinning disks to flash. Now, the million dollar question that came up is whether HBase is capable of efficiently using flash hardware. I benchmarked HDFS and HBase with data on flash storage and posted the results in an earlier post.The short story is that the HDFS/HBase of 2012 had a few software bottlenecks because of which it was not able to use flash storage efficiently. It became clear that if data is stored in flash storage, then we need a new storage engine to be able to serve a random workload on that data efficiently. I started to look out for techniques to build the next generation key-value store, especially designed to serve data from fast storage.

Why do we need an Embedded Database?
Flash is fundamentally different from spinning storage in performance. For the purpose of this discussion, let's assume that a read or write to spinning disk takes about 10 milliseconds while a read or write to flash storage takes about 100 microseconds.  Network network-latency between two machines remains around 50 microseconds. These numbers are not cast in stone and your hardware could be very different from this one, but these numbers demonstrate the relative differences between two scenarios. What does this have anything to do with application-systems architecture? A client wants to store and access data from a database. There are two alternatives, it can store data on locally attached disks or it can store data over the network on a remote server that have disks attached to it. If we consider latency, then the locally attached disks can serve a read request in about 10 milliseconds. And in the client-server architecture, accessing the same data over a network results in a latency of 10.05 milliseconds, the overhead imposed by the network being only a miniscule 0.5%.  Given this fact, it is easy to understand why a majority of currently-deployed systems use the client-server model of accessing data. (For the purpose of the discussion, I am ignoring network bandwidth limitations).

Now, lets consider the same scenario but with disks replaced by flash drives. A data access in the case of locally attached flash storage is 100 microseconds whereas accessing the same data via the network is 150 micros. Network data access is 50% higher overhead than local data access and 50% is a pretty big number. This means that databases that run embedded within an application could have much lower latency than applications that access data over a network. Thus, the necessity of an Embedded Database.

The above hypothesis does not state that client-server models will become extinct. The client-server-data-access-model has inherent advantages in the areas of data management and will continue to remain prominent in application deployment scenarios.

Aren't there any existing embedded databases?
Of course there are many existing embedded  databases: BerkeleyDB, KyotoDB, SQLite3, leveldb, etc. Open-source benchmarks seem to state that leveldb is the fastest of the lot. But not all of them are  suitable for storing data on Flash storage. Flash has limited write-endurance; updates to a block of data on Flash typically introduces write-amplification within the Flash driver. Given that we want to run our database on flash, we focussed on measuring write-amplification for evaluating databases technologies.

HBase and Cassandra are Log Structured Merge (LSM)  style databases, but it will take a lot of engineering work to make HBase and Cassandra be an embeddable library. Both of them are servers with an entire ecosystem of built-in management, configuration and deployments. I was looking for a simple c/c++ library: leveldb was the apparent first choice for our benchmarking.

Why was leveldb insufficient for our purpose?
I started to benchmark leveldb and found that it was unsuitable for our server workloads. Leveldb was a cool piece of work but was not designed for server workloads. The open-source benchmark results looks awesome at first sight, but I quickly realized that those results were for a database whose size was smaller than the size of RAM on the test machine, i.e. the entire database has to fit in the OS page cache. When I performed the same benchmarks on a database that was at least 5 times larger than main memory, the performance results was dismal.

Leveldb's single-threaded compaction process was insufficient to drive server workloads. Frequent write-stalls caused 99-percentile latency to be tremendously large. Mmap-ing a file into the OS cache introduced performance bottlenecks for reads. Leveldb was unable to consume all the IOs offered by the underlying flash storage.

On the other hand, I was seeing server storage hardware evolve fast in different dimensions, For example, an experimental system where a storage volume is striped across 10 flash cards can provide upwards of a million IOs per second. A NVRAM based storage can support a few million data accesses per second. I would like to use a database that can drive these types of fast storage hardware. A natural evolution of flash storage could lead us to storage that has a very limited erase cycles and I envisioned that a database that can allow elegant tradeoffs of read amplification, write amplification and space amplification would be a dire necessity for these type of storage. Leveldb was not designed to achieve these goals. The best path was to fork the leveldb code and change its architecture to suit these needs. So, RocksDB was born!

