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% This is samplepaper.tex, a sample chapter demonstrating the
% LLNCS macro package for Springer Computer Science proceedings;
% Version 2.20 of 2017/10/04
%
\documentclass[runningheads]{llncs}
%
\usepackage{graphicx}
%\usepackage{hyperref}
\usepackage{amsmath,amssymb}
% Used for displaying a sample figure. If possible, figure files should
% be included in EPS format.
%
% If you use the hyperref package, please uncomment the following line
% to display URLs in blue roman font according to Springer's eBook style:
% \renewcommand\UrlFont{\color{blue}\rmfamily}
\usepackage[utf8]{inputenc}
\begin{document}
\title{Robust and Incometransparent Online ECash}
\maketitle
\begin{abstract}
We present an anonymous ecash protocol that protects against economic losses
when protocols are aborted. Furthermore, we ensure that customers have
cryptographic evidence that they
spent ecash on a particular business transaction and introduce \emph{conservation}
as an additional security property to anonymity and unforgeability, and
preserve anonymity when wallets are restored from
backup or synchronized with other devices. We argue from this position that
a protocol for unlinkable change is necessary even in schemes that provide
divisibility. As a na\"ive implementation of a change protocol opens up the
possibility of abuse for tax evasion, we define a new \emph{income transparency}
security property.
We furthermore show that an ecash protocol that fulfills these properties
can be used to implement Camenischstyle fair exchange, tick payments,
and
can be used to provide anonymous refunds.
\keywords{Ecash \and blind signature \and key exchange}
\end{abstract}
\def\Z{\mathbb{Z}}
\def\mathperiod{.}
\def\mathcomma{,}
\newcommand*\ST[5]%
{\left#1\,#4\vphantom{#5} \;\right#2 \left. #5 \vphantom{#4}\,\right#3}
% uniform random selection from set
\newcommand{\randsel}[0]{\ensuremath{\xleftarrow{\text{\$}}}}
\newcommand{\Exp}[1]{\ensuremath{E\left[#1\right]}}
% oracles
\newcommand{\ora}[1]{\ensuremath{\mathcal{O}\mathsf{#1}}}
% oracle set
\newcommand{\oraSet}[1]{\ensuremath{\mathcal{O}\textsc{#1}}}
% algorithm
\newcommand{\algo}[1]{\ensuremath{\mathsf{#1}}}
% party
\newcommand{\prt}[1]{\ensuremath{\mathcal{#1}}}
% longform variable
\newcommand{\V}[1]{\ensuremath{\mathsf{#1}}}
% probability with square brackets of the right size
\newcommand{\Prb}[1]{\ensuremath{\Pr\left [#1 \right ]}}
\newcommand{\mycomment}[1]{~\\ {\small \textcolor{blue}{({#1})}}}
%\theoremstyle{definition}
%\newtheorem{definition}{Definition}[section]
%\theoremstyle{corollary}
%\newtheorem{corollary}{Corollary}[section]
\section{Model and Syntax}
We consider a Chaumstyle~\cite{chaum1983blind}
payment system with an exchange (Chaum: mint) and multiple,
dynamically created customers and merchants.
We model withdrawing digital coins, spending them with
merchants and subsequently depositing them at the exchange, as well as
obtaining unlinkable change for partially spent coins with a
``refresh'' protocol.
The exchange offers digital coins in multiple denominations. We mostly
ignore the denomination
values here, including their impact on anonymity, in keeping with existing
literature~\cite{camenisch2007endorsed,pointcheval2017cut}. For anonymity, we
believe this amounts to assuming that all customers have similar financial
behavior. We note logarithmic storage, computation and bandwidth demands
denominations distributed by powers of a fixed constant.
%We do not model fees taken by the exchange. Reserves\footnote{%
% ``Reserve'' is Taler's terminology for funds submitted to the exchange that
% can be converted to digital coins.
%}
%are also omitted.
Coins can be partially spent by specifying a fraction $0 < f \le 1$ of the
total value associated with the coin's denomination. Unlinkable change below
the smallest denomination cannot be given. In
practice the unspendable, residual value should be seen as a fee
charged by the exchange.
Spending multiple coins is modeled nonatomically: to spend (fractions
of) multiple coins, they must be spent onebyone. The individual
spend/deposit operations are correlated by a unique identifier for the
transaction.
