Abstract
Deviceindependent cryptography goes beyond conventional quantum cryptography by providing security that holds independently of the quality of the underlying physical devices. Deviceindependent protocols are based on the quantum phenomena of nonlocality and the violation of Bell inequalities. This high level of security could so far only be established under conditions which are not achievable experimentally. Here we present a property of entropy, termed “entropy accumulation”, which asserts that the total amount of entropy of a large system is the sum of its parts. We use this property to prove the security of cryptographic protocols, including deviceindependent quantum key distribution, while achieving essentially optimal parameters. Recent experimental progress, which enabled loopholefree Bell tests, suggests that the achieved parameters are technologically accessible. Our work hence provides the theoretical groundwork for experimental demonstrations of deviceindependent cryptography.
Introduction
Deviceindependent (DI) quantum cryptographic protocols achieve an unprecedented level of security—with guarantees that hold (almost) irrespective of the quality, or trustworthiness, of the physical devices used to implement them^{1}. The most challenging cryptographic task in which DI security has been considered is quantum key distribution (QKD); we will use this task as an example throughout the manuscript. In DIQKD, the goal of the honest parties, called Alice and Bob, is to create a shared key, unknown to everybody else but them. To execute the protocol, they hold a device consisting of two parts: each part belongs to one of the parties and is kept in their laboratories. Ideally, the device performs measurements on some entangled quantum states it contains.
In real life, the manufacturer of the device, called Eve, can have limited technological abilities (and hence cannot guarantee that the device’s actions are exact and nonfaulty) or even be malicious. The device itself is far too complex for Alice and Bob to open and assess whether it works as Eve alleges. Alice and Bob must therefore treat the device as a black box with which they can only interact according to the protocol. The protocol must allow them to test the possibly faulty or malicious device and decide whether using it to create their keys poses any security risk. The protocol guarantees that by interacting with the device according to the specified steps, the honest parties will either abort, if they detect a fault, or produce identical and secret keys (with high probability).
Adopting the DI approach is not only crucial for the paranoid cryptographers; even the most skilled experimentalist will recognise that a fully characterised, stable at all times, largescale quantum device that implements a QKD protocol is extremely hard to build. Indeed, implementations of QKD protocols have been attacked by exploiting imperfections of the devices^{2,3,4,5}. Instead of trying to come up with a “patch” each time an imperfection in the device is detected, DI protocols allow us to break the cycle of attacks and countermeasures.
The most important (in fact necessary) ingredient, which forms the basis of all DI protocols, is a “test for quantumness” based on the violation of a Bell inequality^{6,7,8,9}. A Bell inequality^{10,11} can be thought of as a game played by the honest parties using the device they share (Fig. 1). Different devices lead to different winning probabilities when playing the game. The game has a special “feature”—there exists a quantum device which achieves a winning probability ω_{q} greater than all classical, local, devices. Hence, if the honest parties observe that their device wins the game with probability ω_{q} they conclude that it must be nonlocal^{11}. A recent sequence of breakthrough experiments have verified the quantum advantage in such “Bell games” in a loopholefree way^{12,13,14} (in particular, this means that the experiments were executed without making assumptions that could otherwise be exploited by Eve to compromise the security of a cryptographic protocol).
DI security relies on the following deep but wellestablished facts. High winning probability in a Bell game not only implies that the measured system is nonlocal, but more importantly that the kind of nonlocal correlations it exhibits cannot be shared: the higher the winning probability, the less information any eavesdropper can have about the devices’ outcomes. The tradeoff between winning probability and secret randomness, or entropy, can be made quantitative^{15,16}.
The amount of entropy, or secrecy, generated in a single round of the protocol can therefore be calculated from the winning probability in a single game. The major challenge, however, consists in establishing that entropy accumulates additively throughout the multiple rounds of the protocol and use it to bound the total secret randomness produced by the device.
A commonly used assumption^{17,18,19,20,21} to simplify this task is that the device held by the honest parties makes the same measurements on identical and independent quantum states in every round i ∈ {1, …, n} of the protocol. This implies that the device is initialised in some (unknown) state of the form σ^{⊗n}, i.e., an independent and identically distributed (i.i.d.) state, and that the measurements have a similar structure. In that case, the total entropy created during the protocol can be easily related to the sum of the entropies generated in each round separately (as further explained below).
Unfortunately, although quite convenient for the analysis, the i.i.d. assumption cannot be justified a priori. When considering devicedependent protocols, such as the BB84 protocol^{22}, de Finetti theorems^{23,24} can often be applied to reduce the task of proving the security in the most general case to that of proving security with the i.i.d. assumption. This approach was unsuccessful in the DI scenario, where known de Finetti theorems^{23,24,25,26} do not apply. Hence, one cannot simply reduce a general security statement to the one proven under the i.i.d. assumption.
Without this assumption, however, very little is known about the structure of the untrusted device and hence also about its output. As a consequence, previous DIQKD security proofs had to address directly the most general case^{27,28,29}. This led to security statements which are of limited relevance for practical experimental implementations; they are applicable only in an unrealistic regime of parameters, e.g., small amount of tolerable noise and large number of signals.
The work presented here resolves this situation. First, we provide a general informationtheoretic tool that quantifies the amount of entropy accumulated during sequential processes which do not necessarily behave identically and independently in each step. We call this result the “Entropy Accumulation Theorem” (EAT). We then show how it can be applied to essentially reduce the problem of proving DI security in the most general case to that of the i.i.d. case. This allows us to establish simple and modular security proofs for DIQKD that yield tight key rates. Our quantitative results imply that the first proofs of principle experiments implementing a DIQKD protocol are within reach with today’s stateoftheart technology. Aside from its application to security proofs, the EAT can be used in other scenarios in quantum information such as the analysis of quantum random access codes.
Results
In the following, we start by explaining the main steps in a security proof of DIQKD under the i.i.d. assumption using wellestablished techniques. We then present the EAT and show how it can be used to extend the proof and achieve full security (i.e., without assuming an i.i.d. behaviour of the device).
Security under the independent and identically distributed device assumption
The central task when proving the security of cryptographic protocols consists in bounding the information that an adversary, called Eve, may obtain about certain values generated by the protocol, which are supposed to be secret. For QKD, the appropriate measure of knowledge, or rather uncertainty, is given by the smooth conditional minentropy^{30} \(H_{{\mathrm{min}}}^\varepsilon \left( {\left. K \rightE} \right)\), where K is the raw data obtained by the honest parties, E the quantum system held by Eve, and ε a parameter describing the security of the protocol. The quantity \(H_{{\mathrm{min}}}^\varepsilon \left( {\left. K \rightE} \right)\) determines the maximal length of the secret key that can be created by the protocol. Hence, proving the security amounts to establishing a lower bound on \(H_{{\mathrm{min}}}^\varepsilon \left( {\left. K \rightE} \right)\). Evaluating \(H_{{\mathrm{min}}}^\varepsilon \left( {\left. K \rightE} \right)\) can be a daunting task, as the adversary’s system E is out of our control; in particular, it can have arbitrary dimension and share quantum correlations with the users’ devices.
