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Volume 2021 |Article ID 2507670 |

Chen Shan, Lei Fan, "Lodestone: An Efficient Byzantine Fault-Tolerant Protocol in Consortium Blockchains", Security and Communication Networks, vol. 2021, Article ID 2507670, 10 pages, 2021.

Lodestone: An Efficient Byzantine Fault-Tolerant Protocol in Consortium Blockchains

Academic Editor: Yuling Chen
Received29 Oct 2021
Accepted16 Nov 2021
Published03 Dec 2021


We present Lodestone, a chain-based Byzantine fault-tolerant (BFT) state machine replication (SMR) protocol under partial synchrony. Lodestone enables replicas to achieve consensus with two phases of voting and enjoys (1) optimistic responsiveness and (2) linear communication complexity on average. Similar to the state-of-the-art chain-based BFT protocols, Lodestone can be optimized with a pipelining idea elegantly. We implement pipelined Lodestone and deploy experiments to evaluate its performance. The evaluation results demonstrate that Lodestone has a lower latency than HotStuff under various workloads.

1. Introduction

Consensus in blockchain systems, known as state machine replication (SMR), has attracted more and more interest in recent years. When focusing on permissioned blockchains on so-called consortium blockchains, chain-based Byzantine fault-tolerant (BFT) SMR protocols [18] under partial synchrony have been widely used to achieve consistency. In general, chain-based BFT SMR protocols follow the conventional propose-vote paradigm where there exists a special role often called leader who is responsible for packing clients’ requests into proposals, and then all players achieve consensus on these proposals via multiple (two or three) phases of voting.

PBFT [9], as the first practical BFT SMR protocol under the partial synchronous network [10], achieves safety even under the asynchronous network and liveness when the network gets synchronous. However, the view-change subprotocol in PBFT, with an communication complexity, is too heavy to be practical. Tendermint [6] innovatively employs a lock-commit scheme, similar to the paradigm in [11], that a replica should lock on the proposal he has voted COMMIT for. This allows one replica to decide if voting for one proposal according to his own local states and the leader has no need to prove the safety of his proposal. Casper [7, 8] takes a similar strategy and also implies a pipelining idea for a further improvement. However, both Tendermint and Casper sacrifice optimistic responsiveness in that there needs to be a fixed interval between proposals to guarantee liveness since a new leader has to ensure that he has observed all other nonfaulty replicas’ lock state; otherwise, his new proposal may not be accepted. HotStuff [1] creatively introduces another phase of vote to achieve both linear view-change and optimistic responsiveness. The additional phase guarantees that a new leader can construct this proposal safely only with replicas’ states. However, in pipelined HotStuff, the three-phase voting scheme not only brings about an increase in latency but also causes an implicit liveness problem. In pipelined HotStuff, there needs to be four consecutive nonfaulty leaders to make a decision which cannot be guaranteed in the setting. This means pipelined HotStuff cannot provide liveness in the worst case even under the crash fault-tolerant model, which has also been discussed as the silence attack [12].

We present Lodestone, a novel chain-based BFT SMR protocol, which achieves the following combined properties under partial synchrony:(1)Two-phase voting with optimistic responsiveness: Lodestone can achieve responsiveness when liveness is guaranteed after GST. That is, the total time of confirmation on one honest leader’s proposal only relies on the actual network delay instead of any apriori upper bound assumption of network delay.(2)Linear average-case communication complexity: a view-change subprocess in Lodestone costs message complexity on average and in the worst case. Here, we follow the measurement of message complexity in HotStuff [1] which counts the total number of authenticators received by one player in the protocol to achieve a decision.(3)Deterministic liveness under static corruption: pipelined Lodestone can also achieve deterministic liveness under static corruption where leaders are rotated in a round-robin manner.

The difference in detail between these protocols can be shown in figures, of which Figures 1 and 2 show HotStuff and existing two-phase voting protocols, respectively, while Figure 3 demonstrates our solution.

2.1. BFT SMR Protocols in Alternative Assumptions

There are also many works considering about BFT protocols in alternative assumptions.

Firstly, about the network assumption, many recent protocols under a synchronous network [1315] or under an asynchronous network [1618] make efforts to reduce confirmation latency and achieve practical throughput. Secondly, about the corruption assumption, some recent works [19, 20] are aimed to provide higher assurance on blocks even if the adversary corrupts more than f replicas in the future. In this paper, we only concentrate on static security under a partially synchronous network.

2.2. Single-Shot BFT Protocols

There are some related works [2123] focusing on the single-shot BFT problem which explore the optimal latency bound when the leader is nonfaulty under various resilience assumptions. Though it is still a long way to construct a protocol from single shot to multishots, some results are interesting and may be combined with our protocols in the future work.