%In practice this identifier is the hash $\V{transactionId} =
%H(\V{contractTerms})$ of the contract terms\footnote{The contract terms
%are a digital representation of an individual offer for a certain product or service the merchant sells
%for a certain price.}. Contract terms include a nonce to make them
%unique, that merchant and customer agreed upon. Note that this transaction
%identifier and the correlation between multiple spend operations for one
%payment need not be disclosed to the exchange (it might, however, be necessary
%to reveal during a detailed tax audit of the merchant): When spending the $i$th coin
%for the transaction with the identifier $\V{transactionId}$, messages to the
%exchange would only contain $H(i \Vert \V{transactionId})$. This is preferable
%for merchants that might not want to disclose to the exchange the individual
%prices of products they sell to customers, but only the total transaction
%volume over time. For simplicity, we do not include this extra feature in our
%model.
Our system model tracks the total amount of
coins withdrawn by each customer.
Customers are identified by their public key $\V{pkCustomer}$. Every customer's wallet
keeps track of the following data:
\begin{itemize}
\item $\V{wallet}[\V{pkCustomer}]$ contains sets of the customer's coin records,
which individually consist of the coin key pair, denomination and exchange's signature.
\item $\V{acceptedContracts}[\V{pkCustomer}]$ contains the sets of
transaction identifiers accepted by the customer during spending
operations, together with coins spent for it and their contributions $0 < f
\le 1$.
\item $\V{withdrawIds}[\V{pkCustomer}]$ contains the withdraw identifiers of
all withdraw operations that were created for this customer.
\item $\V{refreshIds}[\V{pkCustomer}]$ contains the refresh identifiers of
all refresh operations that were created for this customer.
\end{itemize}
\noindent
The exchange in our model keeps track of the following data:
\begin{itemize}
\item $\V{withdrawn}[\V{pkCustomer}]$ contains the total amount withdrawn by
each customer, i.e. the sum of the financial value of the denominations for
all coins that were withdrawn by $\V{pkCustomer}$.
\item The overspending database of the exchange is modeled by
$\V{deposited}[\V{pkCoin}]$ and $\V{refreshed}[\V{pkCoin}]$, which record
deposit and refresh operations respectively on each coin. Note that since
partial deposits and multiple refreshes to smaller denominations are
possible, one deposit and multiple refresh operations can be recorded for a
single coin.
\end{itemize}
We say that a coin is \emph{fresh} if it appears in neither the $\V{deposited}$
or $\V{refreshed}$ lists nor in $\V{acceptedContracts}$. We say that a coin is
being $\V{overspent}$ if recording an operation in $\V{deposited}$ or
$\V{refreshed}$ would cause the total spent value from both lists to exceed
the value of the coin's denomination.
%
Note that the adversary does not have direct read or write access to these
values; instead the adversary needs to use the oracles (defined in Section~\ref{sec:oracles}) to
interact with the system.
We parameterize our system with two security parameters: The general security
parameter $\lambda$, and the refresh security parameter $\kappa$. While
$\lambda$ determines the length of keys and thus the security level, using a
larger $\kappa$ will only decrease the success chance of malicious merchants
conspiring with customers to obtain unreported (and thus untaxable) income.
\subsection{Algorithms and Protocols} \label{sec:algorithms}
Our ecash scheme is modeled by the following probabilistic\footnote{Our
instantiation is not probabilistic (except key
generation), but we do not want to prohibit this for other instantiations.}
polynomialtime algorithms and interactive protocols. The notation $P(X_1,\dots,X_n)$
stands for a party $P \in \{\prt{E}, \prt{C}, \prt{M}\}$ (Exchange, Customer
and Merchant respectively) in an interactive protocol, with $X_1,\dots,X_n$
being the (possibly private) inputs contributed by the party to the protocol.
Interactive protocols can access the state maintained by party $P$.
While the adversary can freely execute the interactive protocols by creating
their own parties, the adversary is not given direct access to the private data
of parties maintained by the challenger in the security games which we define later.
\begin{itemize}
\item $\algo{ExchangeKeygen}(1^{\lambda}, 1^{\kappa}, \mathfrak{D}) \mapsto (\V{sksE}, \V{pksE})$:
Algorithm executed to generate keys for the exchange, with general security
parameter $\lambda$ and refresh security parameter $\kappa$, both given as
unary numbers. The denomination specification $\mathfrak{D} = d_1,\dots,d_n$ is a
finite sequence of positive rational numbers that defines the financial
value of each generated denomination key pair. We henceforth use $\mathfrak{D}$ to
refer to some appropriate denomination specification, but our analysis is
independent of a particular choice of $\mathfrak{D}$.
