The basic problem is very simple: A **Hamiltonian path** is

The basic problem is very simple: A **Hamiltonian path** is a path through a given graph which visits every vertex exactly once. Hamiltonian paths are related^{1} to the better known *Eulerian paths*, which visit every edge exactly once. As an example, this graph has a Hamiltonian Path, which is highlighted in blue:

Related to Hamiltonian paths are **Hamiltonian cycles**, which add the requirement that there must also be an edge between the start and end vertex, or, equivalently they are a *cycle* instead of a *path* that visits every node exactly once. While knowledge of any Hamiltonian cycle trivially allows us to give a Hamiltonian path by simply deleting any edge from it, the opposite does not necessarily hold true. In fact, the graph used in the example above does not even have a Hamiltonian cycle!

To prove that assume that you start with a node in the "left house". At some point, you must take the one connection $\left(E,J\right)$ into the "right house". Now consider two cases:

- You started in $E$ and immediately take the edge $\left(J,E\right)$ back to $E$. The cycle obviously does not visit every node.
- You take any other edge (and end up in one of $\left\lbrace G,H,K \right\rbrace$). Since you started in the "left house", you cannot close the cycle anymore, since you already visited $J$, which ensures that you cannot use the only connection between the "houses".

Since I, personally prefer cycles over paths (they are 7% more chocolaty2), take the following graph, which is simply augmented to allow a Hamiltonian cycle:

Take note how the path through the "right house" has to change so that we end up in $G$, which is connected to the start point of the Hamiltonian path instead of $K$. Although Hamiltonian cycles are tastier than Hamiltonian paths, finding either one is hard in general.

Since the protocol is going to prove the possession of a secret, we first need to generate a secret. To do so, the *prover* starts with a (Hamiltonian) cycle of $n$ nodes, where $n$ is the desired size of the target graph and then hides that cycle by adding a number of random edges and randomly permutating the graph. By construction, the prover knows a Hamiltonian cycle for this graph $G$.

Since the construction allows *any* graph that has a Hamiltonian cycle to be generated, there is nothing which is special in $G$ beyond the fact that it has such a Hamiltonian Cycle. This means that we can publish $G$ as the public key. Similarly to current asymmetric cryptographic systems, an attacker with enough computational power can conceivably break the public key by finding a Hamiltonian path. *This is not prevented by using a ZKP!* It is however made extremely hard by basing the protocol on a problem that is (expected to be) harder than those on which RSA, DSA, etc are built upon3.

The ZKP we are about to craft is going to be a Σ-Protocol, which means it consists of three basic components: A **Commitment** by the prover, followed by a **Challenge** from the verifier which is answered in the **Response** from the prover.

The commitment is going to consist of a version of $G$, which is hidden in such a way that it is possible to either show that the prover knows a Hamiltonian cycle in the hidden graph or to show that the graph is indeed $G$. It is already obvious that the prover must never perform both steps on the same commitment.

To begin, we add "not-edges" to the graph (displayed red in the example below), or more formally we do something akin to adjacency matrices: We store a $1$ for each edge that is an edge in $G$ and $0$ for each edge that is not an edge in $G$. This ensures that the hidden graph will be a clique, making it impossible to garner any information from the degree of individual nodes4.

The prover then generates an individual cryptographic commitment for each edge and non-edge (in the example below, the edges become black). Additionally, a single commitment is generated for all labels5 (in the example below, the labels fade out).

Finally, the graph is randomly permutated. The result now looks exactly the same for *any* graph with the same number of vertices as $G$.

After receiving the Commitment, the verifier poses a simple, binary Challenge to the prover. With $50\%$ chance each, they request that the prover either prove that the Commitment is a commitment to $G$, or that the commitment is for a graph with a Hamiltonian cycle.

To prove that the Commitment is indeed a commitment to $G$, the prover simply opens all cryptographic commitments. Obviously, this cannot give away any information, as this is nothing more than a random permutation of $G$, which is nothing that the verifier could not have computed on their own.