Most protocols consist of a basic building block, or “round”, which is repeated a large number, n, of times; in each round i, the classical data K_{ i } is generated. The structure of a DIQKD protocol is shown in Box 1. The i.i.d. assumption means that the raw key \(K_1^n = K_1, \ldots ,K_n\) can be treated as a sequence of i.i.d. random variables K_{ i }. That is, all the K_{ i } are identical and independent of one another. The eavesdropper has side information E_{ i } about each K_{ i }. In this case, the total conditional minentropy \(H_{{\mathrm{min}}}^\varepsilon \left( {\left. {K_1^n} \rightE_1^n} \right)\) can be directly related to the singleround conditional von Neumann entropy \(H\left( {\left. {K_i} \rightE_i} \right)\) using the quantum asymptotic equipartition property^{31} (AEP), which asserts that
where c_{ ε } depends only on ε (see the Methods section).
To get a bound on \(H_{{\mathrm{min}}}^\varepsilon \left( {K_1^n\left {E_1^n} \right.} \right)\), we therefore need to analyse the secrecy, \(H\left( {K_i\left {E_i} \right.} \right)\), resulting from a single round of the protocol. Depending on the considered scenario, a lower bound on \(H\left( {K_i\left {E_i} \right.} \right)\) can be found using different techniques. For discrete and continuousvariable QKD, for example, one can use the entropic uncertainty relations^{32,33}. When dealing with DIQKD, a quantum advantage in a Bell game implies a lowerbound on \(H\left( {K_i\left {E_i} \right.} \right)\) as discussed above.
The Clauser–Horne–Shimony–Holt (CHSH) game^{34} (presented in Fig. 1) forms the basis for most DIQKD protocols. For this game, a tight bound on the secrecy as a function of the winning probability in the game was derived^{19}. The bound implies that for any quantum state that wins the CHSH game with probability ω, the entropy evaluated on the state of the system after the game has been played is at least
where h(⋅) is the binary entropy function. This relation is shown in Fig. 2.
To compute the bound on \(H\left( {K_i\left {E_i} \right.} \right)\), Alice and Bob need to collect the statistics they observe while running the protocol and estimate the winning probability ω appearing in Eq. (2); assuming the i.i.d. structure this is easily done using Hoeffding’s inequality.
The conclusion of this section is the following. The i.i.d. assumption plays a crucial role in the above line of proof: it allows us to reduce the problem of calculating the total secrecy of the raw key created by the device to that of bounding the secrecy produced in one round. Instead of dealing with largescale quantum systems, we are only required to understand the physics of small systems associated with just one round (as in Eq. (2)). The AEP appearing as Eq. (1) does the rest.
Extending to full security
Assuming the device behaves in an i.i.d. way goes completely against the DI setting by imposing a severe and even unrealistic restrictions on the implementation of the device. In particular, the assumption implies that the device does not include any, classical or quantum, internal memory (i.e., its actions in one round cannot depend on the previous rounds) and cannot display timedependent behaviour.
Our main contribution can be phrased as follows.
Theorem (Security of DIQKD, informal): Security of DIQKD in the most general case follows from security under the i.i.d. assumption. Moreover, the dependence of the key rate on the number of exchanged signals, n, is the same as the one in the i.i.d. case, up to terms that scale like \(1{\mathrm{/}}\sqrt n\). The key rates are plotted below.
We now explain the above theorem and how it is derived in more detail. A general device is described by an (unknown) tripartite state \(\rho _{Q_{\mathrm{A}}Q_{\mathrm{B}}E}\), where the bipartite quantum state \(\rho _{Q_{\mathrm{A}}Q_{\mathrm{B}}}\) is shared between Alice and Bob and ρ_{E} belongs to Eve, together with the measurements applied to \(\rho _{Q_{\mathrm{A}}Q_{\mathrm{B}}}\) when the device is used. No additional structure is assumed (see Fig. 3).
As mentioned above, the standard DIQKD protocol proceeds in rounds (recall Box 1): Alice and Bob use their components in the first round of the protocol and only then proceed to the second round, etc. We leverage this structure to bound the amount of entropy produced during a complete execution of the protocol.
To do so, we prove a generalisation of the AEP given in Eq. (1) to a scenario in which, instead of the raw key \(K_1^n = K_1, \ldots ,K_n\) being produced by an i.i.d. process, its parts K_{ i } are produced one after the other. In this case, each K_{ i } may depend not only on ith round of the protocol but also on everything that happened in previous rounds (but not on the subsequent ones). We explain our tool, the EAT, in the following.
The entropy accumulation theorem (EAT)
We describe here a simplified and informal version of the EAT, sufficient to understand how it can be used to prove the security of DI protocols; for the most general statements see the Methods section.
We consider processes which can be described by a sequence of n maps \({\cal M}_1, \ldots ,{\cal M}_n\), called “EAT channels”, as shown in Fig. 4. Each \({\cal M}_i\) outputs two systems, O_{ i }, which describes the information that should be kept secret, and S_{ i }, describing some side information leaked by the map, together with a “memory” system R_{ i }, which is passed on as an input to the next map \({\cal M}_{i + 1}\). The systems \(S_1^n\) describe the side information created during the process. A further quantum system, denoted by E, represents additional side information correlated to the initial state in the beginning of the considered process. The systems \(O_1^n\) are then the ones in which entropy is accumulated conditioned on the side information \(S_1^n\) and E.
To bound the entropy of \(O_1^n\), we take into account global statistical properties. These are inferred by tests carried out by the protocol on a small sample of the generated outputs. To incorporate such statistical information, we consider for each round an additional classical value C_{ i } computed from O_{ i } and S_{ i }. Additionally, in each step of the process, the previous outcomes \(O_1^{i  1}\) are independent of future information S_{ i } given all the past side information \(S_1^{i  1}E\). By choosing O_{ i } and S_{ i } properly, this condition can be satisfied by sequential protocols such as DIQKD.
The EAT relates the total amount of entropy to the entropy accumulated in one step of the process. The latter is quantified by the minimal, or worstcase, von Neumann entropy produced by the maps \({\cal M}_i\) when acting on an input state that reproduces the correct statistics on C_{ i }, i.e., states that satisfy \({\cal M}_i(\sigma )_{C_i} = {\rm freq}(c_1^n)\) where \({\rm freq}(c_1^n)\) is the empirical statistics, or frequency distribution, on \({\cal C}\) defined by \({\rm freq}(c_1^n)(c) = \frac{{\left {\{ i \in \{ 1, \ldots ,n\} :c_i = c\} } \right}}{n}\).
To state the explicit result, we define a “mintradeoff function”, f_{min}, from the set of probability distributions over \({\cal C}\) to the real numbers; f_{min} should be chosen as a convex differentiable function which is bounded above by the worstcase entropy just described:
An event Ω is defined by a subset of \({\cal C}^n\) and we write \(p_{{\Omega }} = \mathop {\sum}\nolimits_{c_1^n \in {{\Omega }}} ( {\rho _{O_1^nS_1^nE,C_1^n = c_1^n}} )\) for the probability of the event Ω and
for the state conditioned on Ω. We further define a set \({\hat{ \Omega }}\) over the set of frequencies such that for all \(c_1^n\), \({\rm freq}(c_1^n) \in {\hat{ \Omega }}\) if and only if \(c_1^n \in { {\Omega }}\).
