- Leaders and Followers
- Replication Logs
- Multi-Leader Replication
- Leaderless Replication
Replication means keeping a copy of the same data on multiple machines that are connected via a network. There are several reasons why you might want to replicate data:
• To keep data geographically close to your users (and thus reduce latency)
• To allow the system to continue working even if some of its parts have failed (and thus increase availability)
• To scale out the number of machines that can serve read queries (and thus increase read throughput)
If the data that you’re replicating does not change over time, then replication is easy: you just need to copy the data to every node once, and you’re done.
All the difficulty in replication lies in handling changes to replicated data, and that’s what this chapter is about.
There are three popular algorithms for replicating changes between nodes: single-leader, multi-leader, and leaderless replication. Almost all distributed databases use one of these three approaches.
There are many trade-offs to consider with replication: for example, whether to use synchronous or asynchronous replication, and how to handle failed replicas.
Those are often configuration options in databases, and although the details vary by data-base, the general principles are similar across many implementations.
Leaders and Followers
Each node that stores a copy of the database is called a replica. With multiple replicas, a question inevitably arises: how do we ensure that all the data ends up on all the replicas?
Every write to the database needs to be processed by every replica; otherwise, the replicas would no longer contain the same data. The most common solution for this is called leader-based replication (also known as active/passive or master–slave replication)
One of the replicas is designated the leader (also known as master or primary). When clients want to write to the database, they must send their requests to the leader, which first writes the new data to its local storage.
The other replicas are known as followers (read replicas, slaves, secondaries, or hot standbys).Whenever the leader writes new data to its local storage, it also sends the data change to all of its followers as part of a replication log or change stream. Each follower takes the log from the leader and updates its local copy of the data‐base accordingly, by applying all writes in the same order as they were processed on the leader.
When a client wants to read from the database, it can query either the leader or any of the followers. However, writes are only accepted on the leader (the follow‐ ers are read-only from the client’s point of view).
Synchronous Versus Asynchronous Replication
An important detail of a replicated system is whether the replication happens syn‐ chronously or asynchronously.
Normally, replication is quite fast: most database systems apply changes to followers in less than a second. However, there is no guarantee of how long it might take. There are circumstances when followers might fall behind the leader by several minutes or more;
The advantage of synchronous replication is that the follower is guaranteed to have an up-to-date copy of the data that is consistent with the leader.
If the leader suddenly fails, we can be sure that the data is still available on the follower. The disadvantage is that if the synchronous follower doesn’t respond (because it has crashed, or there is a network fault, or for any other reason), the write cannot be processed.
The leader must block all writes and wait until the synchronous replica is available again.
For that reason, it is impractical for all followers to be synchronous: any one node outage would cause the whole system to grind to a halt.
In practice, if you enable synchronous replication on a database, it usually means that one of the followers is synchronous, and the others are asynchronous. If the synchronous follower becomes unavailable or slow, one of the asynchronous followers is made synchronous. This guarantees that you have an up-to-date copy of the data on at least two nodes: the leader and one synchronous follower. This configuration is sometimes also called semi-synchronous.
Often, leader-based replication is configured to be completely asynchronous. In this case, if the leader fails and is not recoverable, any writes that have not yet been replicated to followers are lost.
This means that a write is not guaranteed to be durable, even if it has been confirmed to the client. However, a fully asynchronous configuration has the advantage that the leader can continue processing writes, even if all of its followers have fallen behind.
Weakening durability may sound like a bad trade-off, but asynchronous replication is nevertheless widely used, especially if there are many followers or if they are geo‐ graphically distributed.
Setting Up New Followers
From time to time, you need to set up new followers—perhaps to increase the number of replicas, or to replace failed nodes. How do you ensure that the new follower has an accurate copy of the leader’s data?
Simply copying data files from one node to another is typically not sufficient: clients are constantly writing to the database, and the data is always in flux, so a standard file copy would see different parts of the database at different points in time. The result might not make any sense.
Take a consistent snapshot of the leader’s database at some point in time—if possible, without taking a lock on the entire database. Most databases have this fea‐ ture, as it is also required for backups. In some cases, third-party tools are needed, such as innobackupex for MySQL.
Copy the snapshot to the new follower node.
The follower connects to the leader and requests all the data changes that have happened since the snapshot was taken. This requires that the snapshot is associated with an exact position in the leader’s replication log. That position has various names: for example, PostgreSQL calls it the log sequence number, and MySQL calls it the binlog coordinates.
