Transactions: FAQ
Dgraph supports distributed ACID transactions through snapshot isolation.
Can we do pre-writes only on leaders?
Seems like a good idea, but has bad implications. If we only do a prewrite in-memory, only on leader, then this prewrite wouldn’t make it to the Raft log, or disk; but would be considered successful.
Then zero could mark the transaction as committed; but this leader could go down, or leadership could change. In such a case, we’d end up losing the transaction altogether despite it having been considered committed.
Therefore, pre-writes do have to make it to disk. And if so, better to propose them in a Raft group.
Consistency Models
[Last updated: Mar 2018] Basing it on this article by aphyr.
- Sequential Consistency: Different users would see updates at different times, but each user would see operations in order.
Dgraph has a client-side sequencing mode, which provides sequential consistency.
Here, let’s replace a “user” with a “client” (or a single process). In Dgraph, each client maintains a linearizable read map (linread map). Dgraph’s data set is sharded into many “groups”. Each group is a Raft group, where every write is done via a “proposal.” You can think of a transaction in Dgraph, to consist of many group proposals.
The leader in Raft group always has the most recent proposal, while replicas could be behind the leader in varying degrees. You can determine this by just looking at the latest applied proposal ID. A leader’s proposal ID would be greater than or equal to some replicas’ applied proposal ID.
linread
map stores a group -> max proposal ID seen, per client. If a client’s
last read had seen updates corresponding to proposal ID X, then linread
map
would store X for that group. The client would then use the linread
map to
inform future reads to ensure that the server servicing the request, has
proposals >= X applied before servicing the read. Thus, all future reads,
irrespective of which replica it might hit, would see updates for proposals >=
X. Also, the linread
map is updated continuously with max seen proposal IDs
across all groups as reads and writes are done across transactions (within that
client).
In short, this map ensures that updates made by the client, or seen by the client, would never be unseen; in fact, they would be visible in a sequential order. There might be jumps though, for e.g., if a value X → Y → Z, the client might see X, then Z (and not see Y at all).
- Linearizability: Each op takes effect atomically at some point between invocation and completion. Once op is complete, it would be visible to all.
Dgraph supports server-side sequencing of updates, which provides linearizability. Unlike sequential consistency which provides sequencing per client, this provide sequencing across all clients. This is necessary to make upserts work across clients. Thus, once a transaction is committed, it would be visible to all future readers, irrespective of client boundaries.
- Causal consistency: Dgraph does not have a concept of dependencies among transactions. So, does NOT order based on dependencies.
- Serializable consistency: Dgraph does NOT allow arbitrary reordering of transactions, but does provide a linear order per key.
{{% notice “outdated” %}}Sections below this one are outdated. You will find Tour of Dgraph a much helpful resource.{{% /notice %}}
Concepts
Edges
Typical data format is RDF NQuad which is:
Subject, Predicate, Object, Label
, akaEntity, Attribute, Other Entity / Value, Label
Both the terminologies get used interchangeably in our code. Dgraph considers edges to be directional,
i.e. from Subject -> Object
. This is the direction that the queries would be run.
{{% notice “tip” %}}Dgraph can automatically generate a reverse edge. If the user wants to run queries in that direction, they would need to define the reverse edge as part of the schema.{{% /notice %}}
Internally, the RDF NQuad gets parsed into this format.
type DirectedEdge struct {
Entity uint64
Attr string
Value []byte
ValueType uint32
ValueId uint64
Label string
Lang string
Op DirectedEdge_Op // Set or Delete
Facets []*facetsp.Facet
}
Note that irrespective of the input, both Entity
and Object/ValueId
get converted in UID
format
as explained in XID <-> UID.
Posting List
Conceptually, a posting list contains all the DirectedEdges
corresponding to an Attribute
, in the
following format:
Attribute: Entity -> sorted list of ValueId // Everything in uint64 representation.