Theorem (EAT, informal): For any EAT channels, an event Ω such that \({\hat{ \Omega }}\) is a convex set, and a convex mintradeoff function for which \(f_{{\mathrm{min}}}\left( {\rm freq}({c_1^n}) \right) \ge t\) for any \({\rm freq}(c_1^n) \in {\hat{ \Omega }}\),
where the conditional smooth minentropy is evaluated on ρ_{Ω} and v depends on the values \(\left\ {\nabla f_{{\mathrm{min}}}} \right\_\infty ,\varepsilon ,p_{{\Omega }}\), and the maximal dimension of the systems O_{ i }.
Equation (5) asserts that, to first order in n, the total conditional smooth minentropy is at least n times the value of the mintradeoff function, evaluated on the empirical statistics observed during the protocol (and hence linear in the number of rounds). In the special case where the EAT channels are independent and identical, the EAT is reduced to the quantum AEP; Eq. (5) is thus a generalisation of Eq. (1).
DIQKD security via the EAT
To gain intuition on how the EAT can be applied to DIQKD, note the following. Define the maps \({\cal M}_i\) to describe the joint behaviour of the honest parties and their respective uncharacterised device while playing a single round of a Bell game such as the CHSH game. Let Ω be the event of the protocol not aborting or a closely related event, e.g., the event that the fraction of CHSH games won is above some threshold ω_{T}. The state for which the smooth minentropy is evaluated is ρ_{Ω}, i.e., the state at the end of the protocol conditioned on not aborting. This implies, in particular, a bound on \(H_{{\mathrm{min}}}^\varepsilon \left( {K_1^n\left E \right.} \right)\).
Furthermore, the condition on the mintradeoff function stated in Eq. (3) corresponds to the requirement that the distribution of C_{ i } equals \(c_1^n\), which ensures that the entropy in Eq. (3) is evaluated on states that can be used to win the CHSH game with probability ω_{T}. Thus, in order to devise an appropriate mintradeoff function, we can use the relation appearing in Eq. (2); the exact details are given in the Methods section. This results in a tight bound on the amount of entropy created in each step of the protocol. In this sense, we reduce the problem of proving the security of the whole protocol to that of a single round.
Using the EAT we get a bound on \(H_{{\mathrm{min}}}^\varepsilon \left( {K_1^n\left E \right.} \right)\) which, to first order in n, coincides with the one derived under the i.i.d. assumption and is thus optimal. The final key rate \(r = \ell {\mathrm{/}}n\) (where \(\ell\) is the length of a key) produced in a DIQKD protocol depends on this amount of entropy and the amount of information leaked during standard classical postprocessing steps. We plot the results for specific choices of parameters in Fig. 5.
To calculate the key rate, one must have some honest implementation of the protocol in mind; this is given by what the experimentalists think (or guess) is happening in their experiment when an adversary is not present. It does not, in any way, restrict the actions of the adversary or the types of imperfections in the device. We consider the following honest implementation, but the analysis can be adapted to any other implementation of interest.
In the realisation of the device, in each round, Alice and Bob share the twoqubit Werner state \(\rho _{Q_{\mathrm{A}}Q_{\mathrm{B}}} = (1  \nu )\left {{\mathrm{\Phi }}^ + } \right\rangle \left\langle {{\mathrm{\Phi }}^ + } \right + \nu {\Bbb I}{\mathrm{/}}4\) resulting from a depolarisation noise acting on the maximally entangled state \(\left {{\mathrm{\Phi }}^ + } \right\rangle\). In every round, the measurements for X_{ i }, Y_{ i } ∈ {0, 1} are as described in Fig. 1 and for Y_{ i } = 2 Bob’s measurement is σ_{ z }. The winning probability in the CHSH game (restricted to X_{ i }, Y_{ i } ∈ {0, 1}) using these measurements on \(\rho _{Q_{\mathrm{A}}Q_{\mathrm{B}}}\) is \(\omega _{{\mathrm{exp}}} = \left[ {2 + \sqrt 2 (1  \nu )} \right]{\mathrm{/}}4\). The quantum bit error rate \(Q = {\mathrm{Pr}}\left[ {\left. {A_i \ne B_i} \right\left( {X_i,Y_i} \right) = (0,2)} \right]\) for the above state and measurements is given by Q = ν/2.
The key rate r is plotted in Fig. 5. For n = 10^{15}, the curve essentially coincides with the rate achieved in the asymptotic i.i.d. case^{19}. Since the latter was shown to be optimal^{19}, it provides an upper bound on the key rate and the amount of tolerable noise. Hence, for large enough n our rates become optimal and the protocol can tolerate up to the maximal error rate Q = 7.1%. For comparison, the previously established explicit rates^{28} are well below the lowest curve presented in Fig. 5, even when the number of signals goes to infinity, with a maximal noise tolerance of 1.6%. Moreover, our key rates are comparable to those achieved in devicedependent QKD protocols^{35}.
Discussion
The information theoretic tool, the EAT, reveals a novel property of entropy: the operationally relevant total uncertainty about an npartite system created in a sequential process corresponds to the sum of the entropies of its parts, even without an independence assumption.
Using the EAT, we show that practical and realistic protocols can be used to achieve the unprecedented level of DI security. The next major challenge in experimental implementations is a field demonstration of a DIQKD protocol. This would provide the strongest cryptographic experiment ever realised. The work presented here provides the theoretical groundwork for such experiments. Our quantitative results imply that the first proofs of principle experiments, with small distances and small rates, are within reach with today’s stateoftheart technology, which recently enabled the violation of Bell inequalities in a loopholefree way.
Methods
We state here the main theorems of our work and sketch the proofs. Using the explicit expressions given below, one can reproduce the key rates presented in Fig. 5.
The formal statement and proof idea of the EAT
In this section, we are interested in the general question of whether entropy accumulates, in the sense that the operationally relevant total uncertainty about an npartite system \(O_1^n\) corresponds to the sum of the entropies of its parts O_{ i }. The AEP, given in Eq. (1), implies that this is indeed the case to first order in n—under the assumption that the parts O_{ i } are identical and independent of each other. Our result shows that entropy accumulation occurs for more general processes, i.e., without an independence assumption, provided one quantifies the uncertainty about the individual systems O_{ i } by the von Neumann entropy of a suitably chosen state.
The type of processes that we consider are those that can be described by a sequence of channels, as illustrated in Fig. 4. Such channels are called EAT channels and are formally defined as follows.
Definition 1 (EAT channels): EAT channels \({\cal M}_i:R_{i  1} \to R_iO_iS_iC_i\), for i ∈ [n], are CPTP (completely positive trace preserving) maps such that for all i ∈ [n]:

1.
C_{ i } are finitedimensional classical systems (random variables). O_{ i }, S_{ i }, and R_{ i } are quantum registers. The dimension of O_{ i } is \(d_{O_i}.\)

2.