When the follower has processed the backlog of data changes since the snapshot, we say it has caught up. It can now continue to process data changes from the leader as they happen.
Handling Node Outages
Being able to reboot individual nodes without downtime is a big advantage for operations and maintenance. Our goal is to keep the system as a whole running despite individual node failures, and to keep the impact of a node outage as small as possible.
Follower failure: Catch-up recovery
On its local disk, each follower keeps a log of the data changes it has received from the leader. If a follower crashes and is restarted, or if the network between the leader and the follower is temporarily interrupted, the follower can recover quite easily: from its log, it knows the last transaction that was processed before the fault occur‐ red.
Thus, the follower can connect to the leader and request all the data changes that occurred during the time when the follower was disconnected. When it has applied these changes, it has caught up to the leader and can continue receiving a stream of data changes as before.
Leader failure: Failover
Handling a failure of the leader is trickier: one of the followers needs to be promoted to be the new leader, clients need to be reconfigured to send their writes to the new leader, and the other followers need to start consuming data changes from the new leader. This process is called failover.
Failover can happen manually (an administrator is notified that the leader has failed and takes the necessary steps to make a new leader) or automatically. An automatic failover process usually consists of the following steps:
Determining that the leader has failed.
Choosing a new leader.
Reconfiguring the system to use the new leader.
Failover is fraught with things that can go wrong:
If asynchronous replication is used, the new leader may not have received all the writes from the old leader before it failed.
Discarding writes is especially dangerous if other storage systems outside of the database need to be coordinated with the database contents.
In certain fault scenarios, it could happen that two nodes both believe that they are the leader. This situation is called split brain, and it is dan‐ gerous
What is the right timeout before the leader is declared dead? A longer timeout means a longer time to recovery in the case where the leader fails. However, if the timeout is too short, there could be unnecessary failovers.
There are no easy solutions to these problems. For this reason, some operations teams prefer to perform failovers manually, even if the software supports automatic failover.
Implementation of Replication Logs
In the simplest case, the leader logs every write request (statement) that it executes and sends that statement log to its followers.
Although this may sound reasonable, there are various ways in which this approach to replication can break down:
Any statement that calls a nondeterministic function, such as NOW() to get the current date and time or RAND() to get a random number, is likely to generate a different value on each replica.
If statements use an autoincrementing column, or if they depend on the existing data in the database (e.g., UPDATE … WHERE
), they must be executed in exactly the same order on each replica, or else they may have a differ‐ ent effect. This can be limiting when there are multiple concurrently executing transactions.
Statements that have side effects (e.g., triggers, stored procedures, user-defined functions) may result in different side effects occurring on each replica, unless the side effects are absolutely deterministic.
the leader can replace any nondeterministic function calls with a fixed return value when the statement is logged so that the followers all get the same value. However, because there are so many edge cases, other replication methods are now generally preferred.
Write-ahead log (WAL) shipping
Usually every write is appended to a log: • In the case of a log-structured storage engine, this log is the main place for storage. Log segments are compacted and garbage-collected in the background.
• In the case of a B-tree, which overwrites individual disk blocks, every modification is first written to a write-ahead log so that the index can be restored to a consistent state after a crash.
We can use the exact same log to build a replica on another node: besides writing the log to disk, the leader also sends it across the network to its followers. When the follower processes this log, it builds a copy of the exact same data structures as found on the leader.
Logical (row-based) log replication
An alternative is to use different log formats for replication and for the storage engine, which allows the replication log to be decoupled from the storage engine internals.
A logical log for a relational database is usually a sequence of records describing writes to database tables at the granularity of a row:
• For an inserted row, the log contains the new values of all columns.
• For a deleted row, the log contains enough information to uniquely identify the row that was deleted. Typically this would be the primary key, but if there is no primary key on the table, the old values of all columns need to be logged.
• For an updated row, the log contains enough information to uniquely identify the updated row, and the new values of all columns (or at least the new values of all columns that changed).
A transaction that modifies several rows generates several such log records, followed by a record indicating that the transaction was committed. MySQL’s binlog (when configured to use row-based replication) uses this approach
A trigger lets you register custom application code that is automatically executed when a data change (write transaction) occurs in a database system. The trigger has the opportunity to log this change into a separate table, from which it can be read by an external process. That external process can then apply any necessary application logic and replicate the data change to another system.
Trigger-based replication typically has greater overheads than other replication methods, and is more prone to bugs and limitations than the database’s built-in replication.