The algorithm generates the exchange's master signing key pair
$(\V{skESig}, \V{pkESig})$ and denomination secret and public keys
$(\V{skD}_1, \dots, \V{skD}_n), (\V{pkD}_1, \dots, \V{pkD}_n)$. We write
$D(\V{pkD}_i)$, where $D : \{\V{pkD}_i\} \rightarrow \mathfrak{D}$ to look
up the financial value of denomination $\V{pkD_i}$.
We collectively refer to the exchange's secrets by $\V{sksE}$ and to the exchange's
public keys together with $\mathfrak{D}$ by $\V{pksE}$.
\item $\algo{CustomerKeygen}(1^\lambda,1^\kappa) \mapsto (\V{skCustomer}, \V{pkCustomer})$:
Key generation algorithm for customers with security parameters $\lambda$
and $\kappa$.
\item $\algo{MerchantKeygen}(1^\lambda,1^\kappa) \mapsto (\V{skMerchant},
\V{pkMerchant})$: Key generation algorithm for merchants. Typically the
same as \algo{CustomerKeygen}.
\item $\algo{WithdrawRequest}(\prt{E}(\V{sksE}, \V{pkCustomer}),
\prt{C}(\V{skCustomer}, \V{pksE}, \V{pkD})) \mapsto (\mathcal{T}_{WR},
\V{wid})$: Interactive protocol between the exchange and a customer that
initiates withdrawing a single coin of a particular denomination.
The customer obtains a withdraw identifier $\V{wid}$ from the protocol
execution and stores it in $\V{withdrawIds}[\V{pkCustomer}]$.
The \algo{WithdrawRequest} protocol only initiates a withdrawal. The coin
is only obtained and stored in the customer's wallet by executing the
\algo{WithdrawPickup} protocol on the withdraw identifier \V{wid}.
The customer and exchange persistently store additional state (if required
by the instantiation) such that the customer can use $\algo{WithdrawPickup}$
Returns a protocol transcript $\mathcal{T}_{WR}$ of all messages exchanged
between the exchange and customer, as well as the withdraw identifier
\V{wid}.
\item $\algo{WithdrawPickup}(\prt{E}(\V{sksE}, \V{pkCustomer}),
\prt{C}(\V{skCustomer}, \V{pksE}, \V{wid})) \mapsto (\mathcal{T}_{WP},
\V{coin})$: Interactive protocol between the exchange and a customer to
obtain the coin from a withdraw operation started with
$\algo{WithdrawRequest}$, identified by the withdraw identifier $\V{wid}$.
The first time $\algo{WithdrawPickup}$ is run with a particular withdraw
identifier $\V{wid}$, the exchange increments
$\V{withdrawn}[\V{pkCustomer}]$ by $D(\V{pkD})$, where $\V{pkD}$ is the
denomination requested in the corresponding $\algo{WithdrawRequest}$
execution. How exactly $\V{pkD}$ is restored depends on the particular instantiation.
The resulting coin $\V{coin} := (\V{skCoin}, \V{pkCoin}, \V{pkD}, \V{coinCert})$,
consisting of secret key $\V{skCoin}$, public key
$\V{pkCoin}$, denomination public key $\V{pkD}$ and certificate $\V{coinCert}$ from the exchange,
is stored in the customers wallet $\V{wallet}[\V{pkCustomer}]$.
Executing the $\algo{WithdrawPickup}$ protocol multiple times with the same
customer and the same withdraw identifier does not result in any change of
the customer's withdraw balance $\V{withdrawn}[\V{pkCustomer}]$,
and results in \mbox{(re)}adding the same coin to the customer's wallet.
Returns a protocol transcript $\mathcal{T}_{WP}$ of all messages exchanged
between the exchange and customer.
\item $\algo{Spend}(\V{transactionId}, f, \V{coin}, \V{pkMerchant}) \mapsto \V{depositPermission}$:
Algorithm to produce and sign a deposit permission \V{depositPermission}
for a coin under a particular transaction identifier. The fraction $0 < f \le 1$ determines the
fraction of the coin's initial value that will be spent.
%
The contents of the deposit permission depend on the instantiation, but it
must be possible to derive the public coin identifier $\V{pkCoin}$ from
them.
\item $\algo{Deposit}(\prt{E}(\V{sksE}, \V{pkMerchant}), \prt{M}(\V{skMerchant}, \V{pksE}, \V{depositPermission})) \mapsto \mathcal{T}_D$:
Interactive protocol between the exchange and a merchant.