So, for, e.g., if we’re storing a list of friends, such as:
Entity | Attribute | ValueId |
---|---|---|
Me | friend | person0 |
Me | friend | person1 |
Me | friend | person2 |
Me | friend | person3 |
Then a posting list friend
would be generated. Seeking for Me
in this PL
would produce a list of friends, namely [person0, person1, person2, person3]
.
The big advantage of having such a structure is that we have all the data to do one join in one Posting List. This means, one RPC to the machine serving that Posting List would result in a join, without any further network calls, reducing joins to lookups.
Implementation wise, a Posting List
is a list of Postings
. This is how they look in
Protocol Buffers format.
message Posting {
fixed64 uid = 1;
bytes value = 2;
enum ValType {
DEFAULT = 0;
BINARY = 1;
INT = 2; // We treat it as int64.
FLOAT = 3;
BOOL = 4;
DATE = 5;
DATETIME = 6;
GEO = 7;
UID = 8;
PASSWORD = 9;
STRING = 10;
}
ValType val_type = 3;
enum PostingType {
REF=0; // UID
VALUE=1; // simple, plain value
VALUE_LANG=2; // value with specified language
// VALUE_TIMESERIES=3; // value from timeseries, with specified timestamp
}
PostingType posting_type = 4;
bytes metadata = 5; // for VALUE_LANG: Language, for VALUE_TIMESERIES: timestamp, etc..
string label = 6;
uint64 commit = 7; // More inclination towards smaller values.
repeated facetsp.Facet facets = 8;
// TODO: op is only used temporarily. See if we can remove it from here.
uint32 op = 12;
}
message PostingList {
repeated Posting postings = 1;
bytes checksum = 2;
uint64 commit = 3; // More inclination towards smaller values.
}
There is typically more than one Posting in a PostingList.
The RDF Label is stored as label
in each posting.
{{% notice “warning” %}}We don’t currently retrieve label via query — but would use it in the future.{{% /notice %}}
Badger
PostingLists are served via Badger, given the latter provides enough knobs to decide how much data should be served out of memory, SSD or disk. Also, it supports bloom filters on keys, which makes random lookups efficient.
To allow Badger full access to memory to optimize for caches, we’ll have one Badger instance per machine. Each instance would contain all the posting lists served by the machine.
Posting Lists get stored in Badger, in a key-value format, like so:
(Predicate, Subject) --> PostingList
Group
A set of Posting Lists sharing the same Predicate
constitute a group. Each server can serve
multiple distinct groups.
A group config file is used to determine which server would serve what groups. In the future versions, live Dgraph alpha would be able to move tablets around depending upon heuristics.
If a groups gets too big, it could be split further. In this case, a single Predicate
essentially
gets divided across two groups.
Original Group:
(Predicate, Sa..z)
After split:
Group 1: (Predicate, Sa..i)
Group 2: (Predicate, Sj..z)
Note that keys are sorted in BadgerDB. So, the group split would be done in a way to maintain that sorting order, i.e. it would be split in a way where the lexicographically earlier subjects would be in one group, and the later in the second.
Replication and Server Failure
Each group should typically be served by atleast 3 servers, if available. In the case of a machine failure, other servers serving the same group can still handle the load in that case.
New Server and Discovery
Dgraph cluster can detect new machines allocated to the cluster, establish connections, and transfer a subset of existing predicates to it based on the groups served by the new machine.
Write Ahead Logs
Every mutation upon hitting the database doesn’t immediately make it on disk via BadgerDB. We avoid
re-generating the posting list too often, because all the postings need to be kept sorted, and it’s
expensive. Instead, every mutation gets logged and synced to disk via append only log files called
write-ahead logs
. So, any acknowledged writes would always be on disk. This allows us to recover
from a system crash, by replaying all the mutations since the last write to Posting List
.
Mutations
In addition to being written to Write Ahead Logs
, a mutation also gets stored in memory as an
overlay over immutable Posting list
in a mutation layer. This mutation layer allows us to iterate
over Posting
s as though they’re sorted, without requiring re-creating the posting list.