Problems with Replication Lag
Being able to tolerate node failures is just one reason for wanting replication. Other reasons are scalability (processing more requests than a single machine can handle) and latency (placing replicas geo‐ graphically closer to users).
Leader-based replication requires all writes to go through a single node, but read- only queries can go to any replica.
For workloads that consist of mostly reads and only a small percentage of writes (a common pattern on the web), there is an attractive option: create many followers, and distribute the read requests across those fol‐ lowers. This removes load from the leader and allows read requests to be served by nearby replicas.
In this read-scaling architecture, you can increase the capacity for serving read-only requests simply by adding more followers.
Unfortunately, if an application reads from an asynchronous follower, it may see out‐dated information if the follower has fallen behind. This leads to apparent inconsistencies in the database: if you run the same query on the leader and a follower at the same time, you may get different results, because not all writes have been reflected in the follower.
This inconsistency is just a temporary state—if you stop writing to the database and wait a while, the followers will eventually catch up and become consistent with the leader. For that reason, this effect is known as eventual consistency
The term “eventually” is deliberately vague: in general, there is no limit to how far a replica can fall behind.
When the lag is so large, the inconsistencies it introduces are not just a theoretical issue but a real problem for applications.
Reading Your Own Writes
Many applications let the user submit some data and then view what they have sub mitted. When new data is submitted, it must be sent to the leader, but when the user views the data, it can be read from a follower. This is especially appropriate if data is frequently viewed but only occasionally written.
With asynchronous replication, there is a problem, if the user views the data shortly after making a write, the new data may not yet have reached the replica. To the user, it looks as though the data they submitted was lost, so they will be understandably unhappy.
In this situation, we need read-after-write consistency, also known as read-your-writes consistency. This is a guarantee that if the user reloads the page, they will always see any updates they submitted themselves. It makes no promises about other users: other users’ updates may not be visible until some later time. However, it reassures the user that their own input has been saved correctly.
There are various possible techniques. To mention a few:
• When reading something that the user may have modified, read it from the leader; otherwise, read it from a follower.
• If most things in the application are potentially editable by the user, that approach won’t be effective, as most things would have to be read from the leader (negating the benefit of read scaling). In that case, other criteria may be used to decide whether to read from the leader. For example, you could track the time of the last update and, for one minute after the last update, make all reads from the leader.
• The client can remember the timestamp of its most recent write—then the sys‐ tem can ensure that the replica serving any reads for that user reflects updates at least until that timestamp.
• If your replicas are distributed across multiple datacenters (for geographical proximity to users or for availability), there is additional complexity. Any request that needs to be served by the leader must be routed to the datacenter that con‐ tains the leader.
Our second example of an anomaly that can occur when reading from asynchronous followers is that it’s possible for a user to see things moving backward in time.
Monotonic reads is a guarantee that this kind of anomaly does not happen. It’s a lesser guarantee than strong consistency, but a stronger guarantee than eventual consistency. When you read data, you may see an old value; monotonic reads only means that if one user makes several reads in sequence, they will not see time go backward.
One way of achieving monotonic reads is to make sure that each user always makes their reads from the same replica.
Solutions for Replication Lag
When working with an eventually consistent system, it is worth thinking about how the application behaves if the replication lag increases to several minutes or even hours.
It would be better if application developers didn’t have to worry about subtle replica‐ tion issues and could just trust their databases to “do the right thing.” This is why transactions exist: they are a way for a database to provide stronger guarantees so that the application can be simpler.
Leader-based replication has one major downside: there is only one leader, and all writes must go through it. If you can’t connect to the leader for any reason, for example due to a network interruption between you and the leader, you can’t write to the database.
A natural extension of the leader-based replication model is to allow more than one node to accept writes. Replication still happens in the same way: each node that processes a write must forward that data change to all the other nodes. We call this a multi-leader configuration (also known as master–master or active/active replication). In this setup, each leader simultaneously acts as a follower to the other leaders.
Use Cases for Multi-Leader Replication
It rarely makes sense to use a multi-leader setup within a single datacenter, because the benefits rarely outweigh the added complexity. However, there are some situations in which this configuration is reasonable.
In a multi-leader configuration, you can have a leader in each datacenter. Within each datacenter, regular leader– follower replication is used; between datacenters, each datacenter’s leader replicates its changes to the leaders in other datacenters.
compare how the single-leader and multi-leader configurations fare in a multi- datacenter deployment:
In a single-leader configuration, every write must go over the internet to the datacenter with the leader. This can add significant latency to writes and might contravene the purpose of having multiple datacenters in the first place. In a multi-leader configuration, every write can be processed in the local datacenter and is replicated asynchronously to the other datacenters. Thus, the inter-datacenter network delay is hidden from users, which means the perceived performance may be better.