%
From the deposit permission we obtain the $\V{pkCoin}$ of the coin to be
deposited. If $\V{pkCoin}$ is being overspent, the protocol is aborted with
an error message to the merchant.
%
On success, we add $\V{depositPermission}$ to $\V{deposited}[\V{pkCoin}]$.
%
Returns a protocol transcript $\mathcal{T}_D$ of all messages exchanged
between the exchange and merchant.
\item $\algo{RefreshRequest}(\prt{E}(\V{sksE}), \prt{C}(\V{pkCustomer}, \V{pksE}, \V{coin}_0, \V{pkD}_u))
\rightarrow (\mathcal{T}_{RR}, \V{rid})$
Interactive protocol between exchange and customer that initiates a refresh
of $\V{coin}_0$. Together with $\algo{RefreshPickup}$, it allows the
customer to convert $D(\V{pkD}_u)$ of the remaining value on coin
$\V{coin}_0 = (\V{skCoin}_0, \V{pkCoin}_0, \V{pkD}_0, \V{coinCert}_0)$
into a new, unlinkable coin $\V{coin}_u$ of denomination $\V{pkD}_u$.
%
Multiple refreshes on the same coin are allowed, but each run subtracts the
respective financial value of $\V{coin}_u$ from the remaining value of
$\V{coin}_0$.
%
The customer only records the refresh operation identifier $\V{rid}$ in
$\V{refreshIds}[\V{pkCustomer}]$, but does not yet obtain the new coin. To
obtain the new coin, \algo{RefreshPickup} must be used.
%
Returns the protocol transcript $\mathcal{T}_{RR}$ and a refresh identifier $\V{rid}$.
\item $\algo{RefreshPickup}(\prt{E}(\V{sksE}, \V{pkCustomer}),
\prt{C}(\V{skCustomer}, \V{pksE}, \V{rid}) \rightarrow (\mathcal{T}_{RP}, \V{coin}_u)$:
Interactive protocol between exchange and customer to obtain the new coin
for a refresh operation previously started with \algo{RefreshRequest},
identified by the refresh identifier $\V{rid}$.
%
The first time \algo{RefreshPickup} is run for a particular refresh
identifier, the exchange tries to record the refresh operation of value $D(\V{pkD}_u)$
in $\V{refreshed}[\V{pkCoin}_0]$, with $\V{pkD}_u$ and $\V{pkCoin}_0$ taken
from the corresponding $\algo{RefreshRequest}$ that resulted in $\V{rid}$.
How exactly the exchange obtains $\V{pkD}_u$ and $\V{pkCoin}_0$ depends on
the instantiation.
%
If $\V{pkCoin}_0$ is being overspent, the refresh operation is not recorded
in $\V{refreshed}[\V{pkCoin}_0]$, the exchange sends the customer the
protocol transcript of the previous deposits and refreshes and aborts the
protocol.
%
If the customer \prt{C} plays honestly in \algo{RefreshRequest} and
\V{RefreshPickup}, the unlinkable coin $\V{coin}_i$ they obtain as change
will be stored in their wallet $\V{wallet}[\V{pkCustomer}]$. If \prt{C} is
caught playing dishonestly, the \algo{RefreshPickup} protocol aborts.
%
An honest customer must be able to repeat a \algo{RefreshPickup} with the
same $\V{rid}$ multiple times and (re)obtain the same coin, even if
previous $\algo{RefreshPickup}$ executions were aborted.
%
Returns a protocol transcript $\mathcal{T}_{RP}$.
\item $\algo{Link}(\prt{E}(\V{sksE}), \prt{C}(\V{skCustomer}, \V{pksE}, \V{coin}_0)) \rightarrow (\mathcal{T}, (\V{coin}_1, \dots, \V{coin}_n))$:
Interactive protocol between exchange and customer. If $\V{coin}_0$ is a
coin that was refreshed, the customer can recompute all the coins obtained
from previous refreshes on $\V{coin}_0$, with data obtained from the
exchange during the protocol. These coins are added to the customer's
wallet $\V{wallet}[\V{pkCustomer}]$ and returned.
\end{itemize}
\subsection{Oracles} \label{sec:oracles}
%We now specify how the adversary can interact with the system by defining
%oracles. Oracles are queried by the adversary, and upon a query the challenger
%will act according to the oracle's specification. Note that the adversary for
%the different security games is run with specific oracles, and does not
%necessarily have access to all oracles simultaneously.
