DIP: 0008 Title: ChainLocks Author(s): Alexander Block Special-Thanks: Andy Freer, Samuel Westrich, Thephez, Udjinm6 Comments-Summary: No comments yet. Status: Final Type: Standard Created: 2018-11-16 License: MIT License
- Abstract
- Motivation
- Prior Work
- Used LLMQ type
- Safe transactions
- Signing attempts
- Finalization of signed blocks
- Handling of signed blocks
- Conflicting successful signing attempts
- Implications of a signed block
- Network partitions
- Initial Block Download
- Calculations
- Security Considerations
- Copyright
This DIP introduces ChainLocks, a technology for near-instant confirmation of blocks and finding near-instant consensus on the longest valid/accepted chain. ChainLocks leverages LLMQ Signing Requests/Sessions to accomplish this.
When a node encounters multiple valid chains, it sets the local "active" chain by selecting the one that has the most accumulated work. This is generally known as the “longest-chain” rule as in most cases it is equivalent to choosing the chain with the most blocks.
If both chains have the same amount of accumulated work (and in most cases the same block count), a decision can’t be made solely based on the longest-chain rule. In that case, the first chain received by the node is chosen to be the active one and the other chain is put aside. If another block is then received which extends the non-active chain so that it has the most accumulated work, it becomes the active one. For example, even if a chain is currently 6 blocks longer than any other chain, it’s still possible that a shorter chain becomes longer and thus the active one. This is generally known as a chain reorganization.
The most common situation where this happens is if two miners find a block at approximately the same time. Such a block would race in the network and one part of the network would accept one block as the new active chain while another part of the network would accept the other block. In most cases, whoever finds the next block also indirectly resolves the situation as the new block’s parent block determines which of the chains will be the longest one. This is generally known as orphaning of blocks.
It might also happen by accident. For example, if parts of the network with a high hashrate are partitioned and miners are unaware of other miners mining on another chain. When the network becomes healthy again, multiple chains will exist that all branch from a common ancestor. While these chains are propagated, one side of the previously partitioned network will have to reorganize their local chain to the chain of the other side.
It can also happen on purpose if a miner with more hashrate than all other miners combined decides to ignore other miner’s blocks and only mine on top of their own blocks. This is generally known as the 51% mining attack. A miner can even go as far as not publishing any blocks for some time so the remainder of the network is not aware of the attack until they suddenly publish the longer secret chain.
In all these cases, uncertainty arises for individual recipients of funds. When a reorganization happens, it is not necessary for the new chain to include the same transactions as the old chain. In addition to including new transactions and excluding old transactions, it is possible to include transactions in the new chain which are in conflict with the old chain. This means that a new chain might send funds from the same inputs to another address. This results in the only valid form of double spending possible in Dash (InstantSend is not double-spendable even for this case) and most other Bitcoin based cryptocurrencies.
This DIP proposes a new method, called ChainLocks, for reducing uncertainty when receiving funds and removing the possibility of 51% mining attacks.
All signing sessions/requests involved in ChainLocks must use the LLMQ_400_60 LLMQ type.
In general, all valid blocks could also be considered for ChainLocks. There is a small risk of confirmed transactions conflicting with InstantSend locks. To resolve these conflicts, ChainLocks need to be prioritized over InstantSend locks when conflicts arise.
As such conflict resolution is quite severe and in some sense voids all the security gained by InstantSend, it is desirable to reduce the probability of such conflicts to practically zero.
This can be achieved by a change in block template generation. Miners are encouraged to only include “safe” transactions in blocks. A transaction is considered safe if it is InstantSend locked or locally known for at least 10 minutes without a conflicting InstantSend lock appearing in the meantime. The default implementation for block template generation (as found in Dash Core) will be changed to honor this.
At the same time, masternodes should only try to sign/lock blocks which include only "safe" transactions. This means that ChainLocks will only be created when the whole block is considered "safe". Consequently, the probability of a conflicting InstantSend lock appearing after the creation of the ChainLock is practically zero.
This assumes that DIP-0010 LLMQ based InstantSend is already implemented as it changes InstantSend behaviour to try to lock all transactions instead of just a few selected ones.
DIP-0010 also implements "retroactive signing of transactions", which guarantees that transactions are InstantSend locked in a retroactive way if they were unknown before appearing in a mined block. This prevents blocks which contain unlocked transactions from suppressing ChainLocks, as it allows the ChainLocks system to sign blocks retroactively as well.
Each masternode should periodically try to sign the current chain tip. When the tip already has a valid ChainLock, this can be skipped.
Before actually signing the tip, each masternode should check if all transactions contained in all blocks between the last ChainLocked block and the current chain tip are "safe" (see previous section). This check should be limited to a depth of 6 blocks in case there is no known previous ChainLock or if the previous ChainLock is deeper than 6 blocks.
When these checks pass, each masternode must invoke the DIP0007
SignIfMember operation.