The vision for RocksDB
1. An embedded key-value store with point lookups and range scans
2. Optimized for fast storage, e.g. Flash and RAM
3. Server Side database with full production support
4. Scale linearly with number of CPU cores and with storage IOPs

RocksDB is not a distributed database. It does not have fault-tolerance or replication built into it. It does not know anything about data-sharding. It is upto the application that is using RocksDB to implement replication, fault-tolerance and sharding if needed.

Architecture of RocksDB
RocksDB is a C++ library that can be used to persistently store keys and values. Keys and values are arbitrary byte streams. Keys are stored in sorted runs. New writes occur to new places in the storage and a background compaction process eliminates duplicates and processes delete markers. There is support for atomically writing a set of keys into the database. Backward and forward iteration over the keys are supported.

RockDB is built using a "pluggable" architecture. This makes it easy to replace parts of it without impacting the overall architecture of the system. This architecture makes me confident that RocksDB will be easily tunable for different workloads and on different hardware.

For example, one can plug in various compression modules (snappy, zlib, bzip, etc) without changing any RocksDB code.  Similarly, an application can plug in their own compaction filter to process keys during compaction; an example application can use it to implement a expiry-time for keys in the database. RocksDB has pluggable memtables so that an application can design a custom data structure to cache their writes, one example is prefix-hash-memtable where a portion of the key is hashed and the remainder of the key is arranged in the form of a btree. The implementation of a sst file is pluggable too and an application can design their own format for sst files. RocksDB supports a MergeType record that allows applications to build higher level constructs (Lists, Counters, etc) by avoiding a read-modify-write.

RocksDB currently supports two styles of compaction: the level style compaction and the universal stye compaction. These two styles offers flexible performance tradeoffs. Compactions are inherently multi-threaded so that large databases can sustain high update rates. I will write a separate post on the pros and cons of these two styles of compaction.

RocksDB exposes interfaces for incremental online backup which is needed for any kind of production usage. It supports setting bloom filters on a sub-part of the key, which is one possible technique to reduce iops needed for range-scans.

RocksDB's apis are stackable, which means that you can wrap lower level apis with higher level easy-to-use wrappers. This is the topic of a future post.

Potential Use-cases of RocksDB
RocksDB can be used by applications that need low latency database accesses. A user-facing application that stores the viewing history and state of users of a website can potentially store this content on RocksDB. A spam detection application that needs fast access to big data sets can use RocksDB. A graph-search query that needs to scan a data set in realtime can use RocksDB. RocksDB can be used to cache data from Hadoop, thereby allowing applications to query Hadoop data in realtime. A message-queue that supports a high number of inserts and deletes can use RocksDB.

The Road Ahead
There is work-in-progress to make RocksDB be able to serve data at memory speeds when the entire database fits in RAM.  RocksDB would evolve to be compatible with highly multi-core machines. With the advent of super-low-endurance-flash storage, we expect RocksDB to store data with minimal write amplification. Maybe sometime in the future, someone will port RocksDB to the Android and iOS Platform. Another nice feature would be to support Column Families to support better clustering of related data.

I am hoping that software programmers and database developers would use, enhance and customize RocksDB for their use-cases. The code base is in You can join the Facebook Group to participate in the engineering design discussions about RocksDB.

Wednesday, October 16, 2013

The SPARROW Theorem for performance of storage systems

I have been measuring the performance of a number of storage systems and here is a post of what I have learnt so far. I do not claim that any of this is new, it is written from my perspective and is a very opinionated piece.

Database and Storage Systems have evolved over the years, starting from storing data in spinning-disk based storage system to solid-state-storage (SSD) and random-access-memory (RAM) storage. These storage hardware have very different performance characteristics. The difference in the physical constraints of these storage hardware means that we might need to use different methodologies to measure their performance. In the following discussions, I refer to OLTP workloads where each record is small (smaller than 4K) and they are read randomly.

Spinning Disks
Spinning disks can typically sustain 100 to 200 IO per sec. Random accesses are typically bottlenecked by the number of IOs that the disk can service. The sequential access speeds of disks can reach upto 80 MBytes/sec. Most database system that run on spinning disks are optimized to reduce the number of random reads to storage that it consumes to service a specified workload. Btree implementations (like Innodb, Berkeley-DB, etc) tries to reduce the number of random reads by caching some part of the working set in RAM. LSM databases (like Apache HBase) converts random writes to sequential writes to get better database performance. For all these systems,  reducing the number of random disk access directly improves performance, especially when the workload is mostly small random reads or short scans.