We refer to customers in the parameters to an oracle query simply by their
public key. For coins, however, the situation is more complicated. The
adversary needs the ability to refer to coins to trigger operations such as
spending and refresh, but to model anonymity we cannot give the adversary access
to the coins' public keys directly. Therefore we allow the adversary to use
the (successful) transcripts of the withdraw, refresh and link protocols to
indirectly refer to coins. We refer to this as a coin handle $\mathcal{H}$.
Since the execution of a link protocol results in a transcript $\mathcal{T}$
that can contain multiple coins, the adversary needs to select a particular
coin from the transcript via the index $i$ as $\mathcal{H} = (\mathcal{T}, i)$.
The respective oracle tries to find the coin that resulted from the transcript
given by the adversary. If the transcript has not been seen before in the
execution of a link, refresh or withdraw protocol; or the index for a link
transcript is invalid, the oracle returns an error to the adversary.
In oracles that trigger the execution of one of the interactive protocols
defined in Section~\ref{sec:algorithms}, we give the adversary the ability to actively
control the communication channels between the exchange, customers and
merchants; i.e. the adversary can effectively record, drop, modify and inject
messages during the execution of the interactive protocol. Note that this
allows the adversary to leave the execution of an interactive protocol in an
unfinished state, where one or more parties are still waiting for messages. We
use $\mathcal{I}$ to refer to a handle to interactive protocols where the
adversary can send and receive messages.
\begin{itemize}
\item $\ora{AddCustomer}() \mapsto \V{pkCustomer}$:
Generates a key pair $(\V{skCustomer}, \V{pkCustomer})$ using the
\algo{CustomerKeygen} algorithm, and sets
\begin{align*}
\V{withdrawn}[\V{pkCustomer}] &:= 0\\
\V{acceptedContracts}[\V{pkCustomer}] &:= \{ \}\\
\V{wallet}[\V{pkCustomer}] &:= \{\} \\
\V{withdrawIds}[\V{pkCustomer}] &:= \{\} \\
\V{refreshIds}[\V{pkCustomer}] &:= \{\}.
\end{align*}
Returns the public key of the newly created customer.
There is no separate oracle for creating merchants, since no information is
tracked for merchants; generating a key pair with \algo{MerchantKeygen}
suffices.
\item $\ora{AddMerchant}() \mapsto \V{pkMerchant}$L
Generate a key pair $(\V{skMerchant}, \V{pkMerchant})$ using the
\algo{MerchantKeygen} algorithm.
Returns the public key of the newly created merchant.
\item $\ora{SendMessage}(\mathcal{I}, P_1, P_2, m) \mapsto ()$:
Send message $m$ on the channel from party $P_1$ to party $P_2$ in the
execution of interactive protocol $\mathcal{I}$.
\item $\ora{ReceiveMessage}(\mathcal{I}, P_1, P_2) \mapsto m$:
Read message $m$ in the channel from party $P_1$ to party $P_2$ in the execution
of interactive protocol $\mathcal{I}$. If no message is queued in the channel,
return $m = \bot$.
\item $\ora{WithdrawRequest}(\V{pkCustomer}, \V{pkD}) \mapsto \mathcal{I}$:
Triggers the execution of the \algo{WithdrawRequest} protocol. the
adversary full control of the communication channels between customer and
exchange.
\item $\ora{WithdrawPickup}(\V{pkCustomer}, \V{pkD}, \mathcal{T}) \mapsto \mathcal{I}$:
Triggers the execution of the \algo{WithdrawPickup} protocol, additionally giving
the adversary full control of the communication channels between customer and exchange.
The customer and withdraw identifier $\V{wid}$ are obtained from the \algo{WithdrawRequest} transcript $\mathcal{T}$.
\item $\ora{RefreshRequest}(\mathcal{H}, \V{pkD}) \mapsto \mathcal{I}$: Triggers the execution of the
\algo{RefreshRequest} protocol with the coin identified by coin handle
$\mathcal{H}$, additionally giving the adversary full control over the communication channels
between customer and exchange.
\item $\ora{RefreshPickup}(\mathcal{T}) \mapsto \mathcal{I}$:
Triggers the execution of the \algo{RefreshPickup} protocol, where the customer and refresh identifier $\V{rid}$
are obtained from the $\algo{RefreshRequest}$ protocol transcript $\mathcal{T}$.