For any input state \(\sigma _{R_{i  1}R{\prime}}\), where R′ is a register isomorphic to R_{i−1}, the output state \(\sigma _{R_iO_iS_iC_iR{\prime}} = \left( {{\cal M}_i \otimes {\Bbb I}_{R{\prime}}} \right)\left( {\sigma _{R_{i  1}R{\prime}}} \right)\) has the property that the classical value C_{ i } can be measured from the system \(\sigma _{O_iS_i}\) without changing the state.

3.
For any initial state \(\rho _{R_0E}^0\), the final state \(\rho _{O_1^nS_1^nC_1^nE} = \left( {\left( {{\mathrm{Tr}}_{R_n} \circ {\cal M}_n \circ \ldots \circ {\cal M}_1} \right) \otimes {\Bbb I}_E} \right)\rho _{R_0E}^0\) fulfils the Markov chain condition \(O_1^{i  1} \leftrightarrow S_1^{i  1}E \leftrightarrow S_i\) for each i ∈ [n].
In the above definition, \(O_1^{i  1} \leftrightarrow S_1^{i  1}E \leftrightarrow S_i\) if and only if their conditional mutual information is 0, \(I\left( {O_1^{i  1}:S_i\left {S_1^{i  1}} \right.E} \right) = 0\).
Next, one should find an adequate way to quantify the amount of entropy which is accumulated in a single step of the process, i.e., in an application of just one channel. To do so, let p be a probability distribution over \({\cal C}\), where \({\cal C}\) denotes the common alphabet of \(C_1, \ldots ,C_n\), and R′ be a system isomorphic to R_{i−1}. We define the set of states
where \(\sigma _{C_i}\) denotes the probability distribution over \({\cal C}\) with the probabilities given by \(\left\langle c \right\sigma _{C_i}\left c \right\rangle\).
The tradeoff functions for the EAT channels are defined below.
Definition 2 (Tradeoff functions): A real function is called a min or maxtradeoff function for \({\cal M}_i\) if it satisfies
or
respectively, and if it is convex or concave, respectively. If the set Σ_{ i }(p) is empty, then the infimum and supremum are by definition equal to ∞ and −∞, respectively, so that the conditions are trivial.
To get some intuition as to why the above definition is the “correct” one, consider the following classical example. Each EAT channel outputs a single bit O_{ i } without any side information S_{ i } about it; the system E is empty as well. Every bit can depend on the ones produced previously. We would like to extract randomness out of the sequence \(O_1^n\); for this we should find a lower bound on \(H_{{\mathrm{min}}}^\varepsilon \left( {O_1^n} \right)\).
We ask the following question—given the randomness of O_{1} which is already accounted for, how much randomness does O_{2} contribute? One possible guess is the conditional von Neumann entropy \(H\left( {\left. {O_2} \rightO_1} \right) = {\Bbb E}_{o_1,o_2}\left[ {  {\mathrm{log}}\,{\mathrm{Pr}}\left( {\left. {o_2} \righto_1} \right)} \right]_{}^{}\). If, however, O_{1} is uniform while O_{2} is fixed when O_{1} = 0 and uniform otherwise, then \(H\left( {\left. {O_2} \rightO_1} \right)\) is too optimistic; the amount of extractable randomness is quantified by the smooth minentropy, which depends on the most probable value of O_{1}O_{2}, and not by an average quantity as the von Neumann entropy.
Another possible guess is a worstcase version of the minentropy \(H_{{\mathrm{min}}}^{{\rm w.c.}} = {\mathrm{min}}_{o_1,o_2}\left[ {  {\mathrm{log}}\,{\mathrm{Pr}}\left( {\left. {o_2} \righto_1} \right)} \right]\). This, however, is too pessimistic; when the O_{ i }’s are independent of each other, the extractable amount of randomness behaves like the von Neumann entropy in first order, and not like the minentropy.
We therefore choose an intermediate quantity—\({\mathrm{min}}_{o_1}\,{\Bbb E}_{o_2}\left[ {  {\mathrm{log}}\,{\mathrm{Pr}}\left( {\left. {o_2} \righto_1} \right)} \right] = {\mathrm{min}}_{o_1}\,H\left( {\left. {O_2} \rightO_1 = o_1} \right)\). That is, this quantity is the von Neumann entropy of O_{2}, evaluated for the worstcase state in the beginning of the second step of the process. The mintradeoff function defined above is the quantum analogue version of this.
Informally, the mintradeoff function can be understood as the amount of entropy available from a single round, conditioned on the outputs of the previous rounds. Since we condition on the previous rounds, one can think of the randomness of the current round as independent from past events. Intuitively, this suggests that, by appropriately generalising the proof of the AEP, one can argue that the entropy that is contributed by this independent randomness in each round accumulates.
The formal statement of the EAT is as follows.
Theorem 3 (EAT, formal): Let \({\cal M}_i:R_{i  1} \to R_iO_iS_iC_i\) for i ∈ [n] be EAT channels, ρ be the final state, Ω an event defined over \({\cal C}^n\), p_{ Ω } the probability of Ω in ρ, and ρ_{Ω} the final state conditioned on Ω. Let ε_{s} ∈ (0, 1).
For f_{min}, a mintradeoff function for \(\left\{ {{\cal M}_i} \right\}\), \({\hat{ \Omega }} = \left\{ {\left. {\rm freq}({c_1^n}) \rightc_1^n \in { {\Omega }}} \right\}\) convex, and any \(t \in {\Bbb R}\) such that \(f_{{\mathrm{min}}}\left( {\rm freq}({c_1^n}) \right) \ge t\) for any \(c_1^n \in {\cal C}^n\) for which \({\mathrm{Pr}}\left[ {c_1^n} \right]_{{\mathrm{\rho }}_{{{\Omega }}}} > 0\),
where
Similarly, for f_{max} a maxtradeoff function and \(t \in {\Bbb R}\) such that \(f_{{\mathrm{max}}}\left({\rm freq}({c_1^n}) \right) \le t\) for any \({\bf{c}} \in {\cal C}^n\) for which \({\mathrm{Pr}}\left[ {\bf{c}} \right]_{{\mathrm{\rho }}_{{{\Omega }}}} > 0\),
The two most important properties of the above statement are that the firstorder term is linear in n and that t is the von Neumann entropy of a suitable state (as explained above). This implies that the EAT is tight to first order in n.
We remark that the Markov chain conditions are important, in the sense that dropping them completely would render the statement invalid.
We now give a rough proof sketch of the \(H_{{\mathrm{min}}}^\varepsilon\) case; the bound for \(H_{{\mathrm{max}}}^\varepsilon\) follows from an almost identical argument. The proof has a similar structure to that of the quantum AEP^{31}, which we can retrieve as a special case. The proof relies heavily on the “sandwiched” Rényi entropies^{36,37}, which is a family of entropies that we will denote here by H_{ α }, where α is a real parameter ranging from \(\frac{1}{2}\) to ∞, and which corresponds to the maxentropy at \(\alpha = \frac{1}{2}\), to the von Neumann entropy when α = 1, and to the minentropy when α = ∞.