Solid State Devices (SSD)
SSDs have physical constraints that are different from spinning disks. An SSD can typically perform about 60K to 100K IO per second. This is orders of magnitude larger than what a spinning disk can possibly do. The throughout of an SSD can very from 300MBytes/sec to 600 MBytes/sec depending on the manufacturer and the percentage of free space on the SSD.

If a storage software issues 100K IO per second and each IO is 4 K, then the total data bandwidth needed is 400 MBytes/sec which is close to the maximum throughput of the SSD. If you run a OLTP benchmark on SSDs, it is likely that you are bottlenecked because of one of two reasons:
     (1) you have used up all the random IOs per second offered by the device or
    (2) you have maxed out the SSD data bandwidth.
It could be either of these two constraints. If you are reaching the maximum data throughput of this device, then reducing the number of random accesses to storage might not improve performance for your system. You need to reduce the total storage bandwidth too.

So, what consumes the storage bandwidth of the device? This bandwidth is consumed by reads and writes done by the storage software on behalf of the application.

Read Amplification (RA)
Read Amplification is the ratio of the number of storage bytes accessed to satisfy a single read request of 1 byte. For example, if a btree-based storage system has to read a single 4K size page from storage (assuming that the index pages are cached in RAM) for every read request of 1 byte, then it has a read amplification of 4K. On the other hand, if an LSM based database needs to consult three storage files to satisfy a single read request of size 1 byte and the block size of each file is 4K, then its read amplification is 12K (ignoring the existence of bloom filters). My colleague Mark Callaghan has an awesome post about Read Amplification.

Write Amplification (WA)
Write Amplification is the ratio of the number of storage bytes accessed to satisfy a single 1 byte write request to the database. WA includes the write amplification caused by the database software as well as the write amplification generated by the SSD itself.

For example, if a btree-based storage system has to write 1 byte into the database, it has to read in the 4K page into memory, update it and then write it back, thereby incurring a write amplification of 8K. A btree-based storage-system typically overwrites the SSD page in place thereby incurring additional write amplification in the SSD layer too.

On the other hand, an LSM storage engine typically writes the new 1 byte to a new place on SSD and has a write amplification of 1 (not including compaction).

For transaction logging, a btree-based system needs a redo log and an undo log which means that WA is further increased by 2 for these systems. An LSM based system needs only an redo log which causes WA to increase by 1.

But the above does not mean that LSM systems are better. Before I try to explain why it is so, please allow me to write about Space Amplification.

Space Amplification (SA)
Space Amplification is the number of storage bytes needed to store 1 byte of information. Storing multiple indices of the same data increase the storage requirements but could decrease read latencies. Space Amplification is also caused by the internal fragmentation and padding. An LSM database can have the same key with older versions of the data in multiple files... this too can cause Space Amplification.

The SPAce, Read Or Write theorem (SPARROW)

The SPARROW Theorem states that:
1. RA is inversely related to WA
2. WA is inversely related to SA

If you want to reduce the Read Amplification caused by a specific workload, then you can achieve it only if you incur higher Write Amplification. Conversely, if you want to reduce the Write Amplification caused by a specific workload, then you have to incur higher Read Amplification to achieve that goal. (This assumes that we maintain Space Amplification constant at all times)

Similarly, the only way to reduce Space Amplification caused by a specific workload is to tune the system in such a way that it can sustain a higher Write Amplification. Conversely, if you want to reduce Write Amplification, then you have to incur higher Space Amplification for the same workload.

Implications of the SPARROW Theorem
There isn't any storage system that can escape from the clutches of the SPARROW Theorem. A single system architecture CANNOT reduce all three SA, RA and WA. I am talking about the architecture and not of implementations.

A practical Database implementations would always try to reduce its current SA, WA and RA by optimizing its code and algorithms. But once all it's code and algorithms are optimized, then it won't be able to improve all three dimensions at the same time. Its performance will be confined by the walls outlined by the SPARROW Theorem.

Given the above fact, it would be great to have most database system be configurable so that an administrator can tune each of these three dimensions based on the workload and the hardware that it is running on. This will result in a highly flexible database system architecture that can sustain a myriad of workloads.