When a posting list has mutations in memory, it’s considered a dirty
posting list. Periodically,
we re-generate the immutable version, and write to BadgerDB. Note that the writes to BadgerDB are
asynchronous, which means they don’t get flushed out to disk immediately, but that wouldn’t lead
to data loss on a machine crash. When Posting lists
are initialized, write-ahead logs get referred,
and any missing writes get applied.
Every time we regenerate a posting list, we also write the max commit log timestamp that was included — this helps us figure out how long back to seek in write-ahead logs when initializing the posting list, the first time it’s brought back into memory.
Queries
Let’s understand how query execution works, by looking at an example.
me(id: m.abcde) {
pred_A
pred_B {
pred_B1
pred_B2
}
pred_C {
pred_C1
pred_C2 {
pred_C21
}
}
}
Let’s assume we have 3 server instances, and instance id = 2 receives this query. These are the steps:
- Determine the UID of provided XID, in this case
m.abcde
using fingerprinting. Say the UID = u. - Send queries to look up keys =
pred_A, u
,pred_B, u
, andpred_C, u
. These predicates could belong to 3 different groups, served by potentially different servers. So, this would typically incur at max 3 network calls (equal to number of predicates at this step). - The above queries would return back 3 list of ids or value. The result of
pred_B
andpred_C
would be converted into queries forpred_Bi
andpred_Ci
. pred_Bi
andpred_Ci
would then cause at max 4 network calls, depending upon where these predicates are located. The keys forpred_Bi
for e.g. would bepred_Bi, res_pred_Bk
, where res_pred_Bk = list of resulting ids frompred_B, u
.- Looking at
res_pred_C2
, you’ll notice that this would be a list of lists aka list matrix. We merge these list of lists into a sorted list with distinct elements to form the query forpred_C21
. - Another network call depending upon where
pred_C21
lies, and this would again give us a list of list ids / value.
If the query was run via HTTP interface /query
, this subgraph gets converted into JSON for
replying back to the client. If the query was run via gRPC interface using
the language clients, the subgraph gets converted to
protocol buffer format, and returned to client.
Network Calls
Compared to RAM or SSD access, network calls are slow.
Dgraph minimizes the number of network calls required to execute queries. As explained above, the
data sharding is done based on predicate
, not entity
. Thus, even if we have a large set of
intermediate results, they’d still only increase the payload of a network call, not the number of
network calls itself. In general, the number of network calls done in Dgraph is directly proportional
to the number of predicates in the query, or the complexity of the query, not the number of
intermediate or final results.
In the above example, we have eight predicates, and so including a call to convert to UID, we’ll have at max nine network calls. The total number of entity results could be in millions.
Worker
In Queries section, you noticed how the calls were made to query for (predicate, uids)
. All those
network calls / local processing are done via workers. Each server exposes a
gRPC interface, which can then be called by the query processor to retrieve data.
Worker Pool
Worker Pool is just a pool of open TCP connections which can be reused by multiple goroutines. This avoids having to recreate a new connection every time a network call needs to be made.
Protocol Buffers
All data in Dgraph that is stored or transmitted is first converted into byte arrays through serialization using Protocol Buffers. When the result is to be returned to the user, the protocol buffer object is traversed, and the JSON object is formed.
Minimizing network calls explained
To explain how Dgraph minimizes network calls, let’s start with an example query we should be able to run.
Find all posts liked by friends of friends of mine over the last year, written by a popular author X.
SQL/NoSQL
In a distributed SQL/NoSQL database, this would require you to retrieve a lot of data.
Method 1:
- Find all the friends (~ 338 friends).
- Find all their friends (~ 338 * 338 = 40,000 people).
- Find all the posts liked by these people over the last year (resulting set in millions).
- Intersect these posts with posts authored by person X.
Method 2:
- Find all posts written by popular author X over the last year (possibly thousands).