2.3. View Synchronization

Another related line of work is about view synchronization [2426] in the partial synchrony setting. It is also necessary in Lodestone for nonfaulty replicas to stay in the same view for a sufficient long time, and then liveness can be guaranteed. However, view synchronization is not the key point of Lodestone, and we assume that replicas are able to stay in the same view for a sufficient long time after GST.

3. Materials and Methods

3.1. Models
(1)Threat model: we consider a permissioned system consisting of n replicas, indexed by , where . We assume a polynomially bounded adversary who can corrupt less than replicas. The replicas corrupted by the adversary can deviate from the prescribed protocol arbitrarily within their capabilities which are also called Byzantine faulty replicas, while the remaining ones are nonfaulty. We only consider a static corruption model in that the adversary chooses which replicas to corrupt prior to the execution.(2)Network model: we assume that each pair of replicas is connected by a reliable authenticated point-to-point channel. Messages are propagated through a partially synchronous network [10] in that there is an unknown global stabilization time (GST). After GST, a message sent by a nonfaulty replica will be delivered to all nonfaulty replicas with a known bound , though the delivery schedule is determined by the adversary.(3)Cryptographic primitives: we assume a cryptographic hash function and a standard digital signature scheme. We also assume a (t, n) threshold signature scheme [27, 28] which provides the following interfaces: generates a pair of key shares for replica i along with a global public key PK. produces a signature share of message m with . verifies if is a valid signature share of message m. produces the threshold signature of message m from t signature shares where and . verifies if is a valid threshold signature of message m with PK.We use an (n − f, n) threshold signature scheme in our following protocols which is assumed to provide robustness and nonforgeability.(4)Problem definition: we now give the definition of a chain-based BFT SMR protocol. Each replica in a chain-based BFT SMR protocol receives requests from clients and maintains a sequence of blocks called a blockchain. Blocks in a blockchain are chained by hash digest, and thus, each block in a blockchain has its own position denoted as its height. Given a blockchain C and a block , all blocks in C lower than b are ancestor blocks of b, and all blocks in C higher than b are descendant blocks of b. Two blocks and are conflicting if and only if is neither an ancestor nor a descendant block of . Each block includes a batch of requests, and one replica's blockchain consists of all blocks which have been finalized. Replicas execute requests in finalized blocks in the sequence of the blockchain and then respond to clients. For simplicity, we do not model clients and assume all requests are sent to all replicas. Thus, we say requests from clients are input to all replicas, and at any time, the output of one replica is his blockchain.

A secure chain-based BFT SMR protocol should satisfy both safety and liveness defined below (in the presence of an arbitrary adversary with all but negligible probability).

Definition 1. (safety). At any time, if two nonfaulty replicas each have their blockchains denoted as and , then it must be either or , where means “is a prefix of or equal to.” In other words, two nonfaulty replicas will never finalize different blocks at the same height.

Definition 2. (liveness). If a request req has been inputted to all replicas at time , then at time , any nonfaulty replica must output a blockchain which includes req, where is the time after which the protocol provides liveness and is a bounded constant.

3.2. Preliminaries
(1)Block: we now format a block b in our protocol as where parent is the hash digest of the parent block of b, view is the view in which b is proposed, txs is a batch of requests from clients, and hash is the hash digest of b (i.e., ). In Lodestone, there are three states for one block, namely, PROPOSED, PREPARED, and COMMITTED. Once a block reaches the COMMITTED state, then the block itself and its all ancestor blocks are finalized and can be executed sequentially.(2)Local states: replicas in Lodestone need to maintain two local states for the protocol execution. The first denoted as currView represents the current view number and implies the current leader. It is noted that using the round-robin manner to rotate leaders guarantees that (i) the protocol will be greeted with a nonfaulty leader after at most f consecutive views; (ii) there exist three consecutive views of which leaders are all nonfaulty which will be proved later. The second is lockedQC which stores qc with the largest view number one replica has voted commit for. Once a nonfaulty replica locks on qc (i.e., sets lockedQC as qc), he will only vote for blocks extending from block b that unless he ensures the majority of nonfaulty replicas have turned to a conflicting branch before locking on qc.(3)Promise and promise-set: during a view-change, each replica additionally sends to the leader a promise, evidence that helps the leader locate the highest COMMITTED block and prove that he does so in an honest manner. A promise for view indicates that the replica has not voted commit at view v. If n − f replicas send promises for the view , then there is no block which has a larger view number than could be COMMITTED. When the leader proposes a block with the n − f promises, all the nonfaulty replicas can accept it safely. A promise from replica P includes a tuple and the corresponding signature share from P. We denote for simplicity. We say p from P for the current view represents that he did not vote commit when he was in the view v. Let be the last view in which P has voted commit. P will generate a set of promises for each view from to currView. We denote it as a promise-set.(4)NullQC: before entering a new view, each nonfaulty replica will send his promise-set promise-set to the new leader, along with his lockedQC. Upon receiving valid promise-sets from n − f distinct replicas, the new leader will select lockedQC with the largest view number as highQC and propose his new block following the block b where . Let I be the set of indexes of the n − f replicas; we have . We use to denote currView and to denote highQC.view. Each promise-set must contain a promise that where . The new leader can combine these signature shares as a threshold signature to prove that highQC he selected is exactly the one with the largest view number among replicas. This extra proof is denoted as nullQC in Lodestone. We have , where .
3.3. Pipelined Lodestone