Proving that the Commitment refers to a graph for which the prover knows a Hamiltonian cycle is just as trivial: To do so, the prover **only** opens the commitments to the edges along the Hamiltonian cycle. The only thing that is revealed is an anonymous Graph that consists of exactly one Hamiltonian cycle — something that can trivially be generated by a simulator, meaning that again no information is revealed.

Checking either response is trivial, but obviously may not be omitted nevertheless.

Although this protocol is great for making pretty pictures, it has the drawback that, for a graph $G$ with $n$ nodes, it needs to create $n^2$ commitments to $1$ bit each, plus one to $n\cdot\log_2\left(n\right)$ bits. Even without switching to a different problem or using any advanced techniques, we can do better.

To do so, we begin by generating a random permutation $pi$ to generate a graph $G'$ that is isomorphic to $G$, just as we did in the previous protocol. For each edge $\left(\alpha, \beta\right)$ we generate a commitment $\text{commit}\left(\alpha, \beta\right)$, where $\alpha$ and $\beta$ are indices into $G'$ (and, most importantly, *not* the labels of the original $G$). The Commitment then consists of these commitments in random order and a commitment to $\pi$.

The Challenge need not be changed in any way, and will again either require a proof that the commitments are bound to $G$ or a proof that they are bound to any graph for which the prover knows a Hamiltonian cycle.

The Response works the same way as before as well. To show that the commitments bind to $G$, they are all opened. To show that they bind to a graph for which the prover knows a Hamiltonian cycle, only that cycle is shown. While the first response is trivially zero knowledge, the second requires some deliberation.

A simulator for this algorithm will begin by generating a list of commitments $\text{commit}\left(1,2\right), \text{commit}\left(2,3\right), \ldots, \text{commit}\left(n-1,n\right)$ plus $m-n$ additional commitments to bogus edges (where $m$ is the number of edges in $G$). All these commitments are then randomly permutated, and, together with a commitment to a bogus permutation, form the simulated Commitment.

For a graph $G$ with $n$ nodes and $m$ edges, this scheme improves upon the previous scheme by only requiring $m$ commitments to $2\cdot\log_2\left(n\right)$ bits each. This means that for $m < n^2$ (which is a given) we require less commitments, and that for $m\cdot 2\cdot \log_2\left(n\right) < n^2$, we need to commit to less bits. However, in practice, this advantage grows even further since graphs with more than $2^{64}$ nodes are far too large to be used for authentication and most commitment schemes commit to at least $128 = 2\cdot 2^{64}$ bit at once, which means that the number of commitments is the more relevant measure in practice.

While the previous technique is commonly used in literature treating this topic, the technique of committing to a series of edges seems to be mentioned rather rarely. Its relevance is considered to be marginal by the cryptography community, as Hamiltonicity is not usually used to construct practical ZKPs anyway.

After exploring the idea of Zero Knowledge Proofs (ZKPs) using simple examples in the previous post, we have explored a mathematical example of a ZKP. By basing it on the hard mathematical problems of finding Hamiltonian paths or cycles in arbitrary graphs, the resulting protocol is both zero knowledge and quantum resistant.

A Hamiltonian path/cycle is a Eulerian path/cycle in the line graph. ↩

Alright, in truth it is only $2\cdot\pi \%\approx 6.283185307179586476925\%$. And of course only on average. ↩

Most importantly, no quantum algorithm for the Hamiltonian path and cycle problems are known to date. ↩

To show why this is necessary, consider that exactly two nodes in the example have degree 5 ($C$ and $G$) and that those two are adjacent in the Hamiltonian cycle. We now have a tiny piece of the Hamiltonian cycle: $\left(C, G\right)$. Additionally, the nodes of degree 5 are adjacent to nodes of degree 2 in the Hamiltonian cycle, of which there are only 2 as well ($A$ and $F$). These can only be added in exactly one way to our piece of the Hamiltonian cycle: $\left(A, C, G, F\right)$. Without going further, it should now be clear that this is hardly "zero knowledge"... ↩

This is only really necessary so that the verifier need not solve the graph isomorphy problem. While recent advances in graph isomorphy solving may make this step technically redundant, it simplifies the protocol significantly and, in practice, can also be expected to be faster to do. ↩