The basic idea is to first lower bound the \(H_{{\mathrm{min}}}^\varepsilon\) term by H_{ α } using the following general bound^{31,38,39,40}:
Then, we lower bound the H_{ α } term by the von Neumann entropy using the following^{31,39}:
Now, we could simply chain these two inequalities and apply them to \(H_{{\mathrm{min}}}^\varepsilon \left( {\left. {O_1^n} \rightS_1^nE} \right)\). However, this would result in a very poor bound due to the \(O\left( {\left( {{\mathrm{log}}\,d_A} \right)^2} \right)\) term in Eq. (13), which in our case would be O(n^{2}). To get the bound we want, we need to reduce this term to O(n); choosing \(\alpha \approx 1 + \frac{1}{{\sqrt n }}\) would then produce a bound with the right scaling.
The trick we use to achieve this is to decompose \(H_\alpha \left( {O_1^n\left {S_1^nE} \right.} \right)\) into n terms of constant size before applying Eq. (13). In the quantum AEP^{31}, this step is immediate since the state is i.i.d. Here, we must use more sophisticated techniques. Specifically, we use the following chain rule for the sandwiched Rényi entropy to decompose \(H_\alpha \left( {O_1^n\left {S_1^nE} \right.} \right)\) into n terms:
Theorem 4: Let \(\rho _{RA_1B_1}^0\) be a density operator on R ⊗ A_{1} ⊗ B_{1} and \({\cal M} = {\cal M}_{A_2B_2 \leftarrow R}\) be a CPTP map. Assuming that \(\rho _{A_1B_1A_2B_2} = {\cal M}( {\rho _{RA_1B_1}^0} )\) satisfies the Markov condition A_{1}↔B_{1}↔B_{2},we have
where the infimum ranges over density operators \(\omega _{RA_1B_1}\) on R ⊗ A_{1} ⊗ B_{1}. Moreover, if \(\rho _{RA_1B_1}^0\) is pure, then we can optimise over pure states \(\omega _{RA_1B_1}\).
Implementing this proof strategy then yields the following chain of inequalities:
However, this does not yet take into account the sampling over the C_{ i } subsystems. To do this, we tweak the EAT channels \({\cal M}_i\) to output two extra systems D_{ i } and \(\bar D_i\) which contain an amount of entropy that depends on the value of C_{ i } observed. To define this, let g be an affine lower bound on f_{min}, let \(\left[ {g_{{\mathrm{min}}},g_{{\mathrm{max}}}} \right]\) be the smallest interval that contains the range of g, and set \(\bar g: = \frac{1}{2}\left( {g_{{\mathrm{min}}} + g_{{\mathrm{max}}}} \right)\). Then, we define \({\cal D}_i:C_i \to C_iD_i\bar D_i\) as
where τ(c) is a mixture between a maximally entangled state and a fully mixed state such that the marginal on \(\bar D_i\) is uniform, and such that \(H_\alpha \left( {D_i\left {\bar D_i} \right.} \right)_{\tau (c)} = \bar g  g(\delta _c)\) (here δ_{ c } stands for the distribution with all the weight on element c). To ensure that this is possible, we need to choose \(d_{D_i}\) large enough, and it turns out that setting \(d_{D_i} = \left\lceil {2^{\parallel \nabla g\parallel _\infty }} \right\rceil\) suffices. We can then define a new sequence of EAT channels by \(\bar {\cal M}_i = {\cal D}_i \circ {\cal M}_i\).
Armed with this, we apply the above argument to our new EAT channels. On the one hand, a more sophisticated version of Eq. (12) yields:
On the other hand, the argument from Eq. (15) can be used here to give
Combining these two bounds then yields the theorem.
We remark that some of the concepts used in this work generalise techniques proposed in the recent security proofs for DI cryptography^{29}.
Entropy accumulation protocol
To analyse the key rates of the DIQKD protocol, we first find a lower bound on the amount of entropy accumulated during the run of the protocol, when the honest parties use their device to play the Bell games repeatedly. To this end, we consider the “entropy accumulation protocol” shown in Box 2. This protocol can be seen as the main building block of many DI cryptographic protocols.
The entropy accumulation protocol creates m blocks of bits, each of maximal length \(s_{\max }\). Each block ends (with high probability) with a test round; this is a round in which Alice and Bob play the CHSH game with their device so that they can verify that the device acts as expected. The probability of each round to be a test round is γ. The rest of the rounds are generation rounds, in which Bob chooses a special input for his component of the device. In the end of the protocol, Alice and Bob check whether they had sufficiently many test rounds in which they won the CHSH game. If not, they abort.
We note that the protocol is complete, in the sense that there exists an honest implementation of it (possibly noisy) which does not abort with high probability. Denoting the completeness error, i.e., the probability that the protocol aborts for an honest implementation of the devices D, by \(\varepsilon _{EA}^c\), one can easily show using Hoeffding’s inequality that for an honest i.i.d. implementation \(\varepsilon _{EA}^c \le {\mathrm{exp}}\left( {  2n\delta _{{\mathrm{est}}}^2} \right)\).
Next, we show that the protocol is also sound. That is, for any device D, if the probability that the protocol does not abort is not too small, then the total amount of smooth minentropy is sufficiently high.
The EAT can be used to bound the total amount of smooth minentropy, \(H_{{\mathrm{min}}}^{\varepsilon _{\mathrm{s}}}\left( {\left. {A_1^nB_1^n} \rightX_1^nY_1^nT_1^nE} \right)_{\rho _{{ {\Omega }}}}\), created when running the entropy accumulation protocol, given that it did not abort. Here n denotes the expected number of rounds of the protocol and ε_{s} is one of the security parameters (to be fixed later).
Below we use the following notation. For each block j ∈ [m], \(\vec A_j\) denotes the string that includes Alice’s outputs in block j (note that the length of this string is unknown, but it is at most \(s_{\max }\)). \(\vec B_j,\vec X_j\), and \(\vec Y_j\) are defined analogously. To use the EAT, we make the following choices of random variables:
The event Ω is the event of not aborting the protocol, as given in Step 11 in Box 2:
The EAT channels are chosen to be
where \({\cal M}_j\) describes Steps 2–10 of block j in the entropy accumulation protocol (Box 2). These channels include both the actions made by Alice and Bob as well as the operations made by the device D in these steps. Note that the Device’s operations can always be described within the formalism of quantum mechanics, although we do not assume we know them. The registers R_{j−1} and R_{ j } hold the quantum state of the device in the beginning and the end of the j’th step of the protocol, respectively.
Lemma 5: The channels \({\cal M}_j\) described above are EAT channels.