We are now ready to describe our protocol pipelined Lodestone. We first define some utilities as shown in Figure 4.

Then, we formalize pipelined Lodestone with algorithms. The protocol runs in a succession of views denoted as currView. Each view number is mapped into a leader in a round-robin manner. The leader will execute Algorithm 1, all replicas will execute Algorithm 2, and then the next leader will execute Algorithm 3. When timeout triggers during any wait-for procedure in one replica’s local view, he will execute Algorithm 4. We omit any check for brevity.

(1)Wait for VIEW-CHANGE message of currView-1
(8)Broadcast MSG (PROPOSE, b, lockedQC, NullQC)
(1)Wait for RROPOSE message m from LEADER (currView)
(3)if m.qc.view + 1 = m.proof.view.
(5) send MSG (GENERIC, b, hash,, v) to LEADER (currView + 1).
(6)if m.qc.view>lockedQC.view
(10) DECIDE ()
(1)Wait for matched GENERIC messages of CurrView
Jump here if TIMEOUT triggers during any wait procedure or before entering currView+1
(2)send MSG (VIEW-CHANGE, , lockedQC, T) to LEADER (currView + 1)

4. Safety, Liveness, and Communication Complexity

4.1. Safety

Lemma 1. For any two valid quorum certificates and , if , then .

Proof. For any valid qc, at least nonfaulty replicas have sent their relevant signature shares, namely, at least nonfaulty replicas have voted for the block which qc.block represents in qc.view. Suppose and are two valid qc such that , but . We must have that at least nonfaulty replicas voted for the and also at least nonfaulty replicas voted for the in the same view. Thus, the intersection of the two sets at least includes one nonfaulty replica since , a contradiction to Algorithm 2 that a nonfaulty replica can only vote once in a view.

Theorem 1. In pipelined Lodestone, two conflicting blocks cannot be both COMMITTED.

Proof. Let and be two conflicting blocks which are both COMMITTED, namely, there exists a two-chain and also one in the form of . W.l.o.g, we can assume that with Lemma 1. Since is COMMITTED, at least nonfaulty replicas have voted for and then locked on . Let be the PREPARED block with the smallest view number that satisfies , and conflicts with . Such must exist since satisfies all these conditions. Now, consider the proposal m of . At least nonfaulty replicas have locked on in . These nonfaulty replicas will only promise on view number larger than since then. Therefore, . And due to the minimality of , we must have . Then, , a contradiction to the validity of nullQC.

4.2. Liveness

Lemma 2. Under an setting where leaders are rotated in a round-robin manner, in any consecutive views, there are three consecutive views all leaded by nonfaulty replicas.

Proof. Consider any n consecutive views denoted as . There are three cases:(i)The leader of is faulty. Then, for the view from + 1 to +n − 1, there are f − 1 faulty leaders while nonfaulty leaders. The faulty leaders can at most separate nonfaulty leaders into segments, with at least one segment owning 3 nonfaulty leaders.(ii)The leader of is nonfaulty, while the leader of + 1 is faulty. Then, for the view from + 2 to +n, there are faulty leaders while nonfaulty leaders. Then, the situation becomes the same as the first case.(iii)Both and + 1 are leaded by nonfaulty leaders. Then, if the leader of + 2 is nonfaulty, three consecutive nonfaulty leaders exist. Otherwise, for the view from + 3 to +n + 1, there are faulty leaders while nonfaulty leaders. Then, the situation becomes the same as the first case.

Lemma 3. If a nonfaulty leader of view v collects n − f valid view-change messages, he can always propose his new block, along with valid nullQC that where highQC points to the parent block of his new block.

Proof. Let be the nonfaulty leader of view . If collects valid view-change messages, he will select qc with the largest view number among them as highQC. Each of the valid view-change messages must also include a promise on ; then, can combine a threshold signature with these signature shares and construct nullQC that .