The request id for the operation is SHA256("clsig-attempt", blockHeight, attemptNum) and the
message hash is the hash of the new block (newBlockHash). The first time this
is attempted, attemptNum must be set to 0. "clsig-attempt" is a static string that
must be prepended by its length (13, as a compactSize uint).
In most cases, the majority of the LLMQ will sign the same message hash in the
first attempt and thus find consensus. This can be checked with the DIP007
HasRecoveredSig operation.
This will even hold true in most cases where 2 competing blocks are being
propagated inside the network, as only one is able to reach more LLMQ members
faster than the other and thus gain a majority in the signing request.
In some cases however, it is possible that no majority can be reached in the
first attempt. This could happen if too many members of the LLMQ are
malfunctioning or if more than two blocks are competing. If this happens, a
second signing request with an incremented attemptNum value must be
initiated. To check for a failed attempt, the DIP007 IsMajorityPossible
operation must be used. An attempt
is also considered as failed when it did not succeed after some timeout.
On failure, another signing request with an incremented attemptNum value
should be initiated. The new request should use the message hash returned by
the DIP007 GetMostSignedSession operation, which is the hash of the block
which had the most signatures in the last attempt. After a few attempts, a
request should result in a recovered threshold signature which indicates
consensus has been reached.
After a signing attempt has succeeded, another LLMQ must sign the successful
attempt. This is only performed once for each blockHeight and thus either
succeeds or fails without performing additional attempts.
The request id is SHA256("clsig", blockHeight) and the message hash is the block hash of the
previously successful attempt.
After a LLMQ member has successfully recovered the final ChainLocks
signature, it must create a P2P message and propagate it to all nodes. The
message is called CLSIG and has the following structure:
| Field | Type | Size | Description |
|---|---|---|---|
| height | int32 | 4 | Height of the signed block |
| blockHash | uint256 | 32 | Hash of the signed block from the successful attempt |
| sig | BLSSig | 96 | Recovered signature |
This message is propagated through the inventory system.
Upon receipt, each node must perform the following verification before announcing it to other nodes:
- Based on the deterministic masternode list at the given height, a quorum must be selected that was active at the time this block was mined
- The signature must verify against the quorum public key and
SHA256(llmqType, quorumHash, SHA256(height), blockHash).llmqTypeandquorumHashmust be taken from the quorum selected in 2.
When a new block has been successfully signed by a LLMQ and the CLSIG message
is received by a node, it should ensure that only this block is locally
accepted as the next block.
If an alternative block for the same height is received, it must be invalidated and removed from the currently active chain since a signed block has already been received. If the correct block is already present locally, its chain should be activated as the new active chain. If the correct block is not known locally, it must wait for this block to arrive and request it from other nodes if necessary.
If a block has been received locally and no CLSIG message has been received
yet, it should be handled the same way it was handled before the introduction
of ChainLocks. This means the longest-chain and first-seen rules must be
applied. When the CLSIG message for this (or another) block is later
received, the above logic must be applied.
While the network is operating as expected, it’s not possible to encounter two conflicting recovered signatures for two signing attempts of the same parent block. It is possible for a malicious masternode operator to manually double-sign two different attempts when a close race between two competing blocks occurs. If one of the conflicting signature shares is withheld until the second attempt succeeds and the conflicting signature is then propagated to the network, the two attempts will result in two valid recovered signatures.
When performing the finalization of successful attempts, the LLMQ members will
only try to finalize a single attempt, which is usually the first one to
succeed. Only a single attempt will be able to gain a majority during
finalization, which removes the possibility of conflicts. In the worst case,
finalization completely fails, no CLSIG message is created and nodes must
fall back to the first-seen and longest-chain rules.
If a block was successfully signed, it can be safely assumed that no chain reorganization before this block can happen, as all nodes would agree to reject blocks with a lower height. This means that each transaction in this block and all previous blocks can be considered irreversibly and instantly confirmed.
For InstantSend, this also means that the minimum of 6 confirmations of the parent transaction can be removed if the parent transaction is inside or below a signed block.
If there is a network partition, the most likely thing to happen is that just one side is able to mine a signed chain. The other side will encounter non-signed blocks building on top of the last signed block. Miners who observe this must assume that another currently unobserved chain is being built in parallel. Since the parallel chain might be signed and could possibly overtake their own chain after the network is healthy again, miners should act accordingly (e.g. reduce hash power to reduce costs).
If the network is partitioned to a degree that makes a majority in the responsible LLMQ impossible, all partitions in the network will be unable to produce a signed chain. After the network is healthy again, one part of the network will reorganize itself to the other’s chain after which the responsible LLMQ will sign the new chain tip.
While fully synced, nodes will usually receive CLSIG messages for new blocks
shortly after they are mined. If a node was offline for some time or has to
perform an initial block download, the signatures for old blocks will not be
present in the initial implementation.
Nodes should fall back to the plain “longest-chain” and “first-seen” rules in this case until the first block signature for a new block is received.