Proof. For the channels to be EAT channels, they need to fulfil the conditions given in Definition 1. We show that this is indeed the case. First, \(\tilde C_j\) are classical registers with \(\tilde C_j \in \{ 0,1, \bot \}\) and \(d_{\vec A_j} \times d_{\vec B_j} \le 6^{s_{{\mathrm{max}}}}\). Second, \(\tilde C_j\) is determined by the classical registers \(\vec A_j,\vec B_j,\vec X_j,\vec Y_j,\vec T_j\) as shown in Box 2. Therefore, \(\tilde C_j\) can be calculated without modifying the marginal on those registers. The third condition is also fulfilled since the inputs are chosen independently in each round and hence \(\vec A_{1 \ldots j  1}\vec B_{1 \ldots j  1} \leftrightarrow \vec X_{1 \ldots j  1}\vec Y_{1 \ldots j  1}\vec T_{1 \ldots j  1}E \leftrightarrow \vec X_j\vec Y_j\vec T_j\) trivially holds.
To continue one should devise a mintradeoff function. Let \(\tilde p\) be the probability distribution describing \(\tilde C_j\). We remark that due to the structure of our EAT channels, it is sufficient to consider \(\tilde p_{}^{}\) for which \(\tilde p(1) + \tilde p(0) = 1  (1  \gamma )^{s_{{\mathrm{max}}}}\) (otherwise the set Σ defined in Eq. (6) is an empty set).
The following lemma gives a lower bound on the von Neumann entropy of the outputs in a single block.
Lemma 6: Let \(\bar s = \frac{{1  (1  \gamma )^{s_{{\mathrm{max}}}}}}{\gamma }\) be the expected length of a block and h the binary entropy function. Then,
where the entropy is evaluated on a state that wins the CHSH game, in the test round, with probability
Proof sketch. The amount of entropy accumulated in a single round in a block is given in Eq. (2) in the main text. To get the amount of entropy accumulated in a block, one can use the chain rule for the von Neumann entropy. The result is then
where the prefactors (1 − γ)^{(i−1)} are attributed to the fact that the entropy in each round is nonzero only if the round is part of the block, i.e., if a test round was not performed before the i’th round in the block, and ω_{ i } denotes the winning probability in the i’th round (given that a test was not performed before).
The value of each ω_{ i } is not fixed completely given ω^{*}. However, by the operation of the EAT channels the following relation holds:
To conclude the proof, we thus need to minimise \(H\left( {\vec A_j\vec B_j\vec X_j\vec Y_j\vec T_jR{\prime}} \right)\) under the above constraint. Using standard techniques, e.g., Lagrange multipliers, one can see that the minimal value of this entropy is achieved for ω_{ i } = ω^{*} for all i and the lemma follows.
The bound given in the above lemma can now be used to define the mintradeoff function \(f_{{\mathrm{min}}}\left( {\tilde p} \right)\). As the derivative of the function plays a role in the final bound, we must make sure it is not too large at any point. This can be enforced by “cutting” the function at a chosen point \(\tilde p_t\) and “gluing” it to a linear function starting at that point, as shown in Fig. 6. \(\tilde p_t\) can be chosen depending on the other parameters such that the total amount of smooth minentropy is maximal. Following this idea, the resulting mintradeoff function is given by
Let ε_{ EA } be the desired error probability of the entropy accumulation protocol. We can then use Theorem 3 to say that either the probability of the protocol aborting is greater than 1 − ε_{ EA } or the following bound on the total smooth minentropy holds:
where
To illustrate the behaviour of the entropy rate η_{opt}, we plot it as a function of the expected Bell violation ω_{exp} in Fig. 7 for γ = 1 and s_{max} = 1. For comparison, we also plot in Fig. 7 the asymptotic rate (n → ∞) under the assumption that the state of the device is an (unknown) i.i.d. state. In this case, the quantum AEP appearing in Eq. (1) implies that the optimal rate is the von Neumann entropy accumulated in one round of the protocol (as given in Eq. (2)). This rate, appearing as the dashed line in Fig. 7, is an upper bound on the entropy that can be accumulated. One can see that as the number of rounds in the protocol increases, our rate η_{opt} approaches this optimal rate.
For the calculations of the DIQKD rates later on, we choose \(s_{{\mathrm{max}}} = \left\lceil {{\textstyle{1 \over \gamma }}} \right\rceil\). For this choice, the firstorder term of η_{opt} is linear in n and a short calculation reveals that the secondorder term scales, roughly, as \(\sqrt {n/\gamma }\).
Our DIQKD protocol, shown in Box 3, is based on the entropy accumulation protocol described above. In the first part of the protocol Alice and Bob use their devices to produce the raw data, similarly to what is done in the entropy accumulation protocol. The main difference is that Bob’s outputs always contains his measurement outcomes (instead of being set to \(\bot\) in the generation rounds); to make the distinction explicit, we denote Bob’s outputs in the DIQKD protocol with a tilde, \(\tilde B_1^n\).
We now describe the three postprocessing steps, error correction, parameter estimation, and privacy amplification, in more detail.
Error correction
Alice and Bob use an error correction protocol EC to obtain identical raw keys K_{A} and K_{B} from their bits \(A_1^n\), \(\tilde B_1^n\). In our analysis, we use a protocol, based on universal hashing, which minimises the amount of leakage to the adversary^{41,42}. To implement this protocol, Alice chooses a hash function and sends the chosen function and the hashed value of her bits to Bob. We denote this classical communication by O. Bob uses O, together with his prior knowledge \(\tilde B_1^nX_1^nY_1^nT_1^n\), to compute a guess \(\hat A_1^n\) for Alice’s bits \(A_1^n\). If EC fails to produce a good guess, Alice and Bob abort; in an honest implementation, this happens with probability at most \(\varepsilon _{{\mathrm{EC}}}^c\). If Alice and Bob do not abort, then they hold raw keys \(K_{\mathrm{A}} = A_1^n\) and \(K_{\mathrm{B}} = \hat A_1^n\) and K_{A} = K_{B} with probability at least 1 − ε_{EC}.
Due to the communication from Alice to Bob, leak_{EC} bits of information are leaked to the adversary. The following guarantee holds for the described protocol^{42}:
for \(\varepsilon _{{\mathrm{EC}}}^c = \varepsilon\prime _{{\mathrm{EC}}}+ \varepsilon _{{\mathrm{EC}}}\) and where \(H_0^{\varepsilon\prime _{{\mathrm{EC}}} }\left( {A_1^n\tilde B_1^nX_1^nY_1^nT_1^n} \right)\) is evaluated on the state in an honest implementation of the protocol. If a larger fraction of errors occur when running the actual DIQKD protocol (for instance due to adversarial interference) the error correction might not succeed, as Bob will not have a sufficient amount of information to obtain a good guess of Alice’s bits. If so, this will be detected with probability at least 1 − ε_{EC} and the protocol will abort. In an honest implementation of the device, Alice and Bob’s outputs in the generation rounds should be highly correlated in order to minimise the leakage of information.
Parameter estimation
After the error correction step, Bob has all of the relevant information to perform parameter estimation from his data alone, without any further communication with Alice. Using \(\tilde B_1^n\) and K_{B}, Bob sets \(\tilde C_j = w_{{\mathrm{CHSH}}}\left( {\hat A_i,\tilde B_i,X_i,Y_i} \right) = w_{{\mathrm{CHSH}}}\left( {K_{{\mathrm{B}}i},\tilde B_i,X_i,Y_i} \right)\) for the blocks with a test round (which was done at round i of the block) and \(\tilde C_j = \bot\) otherwise. He aborts if the fraction of successful test rounds is too low, that is, if \(\mathop {\sum}\nolimits_{j \in \left[ m \right]} {\kern 1pt} \tilde C_j < \left[ {\omega _{{\mathrm{exp}}}\left( {1  (1  \gamma )^{s_{{\mathrm{max}}}}} \right)  \delta _{{\mathrm{est}}}} \right] \cdot m\).