Tolerance of datacenter outages
In a single-leader configuration, if the datacenter with the leader fails, failover can promote a follower in another datacenter to be leader. In a multi-leader con‐ figuration, each datacenter can continue operating independently of the others, and replication catches up when the failed datacenter comes back online.
Tolerance of network problems
Traffic between datacenters usually goes over the public internet, which may be less reliable than the local network within a datacenter. A single-leader configuration is very sensitive to problems in this inter-datacenter link, because writes are made synchronously over this link. A multi-leader configuration with asynchronous replication can usually tolerate network problems better: a temporary network interruption does not prevent writes being processed.
Although multi-leader replication has advantages, it also has a big downside: the same data may be concurrently modified in two different datacenters, and those write conflicts must be resolved
If you want to guarantee that there will be no editing conflicts, the application must obtain a lock on the document before a user can edit it.
However, for faster collaboration, you may want to make the unit of change very small (e.g., a single keystroke) and avoid locking. This approach allows multiple users to edit simultaneously, but it also brings all the challenges of multi-leader replication, including requiring conflict resolution
Handling Write Conflicts
The biggest problem with multi-leader replication is that write conflicts can occur, which means that conflict resolution is required.
Synchronous versus asynchronous conflict detection
In a single-leader database, the second writer will either block and wait for the first write to complete, or abort the second write transaction, forcing the user to retry the write. On the other hand, in a multi-leader setup, both writes are successful, and the conflict is only detected asynchronously at some later point in time. At that time, it may be too late to ask the user to resolve the conflict.
The simplest strategy for dealing with conflicts is to avoid them: if the application can ensure that all writes for a particular record go through the same leader, then conflicts cannot occur.
you can ensure that requests from a particular user are always routed to the same datacenter and use the leader in that datacenter for reading and writing. Different users may have differ‐ ent “home” datacenters (perhaps picked based on geographic proximity to the user), but from any one user’s point of view the configuration is essentially single-leader.
Multi-Leader Replication Topologies
A replication topology describes the communication paths along which writes are propagated from one node to another. If you have two leaders, there is only one plausible topology: leader 1 must send all of its writes to leader 2, and vice versa. With more than two leaders, various different topologies are possible.
The most general topology is all-to-all, in which every leader sends its writes to every other leader. However, more restricted topologies are also used: for example, MySQL by default supports only a circular topology, in which each node receives writes from one node and forwards those writes (plus any writes of its own) to one other node.
Another popular topology has the shape of a star: one designated root node forwards writes to all of the other nodes. The star topology can be generalized to a tree. In circular and star topologies, a write may need to pass through several nodes before it reaches all replicas. Therefore, nodes need to forward data changes they receive from other nodes.
To prevent infinite replication loops, each node is given a unique identifier, and in the replication log, each write is tagged with the identifiers of all the nodes it has passed through.
A problem with circular and star topologies is that if just one node fails, it can interrupt the flow of replication messages between other nodes, causing them to be unable to communicate until the node is fixed. The topology could be reconfigured to work around the failed node, but in most deployments such reconfiguration would have to be done manually.
The fault tolerance of a more densely connected topology (such as all-to-all) is better because it allows messages to travel along different paths, avoiding a single point of failure.
On the other hand, all-to-all topologies can have issues too. In particular, some net‐ work links may be faster than others (e.g., due to network congestion), with the result that some replication messages may “overtake” others,
However, leader 2 may receive the writes in a different order: it may first receive the update (which, from its point of view, is an update to a row that does not exist in the database) and only later receive the corresponding insert (which should have preceded the update).
the update depends on the prior insert, so we need to make sure that all nodes process the insert first, and then the update. Simply attaching a timestamp to every write is not sufficient, because clocks cannot be trusted to be sufficiently in sync to correctly order these events at leader 2
To order these events correctly, a technique called version vectors can be used. However, conflict detection techniques are poorly implemented in many multi-leader replication systems. For example, at the time of writing, PostgreSQL BDR does not provide causal ordering of writes, and Tungsten Replicator for MySQL doesn’t even try to detect conflicts.
In some leaderless implementations, the client directly sends its writes to several replicas, while in others, a coordinator node does this on behalf of the client. However, unlike a leader database, that coordinator does not enforce a particular ordering of writes. As we shall see, this difference in design has profound consequences for the way the database is used.