- Find all people who liked those posts (easily millions)
result set 1
. - Find all your friends.
- Find all their friends
result set 2
. - Intersect
result set 1
withresult set 2
.
Both of these approaches would result in a lot of data going back and forth between database and application; would be slow to execute, or would require you to run an offline job.
Dgraph
This is how it would run in Dgraph:
- Node X contains posting list for predicate
friends
. - Seek to caller’s userid in Node X (1 RPC). Retrieve a list of friend uids.
- Do multiple seeks for each of the friend uids, to generate a list of friends of friends uids.
result set 1
- Node Y contains posting list for predicate
posts_liked
. - Ship result set 1 to Node Y (1 RPC), and do seeks to generate a list of all posts liked by
result set 1.
reult set 2
- Node Z contains posting list for predicate
author
. - Ship result set 2 to Node Z (1 RPC). Seek to author X, and generate a list of posts authored
by X.
result set 3
- Intersect the two sorted lists,
result set 2
andresult set 3
.result set 4
- Node N contains names for all uids.
- Ship
result set 4
to Node N (1 RPC), and convert uids to names by doing multiple seeks.result set 5
- Ship
result set 5
back to caller.
In 4-5 RPCs, we have figured out all the posts liked by friends of friends, written by popular author X.
This design allows vast scalability, and yet consistent production level latencies, to support running complicated queries requiring deep joins.
RAFT
This section aims to explain the RAFT consensus algorithm in simple terms. The idea is to give you just enough to make you understand the basic concepts, without going into explanations about why it works accurately. For a detailed explanation of RAFT, please read the original thesis paper by Diego Ongaro.
Term
Each election cycle is considered a term, during which there is a single leader (just like in a democracy). When a new election starts, the term number is increased. This is straightforward and obvious but is a critical factor for the accuracy of the algorithm.
In rare cases, if no leader could be elected within an ElectionTimeout
, that term can end without
a leader.
Server States
Each server in cluster can be in one of the following three states:
- Leader
- Follower
- Candidate
Generally, the servers are in leader or follower state. When the leader crashes or the communication breaks down, the followers will wait for election timeout before converting to candidates. The election timeout is randomized. This would allow one of them to declare candidacy before others. The candidate would vote for itself and wait for the majority of the cluster to vote for it as well. If a follower hears from a candidate with a higher term than the current (dead in this case) leader, it would vote for it. The candidate who gets majority votes wins the election and becomes the leader.
The leader then tells the rest of the cluster about the result (Heartbeat Communication) and the other candidates then become followers. Again, the cluster goes back into leader-follower model.
A leader could revert to being a follower without an election, if it finds another leader in the cluster with a higher Term). This might happen in rare cases (network partitions).
Communication
There is unidirectional RPC communication, from leader to followers. The followers never ping the
leader. The leader sends AppendEntries
messages to the followers with logs containing state
updates. When the leader sends AppendEntries
with zero logs, that’s considered a
Heartbeat. Leader sends all followers Heartbeats at regular intervals.
If a follower doesn’t receive Heartbeat for ElectionTimeout
duration (generally between
150ms to 300ms), it converts it’s state to candidate (as mentioned in Server States*.
Every communication request contains a term number. If a server receives a request with a stale term number, it rejects the request.
Raft believes in retrying RPCs indefinitely.
Log Entries
Log Entries are numbered sequentially and contain a term number. Entry is considered committed if it has been replicated to a majority of the servers.
On receiving a client request, the leader does four things (aka Log Replication):
- Appends and persists to its log.
- Issue
AppendEntries
in parallel to other servers. - On majority replication, consider the entry committed and apply to its state machine.
- Notify followers that entry is committed so that they can apply it to their state machines.
A leader never overwrites or deletes its entries. There is a guarantee that if an entry is committed, all future leaders will have it. A leader can, however, force overwrite the followers’ logs, so they match leader’s logs (elected democratically, but got a dictator).