Theorem 2. In pipelined Lodestone, after GST, any request from clients will be included in the finalized block in a bounded time.

Proof. According to Lemma 2, there are three consecutive views leaded by nonfaulty replicas in any consecutive views, denoted as . There exists a bounded time duration T that if all nonfaulty replicas stay in the same view during T, then the leader can propose his new block with according to Lemma 3. Thus, the first nonfaulty leader of v will propose a block b, and all nonfaulty replicas will vote for b. Then, the second nonfaulty leader of + 1 will collect n − f votes for b and propose a block b' with . All nonfaulty replicas will also vote for b'. Then, the third nonfaulty leader of + 2 will collect votes for b' and propose a block b'' with . Therefore, b, b', and b'' form a two-chain , and all nonfaulty replicas will consider b as COMMITTED after receiving b''. For any time t, the three consecutive views leaded by nonfaulty players will come within a bounded time duration after t. Therefore, there exists a bounded constant in that for any request input to all replicas at time t, it will be included in the finalized block at .

4.3. Communication Complexity

We now discuss the communication complexity of pipelined Lodestone. It is noted that we only consider the communication overhead when liveness can be guaranteed after GST. When the network is asynchronous, there may be unbounded views without making a decision. In fact, we can allow replicas at most send promises during a view-change which has no effect on the safety property since it is impossible to guarantee liveness under an asynchronous network [29].

Theorem 3. After GST, pipelined Lodestone achieves a linear communication complexity on average while in the worst case.

Proof. When liveness can be guaranteed after GST, whenever two consecutive views are leaded by nonfaulty replicas, the second leader can generate qc, and all nonfaulty replicas will update their lockedQC. Therefore, in the worst case, there are consecutive views without generating new qc. And then, one replica needs to send promises on these views. And thus, the total complexity of a view-change is in the worst case.
Now, we discuss the communication complexity of a view-change in the average case. The probability of one view leaded by a faulty replica can be considered as while by a nonfaulty leader approximately with the assumption . Given any view , let X be the length of views before without any two consecutive views leaded by nonfaulty replicas. Let P(X) be the probability distribution of X. For any view , X = k occurs when(i)Either the leader of v − 1 is faulty and X = k − 1 holds for v − 1(ii)Or the leader of v − 1 is nonfaulty while the leader of v − 2 is faulty and X = k − 2 holds for v − 2As the assumption, v − 1 and v − 2 also follow the same probability distribution which is denoted as P(X = k − 1) and P(X = k − 2), respectively. Then, P(X = k) can be expressed asThe expectation E(X) is the length of views before without any two consecutive views leaded by nonfaulty players on average. E(k) is also the total number of signatures in the promise-set of one replica on average. We have after simplification. Therefore, the expectation of the total communication overhead of a view-change is for all n replicas in a view.

5. Results and Discussion

5.1. Implementation and Setup

We have implemented both pipelined Lodestone and pipelined HotStuff in C++ language with the same codebase for a fair comparison, taking the implementation ( in HotStuff’s paper [1] as a reference. We use Ed25519 for common digital signatures and BLS threshold signatures ( for combining signatures in our protocol.

We deploy our experiments on Alibaba Cloud using ecs.ic5.4xlarge instances. The round-trip delay between two instances is less than 1 millisecond, with the bandwidth about 5 Gbps.

In all experiments, besides all players and clients, we develop a coordinator who is responsible for notifying all clients sending requests to players. The coordinator collects measurement data to compute the throughput and end-to-end latency of clients. Figure 5 shows the architecture in our experiments.

We first evaluate these two protocols with zero-sized payload and different choices of batch sizes to get rid of the effects of payload size. Figure 6 shows that Lodestone has a prominent lower latency compared with HotStuff under the batch size of both 400 and 800 benefiting from conserving one phase of the vote.

Figure 7 depicts different payload sizes for both systems as 512 bytes and 1024 bytes, with a fixed batch size of 800. In such settings, Lodestone still enjoys a remarkable lower latency compared with HotStuff and provides comparable throughput.

6. Conclusions

We presented pipelined Lodestone, a chain-based BFT SMR protocol, which achieves linear view-change on average and optimistic responsiveness with only two phases of voting. Through the experimental results, pipelined Lodestone provides a lower latency and comparable throughput compared with HotStuff in various workloads and network scales.

Data Availability

The data used to support the findings of this study are available from the corresponding author upon request.

Conflicts of Interest

The authors declare that there are no conflicts of interest regarding the publication of this paper.


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Copyright © 2021 Chen Shan and Lei Fan. This is an open access article distributed under the Creative Commons Attribution License, which permits unrestricted use, distribution, and reproduction in any medium, provided the original work is properly cited.

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