Writing to the Database When a Node Is Down
In a leader-based, configuration, if you want to continue processing writes, you may need to perform a failover
On the other hand, in a leaderless configuration, failover does not exist. it’s sufficient for two out of three replicas to acknowledge the write: after user 1234 has received two ok responses, we consider the write to be successful. The client simply ignores the fact that one of the replicas missed the write.
To solve that problem, when a client reads from the database, it doesn’t just send its request to one replica: read requests are also sent to several nodes in parallel.
Read repair and anti-entropy
The replication scheme should ensure that eventually all the data is copied to every replica.
Two mechanisms are often used in Dynamo-style datastores:
When a client makes a read from several nodes in parallel, it can detect any stale responses. For example, in Figure 5-10, user 2345 gets a version 6 value from rep‐ lica 3 and a version 7 value from replicas 1 and 2. The client sees that replica 3 has a stale value and writes the newer value back to that replica. This approach works well for values that are frequently read.
In addition, some datastores have a background process that constantly looks for differences in the data between replicas and copies any missing data from one replica to another. Unlike the replication log in leader-based replication, this anti-entropy process does not copy writes in any particular order, and there may be a significant delay before data is copied.
Limitations of Quorum Consistency
If you have n replicas, and you choose w and r such that w + r > n, you can generally expect every read to return the most recent value written for a key. This is the case because the set of nodes to which you’ve written and the set of nodes from which you’ve read must overlap.
That is, among the nodes you read there must be at least one node with the latest value
However, even with w + r > n, there are likely to be edge cases where stale values are returned. These depend on the implementation, but possible scenarios include:
Thus, although quorums appear to guarantee that a read returns the latest written value, in practice it is not so simple. Dynamo-style databases are generally optimized for use cases that can tolerate eventual consistency.
From an operational perspective, it’s important to monitor whether your databases are returning up-to-date results. Even if your application can tolerate stale reads, you need to be aware of the health of your replication.
If it falls behind significantly, it should alert you so that you can investigate the cause (for example, a problem in the network or an overloaded node).
However, in systems with leaderless replication, there is no fixed order in which writes are applied, which makes monitoring more difficult. Moreover, if the database only uses read repair (no anti-entropy), there is no limit to how old a value might be —if a value is only infrequently read, the value returned by a stale replica may be ancient.
Sloppy Quorums and Hinted Handoff
Databases with appropriately configured quorums can tolerate the failure of individual nodes without the need for failover. They can also tolerate individual nodes going slow, because requests don’t have to wait for all n nodes to respond—they can return when w or r nodes have responded.
In that case, database designers face a trade-off:
• Is it better to return errors to all requests for which we cannot reach a quorum of w or r nodes?
• Or should we accept writes anyway, and write them to some nodes that are reachable but aren’t among the n nodes on which the value usually lives?
By analogy, if you lock yourself out of your house, you may knock on the neighbor’s door and ask whether you may stay on their couch temporarily.
Once the network interruption is fixed, any writes that one node temporarily accepted on behalf of another node are sent to the appropriate “home” nodes.
Once you find the keys to your house again, your neighbor politely asks you to get off their couch and go home.)
Sloppy quorums are particularly useful for increasing write availability: as long as any w nodes are available, the database can accept writes. However, this means that even when w + r > n, you cannot be sure to read the latest value for a key, because the latest value may have been temporarily written to some nodes outside of n.
Thus, a sloppy quorum actually isn’t a quorum at all in the traditional sense. It’s only an assurance of durability, namely that the data is stored on w nodes somewhere. There is no guarantee that a read of r nodes will see it until the hinted handoff has completed.
Leaderless replication is also suitable for multi-datacenter operation, since it is designed to tolerate conflicting concurrent writes, network interruptions, and latency spikes.
Each write from a client is sent to all replicas, regardless of datacen‐ ter, but the client usually only waits for acknowledgment from a quorum of nodes within its local datacenter so that it is unaffected by delays and interruptions on the cross-datacenter link.
Detecting Concurrent Writes
Dynamo-style databases allow several clients to concurrently write to the same key, which means that conflicts will occur even if strict quorums are used. Although in Dynamo-style databases conflicts can also arise during read repair or hinted handoff.
The problem is that events may arrive in a different order at different nodes, due to variable network delays and partial failures.
If each node simply overwrote the value for a key whenever it received a write request from a client, the nodes would become permanently inconsistent.
if you want to avoid losing data, you—the application developer—need to know a lot about the internals of your database’s conflict handling.