Voting
Each server persists its current term and vote, so it doesn’t end up voting twice in the same term.
On receiving a RequestVote
RPC, the server denies its vote if its log is more up-to-date than the
candidate. It would also deny a vote, if a minimum ElectionTimeout
hasn’t passed since the last
Heartbeat from the leader. Otherwise, it gives a vote and resets its ElectionTimeout
timer.
Up-to-date property of logs is determined as follows:
- Term number comparison
- Index number or log length comparison
{{% notice “tip” %}}To understand the above sections better, you can see this interactive visualization.{{% /notice %}}
Cluster membership
Raft only allows single-server changes, i.e. only one server can be added or deleted at a time.
This is achieved by cluster configuration changes. Cluster configurations are communicated using
special entries in AppendEntries
.
The significant difference in how cluster configuration changes are applied compared to how typical Log Entries are applied is that the followers don’t wait for a commitment confirmation from the leader before enabling it.
A server can respond to both AppendEntries
and RequestVote
, without checking current
configuration. This mechanism allows new servers to participate without officially being part of
the cluster. Without this feature, things won’t work.
When a new server joins, it won’t have any logs, and they need to be streamed. To ensure cluster availability, Raft allows this server to join the cluster as a non-voting member. Once it’s caught up, voting can be enabled. This also allows the cluster to remove this server in case it’s too slow to catch up, before giving voting rights (sort of like getting a green card to allow assimilation before citizenship is awarded providing voting rights).
{{% notice “tip” %}}If you want to add a few servers and remove a few servers, do the addition before the removal. To bootstrap a cluster, start with one server to allow it to become the leader, and then add servers to the cluster one-by-one.{{% /notice %}}
Log Compaction
One of the ways to do this is snapshotting. As soon as the state machine is synced to disk, the logs can be discarded.
Clients
Clients must locate the cluster to interact with it. Various approaches can be used for discovery.
A client can randomly pick up any server in the cluster. If the server isn’t a leader, the request should be rejected, and the leader information passed along. The client can then re-route it’s query to the leader. Alternatively, the server can proxy the client’s request to the leader.
When a client first starts up, it can register itself with the cluster using RegisterClient
RPC.
This creates a new client id, which is used for all subsequent RPCs.
Linearizable Semantics
Servers must filter out duplicate requests. They can do this via session tracking where they use the client id and another request UID set by the client to avoid reprocessing duplicate requests. RAFT also suggests storing responses along with the request UIDs to reply back in case it receives a duplicate request.
Linearizability requires the results of a read to reflect the latest committed write. Serializability, on the other hand, allows stale reads.
Read-only queries
To ensure linearizability of read-only queries run via leader, leader must take these steps:
- Leader must have at least one committed entry in its term. This would allow for up-to-dated-ness. (C’mon! Now that you’re in power do something at least!)
- Leader stores it’s latest commit index.
- Leader sends Heartbeats to the cluster and waits for ACK from majority. Now it knows that it’s the leader. (No successful coup. Yup, still the democratically elected dictator I was before!)
- Leader waits for its state machine to advance to readIndex.
- Leader can now run the queries against state machine and reply to clients.
Read-only queries can also be serviced by followers to reduce the load on the leader. But this could lead to stale results unless the follower confirms that its leader is the real leader(network partition). To do so, it would have to send a query to the leader, and the leader would have to do steps 1-3. Then the follower can do 4-5.
Read-only queries would have to be batched up, and then RPCs would have to go to the leader for each batch, who in turn would have to send further RPCs to the whole cluster. (This is not scalable without considerable optimizations to deal with latency.)
An alternative approach would be to have the servers return the index corresponding to their state machine. The client can then keep track of the maximum index it has received from replies so far. And pass it along to the server for the next request. If a server’s state machine hasn’t reached the index provided by the client, it will not service the request. This approach avoids inter-server communication and is a lot more scalable. (This approach does not guarantee linearizability, but should converge quickly to the latest write.)