We assume that old blocks are secure enough to not encounter any significant forks which could lead to a different chain tip after initial block download is finished. When the chain tip is reached, the first received signature will resolve any ambiguities which might occur in the last few blocks.
If the need arises to include block signatures in initial block download, we will update this DIP and implementations accordingly.
We consider the scenario where an attacker has assumed control of a number of masternodes. We will then calculate two probabilities that the attacker will be able to disrupt the network. Specifically, we will calculate the probability that an attacker will be able to prevent all ChainLocks for a twelve hour period, and we will calculate the probability that an attacker can successfully create a malicious or erroneous ChainLock.
The consequences of an attacker withholding ChainLocks would be that the network would fall back on proof of work. The consequence of a malicious or erroneous ChainLock might depend on the attackers motivations. An attacker with control of a ChainLock quorum would be able to:
- Produce conflicting ChainLocks.
- Withhold ChainLocks and then issue a ChainLock for a conflicting previous block not included in the current chain of most work.
- Secure undo advantage.
We assume that all masternodes not under malicious control are behaving as expected. In practice this is generally observed. In theory, it is possible for some nodes to be unreachable, or the attacker could combine a DoS attack in conjuction with compromising nodes. If this is the case, it would become more likely an attacker could withhold ChainLocks, and not be any more likely the attacker could produce a malicious ChainLock. Our calculations can easily be modified to model the network operating under other assumptions.
In this section we will show that the random selection of a LLMQ will not allow an attacker to control a ChainLock quorum without control of a supermajority of masternodes.
Quorum selection is designed to provide for an effectively simple random
choice of masternodes in each quorum. Thus we may assume that each quorum is
a simple random selection. If there are a total of 'N' masternodes the
number of possible quorums is the number of ways that 400 objects can be
chosen out of N. In some literature this number is
written NC400 and is enunciated as
'N choose 400'. Numbers that arise in this fashion are all
binomial coefficients and have a delightful relation to Pascal's Triangle.
For any process with an observable outcome the set of all outcomes is called a sample space. If all outcomes are equally likely, as in the selection of a quorum, the probability of any event is the number of outcomes that satisfy that event over the total number outcomes.
An attacker can temporarily prevent ChainLocks if they control
161 out of 400 masternodes in the quorum. If the attacker controls
m out of the N total masternodes we can compute the
number of outcomes that satisfy the event "exactly 161
nodes of the quorum are controlled by the masternode."
Specifically there are
mC161 * N - mC239
outcomes where the attacker controls exactly 161 masternodes.
Thus, the probability that the attacker controls exactly 161 nodes is
The attacker could also be successful if they control more than 160
of the nodes in the quorum selection. Thus the exact probability that an
attacker can successfully prevent ChainLocks for that quorum round is
computed by counting the number of outcomes where the attacker controls
160 or more nodes in the quorum divided by the total number of outcomes.
This is computed by the sum
The numbers involved in this sum can be quite large. We may use a Python Script to effectively evaluate this sum. Note that we made the necessary modifications to this formula to calculate the probability that an attacker controls 60% or more of the quorum.
| N | m | Success probability of withholding a ChainLock | Success probability of creating a malicious ChainLock |
|---|---|---|---|
| 5000 | 500 | 3.32e-65 | 7.11e-157 |
| 5000 | 1000 | 1.69e-22 | 2.89e-76 |
| 5000 | 1500 | 3.36e-6 | 1.29e-38 |
| 2000 | 200 | 2.12e-87 | 0 |
| 2000 | 400 | 1.80e-26 | 9.49e-94 |
| 2000 | 600 | 6.20e-7 | 3.94e-45 |
In all scenarios above, an attacker should not expect to withhold a ChainLock in the next century. The attacker would have a less than one in 100 trillion chance of producing at least one malicious ChainLock in the next sextillion (10^21) years.
From these calculations we conclude that an attacker with control of fewer than 30% of all masternodes will not have any influence on ChainLocks. For completeness we note that if an attacker controlled 50% of masternodes an attacker would generally be able to prevent ChainLocks, but there is less than 30 out of one billion chance that the attacker will be able to produce at least one malicious ChainLock in the next million years.
It is of note that it is not required for an attacker to actually own the collateral of masternodes that the attacker controls. A company offering the service of hosting masternodes for clients could be persuaded to run malicious software. Also if node operators do not check the digital signature there could be a malicious download or software that spoofs being an official download. We recommend mitigations of both these side channel attacks.
It is recommended that Dash users require some transparency from masternode hosting services. It should be expected that a masternode hosting service disclose all masternodes that it hosts as a service. It is then recommended that those who utilize a masternode hosting service check that their node is listed in the disclosure. This information could help identify if nodes hosted by services are behaving as expected.
It is always recommended that users check the signature on any software they download. As long as the correct public key is acquired then it is cryptographically not possible for the official download to be spoofed.
Copyright (c) 2018 Dash Core Group, Inc. Licensed under the MIT License