Last write wins (discarding concurrent writes)
One approach for achieving eventual convergence is to declare that each replica need only store the most “recent” value and allow “older” values to be overwritten and dis‐ carded. Then, as long as we have some way of unambiguously determining which write is more “recent,” and every write is eventually copied to every replica, the replicas will eventually converge to the same value.
we can attach a timestamp to each write, pick the biggest timestamp as the most “recent,” and discard any writes with an earlier timestamp.
LWW achieves the goal of eventual convergence, but at the cost of durability: if there are several concurrent writes to the same key, even if they were all reported as successful to the client (because they were written to w replicas), only one of the writes will survive and the others will be silently discarded.
If losing data is not acceptable, LWW is a poor choice for conflict resolution.
The “happens-before” relationship and concurrency
An operation A happens before another operation B if B knows about A, or depends on A, or builds upon A in some way. Whether one operation happens before another operation is the key to defining what concurrency means. In fact, we can simply say that two operations are concurrent if neither happens before the other (i.e., neither knows about the other).
Thus, whenever you have two operations A and B, there are three possibilities: either A happened before B, or B happened before A, or A and B are concurrent. What we need is an algorithm to tell us whether two operations are concurrent or not. If one operation happened before another, the later operation should overwrite the earlier operation, but if the operations are concurrent, we have a conflict that needs to be resolved.
Concurrency, Time, and Relativity
It may seem that two operations should be called concurrent if they occur “at the same time”—but in fact, it is not important whether they literally overlap in time.
For defining concurrency, exact time doesn’t matter: we simply call two operations concurrent if they are both unaware of each other, regardless of the physical time at which they occurred.
When a write includes the version number from a prior read, that tells us which previous state the write is based on. If you make a write without including a version number, it is concurrent with all other writes, so it will not overwrite anything—it will just be returned as one of the values on subsequent reads.
Merging concurrently written values This algorithm ensures that no data is silently dropped, but it unfortunately requires that the clients do some extra work: if several operations happen concurrently, clients have to clean up afterward by merging the concurrently written values.
A simple approach is to just pick one of the values based on a version number or timestamp (last write wins), but that implies losing data. So, you may need to do something more intelligent in application code.
As merging siblings in application code is complex and error-prone, there are some efforts to design data structures that can perform this merging automatically
Replication can serve several purposes
Keeping the system running, even when one machine (or several machines, or an entire datacenter) goes down
Allowing an application to continue working when there is a network interrup‐ tion
Placing data geographically close to users, so that users can interact with it faster
Being able to handle a higher volume of reads than a single machine could han‐ dle, by performing reads on replicas
Despite being a simple goal—keeping a copy of the same data on several machines— replication turns out to be a remarkably tricky problem.
It requires carefully thinking about concurrency and about all the things that can go wrong, and dealing with the consequences of those faults. At a minimum, we need to deal with unavailable nodes and network interruptions (and that’s not even considering the more insidious kinds of fault, such as silent data corruption due to software bugs).
Clients send all writes to a single node (the leader), which sends a stream of data change events to the other replicas (followers). Reads can be performed on any replica, but reads from followers might be stale.
Clients send each write to one of several leader nodes, any of which can accept writes. The leaders send streams of data change events to each other and to any follower nodes.
Clients send each write to several nodes, and read from several nodes in parallel in order to detect and correct nodes with stale data.
Each approach has advantages and disadvantages. Single-leader replication is popular because it is fairly easy to understand and there is no conflict resolution to worry about. Multi-leader and leaderless replication can be more robust in the presence of faulty nodes, network interruptions, and latency spikes—at the cost of being harder to reason about and providing only very weak consistency guarantees.
Replication can be synchronous or asynchronous, which has a profound effect on the system behavior when there is a fault. Although asynchronous replication can be fast when the system is running smoothly, it’s important to figure out what happens when replication lag increases and servers fail. If a leader fails and you promote an asynchronously updated follower to be the new leader, recently committed data may be lost.
Users should always see data that they submitted themselves.
After users have seen the data at one point in time, they shouldn’t later see the data from some earlier point in time.
Consistent prefix reads
Users should see the data in a state that makes causal sense: for example, seeing a question and its reply in the correct order.
because they allow multiple writes to happen con‐ currently, conflicts may occur. We examined an algorithm that a database might use to determine whether one operation happened before another, or whether they happened concurrently. We also touched on methods for resolving conflicts by merging together concurrent updates.