As Bob does the estimation using his guess of Alice’s bits, the probability of aborting in this step in an honest implementation, \(\varepsilon _{{\mathrm{PE}}}^c\), is bounded by \(\varepsilon _{{\mathrm{EA}}}^c + \varepsilon _{{\mathrm{EC}}}\).
Privacy amplification
Finally, Alice and Bob use a (quantumproof) privacy amplification protocol PA (which takes some random seed S as input) to create their final keys \(\tilde K_{\mathrm{A}}\) and \(\tilde K_{\mathrm{B}}\) of length \(\ell\), which are close to ideal keys, i.e., uniformly random and independent of the adversary’s knowledge.
For simplicity, we use universal hashing^{43} as the privacy amplification protocol in the analysis below. Any other quantumproof strong extractor, e.g., Trevisan’s extractor^{44}, can be used for this task and the analysis can be easily adapted.
The secrecy of the final key depends only on the privacy amplification protocol used and the value of \(H_{{\mathrm{min}}}^{\varepsilon _{\mathrm{s}}}\left( {A_1^nX_1^nY_1^nT_1^nOE} \right)\), evaluated on the state at the end of the protocol, conditioned on not aborting. For universal hashing, for every ε_{PA},ε_{s}∈(0, 1), a secure key of maximal length
is produced with probability at least 1 − ε_{PA} − ε_{s}.
Correctness, secrecy, and overall security of a DIQKD protocol are defined as follows^{45}:
Definition 7 (Correctness): A DIQKD protocol is said to be ε_{corr}correct, when implemented using a device D, if Alice and Bob’s keys, \(\tilde K_{\mathrm{A}}\) and \(\tilde K_{\mathrm{B}}\) respectively, are identical with probability at least 1 − ε_{corr}. That is, \({\mathrm{Pr}}\left( {\tilde K_{\mathrm{A}} \ne \tilde K_{\mathrm{B}}} \right) \le \varepsilon _{{\mathrm{corr}}}\).
Definition 8 (Secrecy): A DIQKD protocol is said to be ε_{sec}secret, when implemented using a device D, if for a key of length l, \(\left( {1  {\mathrm{Pr}}\left[ {{\mathrm{abort}}} \right]} \right)\left\ {\rho _{\tilde K_{\mathrm{A}}E}  \rho _{U_l} \otimes \rho _E} \right\_1 \le \varepsilon _{{\mathrm{sec}}}\), where E is a quantum register that may initially be correlated with D.
ε_{sec} in the above definition can be understood as the probability that some nontrivial information leaks to the adversary^{45}. If a protocol is ε_{corr}correct and ε_{sec}secret (for a given D), then it is \(\varepsilon _{{\mathrm{QKD}}}^s\)correctandsecret for any \(\varepsilon _{{\mathrm{QKD}}}^s \ge \varepsilon _{{\mathrm{corr}}} + \varepsilon _{{\mathrm{sec}}}\).
Definition 9 (Security): A DIQKD protocol is said to be \(\left( {\varepsilon _{{\mathrm{QKD}}}^s,\varepsilon _{{\mathrm{QKD}}}^c,l} \right)\)secure if:

1.
(Soundness) For any implementation of the device D it is \(\varepsilon _{{\mathrm{QKD}}}^s\)correctandsecret.

2.
(Completeness) There exists an honest implementation of the device D such that the protocol does not abort with probability greater than \(1  \varepsilon _{{\mathrm{QKD}}}^c\).
Below we show that the following theorem holds.
Theorem 10: The DIQKD protocol described above is \(\left( {\varepsilon _{{\mathrm{QKD}}}^s,\varepsilon _{{\mathrm{QKD}}}^c,\ell } \right)\)secure, with \(\varepsilon _{{\mathrm{QKD}}}^s \le \varepsilon _{{\mathrm{EC}}} + \varepsilon _{{\mathrm{PA}}} + \varepsilon _{\mathrm{s}} + \varepsilon _{{\mathrm{EA}}}\), \(\varepsilon _{{\mathrm{QKD}}}^c \le \varepsilon _{EC}^c + \varepsilon _{{\mathrm{EA}}}^c + \varepsilon _{{\mathrm{EC}}}\), and
where \(t = \sqrt {  m\left( {1  \gamma } \right)^2{\kern 1pt} {\mathrm{log}}{\kern 1pt} \varepsilon _{\mathrm{t}}{\mathrm{/}}2\gamma ^2}\) for any ε_{t} ∈ (0, 1).
We now explain the steps taken to prove Theorem 10. The completeness part follows trivially from the completeness of the “subprotocols”.
To establish soundness, first note that by definition, as long as the protocol does not abort it produces a key of length \(\ell\). Therefore, it remains to verify correctness, which depends on the error correction step, and security, which is based on the privacy amplification step. To prove security we start with Lemma 11, in which we assume that the error correction step is successful. We then use it to prove soundness in Lemma 12.
Let \({\tilde{ \Omega }}\) denote the event of the DIQKD protocol not aborting and the EC protocol being successful, and let \(\tilde \rho _{A_1^n\tilde B_1^nX_1^nY_1^nT_1^nO_1^nE{\tilde{ \Omega }}}\) be the state at the end of the protocol, conditioned on this event.
Success of the privacy amplification step relies on the minentropy \(H_{{\mathrm{min}}}^{\varepsilon _{\mathrm{s}}}\left( {A_1^nX_1^nY_1^nT_1^nOE} \right)_{\tilde \rho _{{\tilde{ \Omega }}}}\) being sufficiently large. The following lemma connects this quantity to \(H_{{\mathrm{min}}}^{\frac{{\varepsilon _{\mathrm{s}}}}{4}}\left( {A_1^nB_1^nX_1^nY_1^nT_1^nE} \right)_{\rho _{{ {\Omega }}}}\), on which a lower bound is provided in Eq. (31) above.
Lemma 11: For any device D, let \(\tilde \rho\) be the state generated in the protocol right before the privacy amplification step. Let \(\tilde \rho _{{\tilde{ \Omega }}}\) be the state conditioned on not aborting the protocol and success of the EC protocol. Then, for any ε_{EA}, ε_{EC}, ε_{s}, ε_{t} ∈ (0, 1), either the protocol aborts with probability greater than 1 − ε_{EA} − ε_{EC} or
Proof sketch. Before deriving a bound on the entropy of interest, we remark that the t is chosen such that the probability that the actual number of rounds in the protocol, N, is larger than the expected number of rounds n plus t is ε_{t}. The above value for t can be derived by noticing that the sizes of the blocks are i.i.d. random variables which take values in [1, 1/γ].
The key idea of the proof is to consider the following events:

1.
Ω: the event of not aborting in the entropy accumulation protocol. This happens when the Bell violation, calculated using Alice and Bob’s outputs (and inputs), is sufficiently high.

2.
\({\hat{ \Omega }}\): Suppose Alice and Bob run the entropy accumulation protocol, and then execute the EC protocol. The event \({\hat{ \Omega }}\) is defined by Ω and \(K_{\mathrm{B}} = A_1^n\).

3.
\({\tilde{ \Omega }}\): the event of not aborting the DIQKD protocol and \(K_{\mathrm{B}} = A_1^n\).
The state \(\rho _{{\hat{ \Omega }}}\) then denotes the state at the end of the entropy accumulation protocol conditioned on \({\hat{ \Omega }}\).
Using a sequence of chain rules for smooth entropies^{46} and the fact that \(\tilde \rho _{A_1^nX_1^nY_1^nT_1^nE{\tilde{ \Omega }}} = \rho _{A_1^nX_1^nY_1^nT_1^nE{\hat{ \Omega }}}\) (\(\tilde B_1^n\) and \(B_1^n\) were traced out from \(\tilde \rho\) and ρ, respectively), one can conclude
\(H_{{\mathrm{max}}}^{\frac{{\varepsilon _{\mathrm{s}}}}{4}}\left( {B_1^nT_1^nE} \right)_{\rho _{{\hat{ \Omega }}}}\) can be bounded from above. The intuition is that \(B_i \ne \bot\) only when T_{ i } = 0, which happens with probability γ. The exact bound can be calculated using the EAT and is given by
The above steps together with Eq. (31) conclude the proof.
Using Lemma 11, one can prove that our DIQKD protocol is sound.
Lemma 12: For any device D let \(\tilde \rho\) be the state generated by the DIQKD protocol. Then either the protocol aborts with probability greater than 1 − ε_{EA} − ε_{EC} or it is (ε_{EC} + ε_{PA} + ε_{s})correctandsecret while producing keys of length \(\ell\), as defined in Eq. (35).
Proof sketch. Assume the DIQKD protocol did not abort. We consider two cases. First, assume that the EC protocol was not successful (but did not abort). Then Alice and Bob’s final keys might not be identical. This happens with probability at most ε_{EC}. Otherwise, assume the EC protocol was successful, i.e., \(K_{\mathrm{B}} = A_1^n\). In that case, Alice and Bob’s keys must be identical also after the final privacy amplification step.
The secrecy depends only on the privacy amplification step, and for universal hashing a secure key is produced as long as Eq. (34) holds. Hence, a uniform and independent key of length \(\ell\) as in Eq. (35) is produced by the privacy amplification step unless the smooth minentropy is not high enough or the privacy amplification protocol was not successful, which happens with probability at most ε_{PA} + ε_{s}.
According to Lemma 11, either the protocol aborts with probability greater than 1 − ε_{EA} − ε_{EC} or the entropy is sufficiently high to create the secret key.
The expected key rates appearing in Fig. 5 in the main text are given by \(r = \ell {\mathrm{/}}n\). The key rate depends on the amount of leakage of information due to the error correction step, which in turn depends on the honest implementation of the protocol as mentioned above. To have an explicit bound, we consider the honest implementation described in the main text. Using Eq. (33) and the AEP, one can show that the amount of leakage in the error correction step is then given by
To get the optimal key rates, one should fix the parameters of interest (e.g., \(\varepsilon _{{\mathrm{QKD}}}^s\), \(\varepsilon _{{\mathrm{QKD}}}^c\), and n) and optimise over all other parameters.
DI randomness expansion
The entropy accumulation protocol can be used to perform DI randomness expansion as well. In a DI randomness expansion protocol, the honest parties start with a short seed of perfect randomness and use it to create a longer random secret string. For the purposes of randomness expansion, we may assume that the parties are colocated, therefore, the main difference from the DIQKD scheme is that there is no need for error correction (and hence there is no leakage of information due to public communication).
In order to minimise the amount of randomness required to execute the protocol, we adapt the main entropy accumulation protocol by deterministically choosing inputs in the generation rounds X_{ i },Y_{ i }∈{0,1}. In particular, there is no use for the input 2 to Bob’s device, and no randomness is required for the generation rounds. Aside from the last step of privacy amplification, the remainder of the protocol is essentially the same as the entropy accumulation protocol.
The plotted entropy rates in Fig. 7 are therefore also the ones relevant for a DI randomness expansion.
Since we are concerned here not only with generating randomness but also with expanding the amount of randomness initially available to the users of the protocol, we should also evaluate the total number of random bits that are needed to execute the protocol. Random bits are required to select which rounds are generation rounds, i.e., the random variable \(T_1^n\), to select inputs to the devices in the testing rounds, i.e., those for which T_{ i } = 0, and to select the seed for the quantum proof extractor used for privacy amplification. All of these can be accounted for using standard techniques and so we omit the detailed explanation and formulas.
Data availability
No data sets were generated or analysed during the current study.
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Acknowledgements
We thank Asher Arnon for the illustrations presented in Figs. 1, 4, and 5. R.A.F. and R.R. were supported by the Stellenbosch Institute for Advanced Study (STIAS), by the European Commission via the project “RAQUEL”, by the Swiss National Science Foundation (grant No. 200020–135048) and the National Centre of Competence in Research “Quantum Science and Technology”, by the European Research Council (grant No. 258932), and by the US Air Force Office of Scientific Research (grant No. FA95501610245). F.D. acknowledges the financial support of the Czech Science Foundation (GA ČR) project no GA1622211S and of the European Commission FP7 Project RAQUEL (grant No. 323970). O.F. acknowledges support from the LABEX MILYON (ANR10LABX0070) of Université de Lyon, within the program “Investissements d’Avenir” (ANR11IDEX0007) operated by the French National Research Agency (ANR). T.V. was partially supported by NSF CAREER Grant CCF1553477, an AFOSR YIP award, the IQIM, and NSF Physics Frontiers Center (NFS Grant PHY1125565) with support of the Gordon and Betty Moore Foundation (GBMF12500028).
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Institute for Theoretical Physics, ETHZürich, WolfgangPauliStr. 27, 8093, Zürich, Switzerland
 Rotem ArnonFriedman
 & Renato Renner
Faculty of Informatics, Masaryk University, Brno, Czech Republic
 Frédéric Dupuis
CNRS, LORIA, Université de Lorraine, Campus scientifique, BP 239, 54506, VandoeuvrelèsNancy Cedex, France
 Frédéric Dupuis
Laboratoire de l’Informatique du Parallélisme, LIP, ENS de Lyon, 46 Allee d’Italie, 69364, Lyon Cedex 07, France
 Omar Fawzi
Department of Computing and Mathematical Sciences, California Institute of Technology, 1200 E. California Blvd, Pasadena, CA, 91125, USA
 Thomas Vidick
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R.A.F., F.D., O.F., R.R., and T.V. contributed equally to this work.
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The authors declare that they have no competing financial interests.
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Correspondence to Rotem ArnonFriedman.
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