Additionally gives the adversary full control over the communication channels
between customer and exchange.
\item $\ora{Link}(\mathcal{H}) \mapsto \mathcal{I}$: Trigger the execution of the
\algo{Link} protocol for the coin referenced by handle $\mathcal{H}$,
additionally giving the adversary full control over the communication channels
between customer and exchange.
\item $\ora{Spend}(\V{transactionId}, \V{pkCustomer}, \mathcal{H}, \V{pkMerchant}) \mapsto \V{depositPermission}$:
Makes a customer sign a deposit permission over a coin identified by handle
$\mathcal{H}$. Returns the deposit permission on success, or $\bot$ if $\mathcal{H}$
is not a coin handle that identifies a coin.
Note that $\ora{Spend}$ can be used to generate deposit permissions that,
when deposited, would result in an error due to overspending
Adds $(\V{transactionId}, \V{depositPermission})$ to $\V{acceptedContracts}[\V{pkCustomer}]$.
\item $\ora{Share}(\mathcal{H}, \V{pkCustomer}) \mapsto ()$:
Shares a coin (identified by handle $\mathcal{H}$) with the customer
identified by $\V{pkCustomer}$, i.e. puts the coin identified by $\mathcal{H}$
into $\V{wallet}[\V{pkCustomer}]$. Intended to be used by the adversary in attempts to
violate income transparency.
Note that this trivially violates anonymity (by sharing with a corrupted customer), thus the usage must
be restricted in some games.
% the share oracle is the reason why we don't need a second withdraw oracle
\item $\ora{CorruptCustomer}(\V{pkCustomer})\mapsto
\newline{}\qquad (\V{skCustomer}, \V{wallet}[\V{pkCustomer}],\V{acceptedContracts}[\V{pkCustomer}],
\newline{}\qquad \phantom{(}\V{refreshIds}[\V{pkCustomer}], \V{withdrawIds}[\V{pkCustomer}])$:
Used by the adversary to corrupt a customer, giving the adversary access to
the customer's secret key, wallet, withdraw/refresh identifiers and accepted contracts.
Permanently marks the customer as corrupted. There is nothing ``special''
about corrupted customers, beyond that the adversary has used
\ora{CorruptCustomer} on them in the past. The adversary cannot modify
corrupted customer's wallets directly, and must use the oracle again to
obtain an updated view on the corrupted customer's private data.
\item $\ora{Deposit}(\V{depositPermission}) \mapsto \mathcal{I}$:
Triggers the execution of the \algo{Deposit} protocol, additionally giving
the adversary full control over the communication channels between merchant and exchange.
Returns an error if the deposit permission is addressed to a merchant that was not registered
with $\ora{AddMerchant}$.
This oracle does not give the adversary new information, but is used to
model the situation where there might be multiple conflicting deposit
permissions generated via $\algo{Spend}$, but only a limited number can be
deposited.
\end{itemize}
We write \oraSet{All} for the set of all the oracles we just defined.
We also let $\oraSet{NoShare} := \oraSet{All}  \{ \ora{Share} \}$
stand for access to all oracles except the share oracle.
The exchange does not need to be corrupted with an oracle. A corrupted exchange
is modeled by giving the adversary the appropriate oracles and the exchange
secret key from the exchange key generation.
If the adversary determines the exchange's secret key during the setup,
invoking \ora{WithdrawRequest}, \ora{WithdrawPickup}, \ora{RefreshRequest},
\ora{RefreshPickup} or \ora{Link} can be seen as the adversary playing the
exchange. Since the adversary is an active maninthemiddle in these oracles,
it can drop messages to the simulated exchange and make up its own response.
If the adversary calls these oracles with a corrupted customer, the adversary
plays as the customer.
%\begin{mdframed}
%The difference between algorithms and interactive protocols
%is that the ``pure'' algorithms only deal with data, while the interactive protocols
%take ``handles'' to parties that are communicating in the protocol. The adversary can
%always execute algorithms that don't depend on handles to communication partners.
%However the adversary can't run the interactive protocols directly, instead it must
%rely on the interaction oracles for it. Different interaction oracles might allow the
%adversary to play different roles in the same interactive protocol.
%
%While most algorithms in Taler are not probabilistic, we still say that they are, since
%somebody else might come up with an instantiation of Taler that uses probabilistic algorithms,
%and then the games should still apply.
%
%
%While we do have a \algo{Deposit} protocol that's used in some of the games, having a deposit oracle is not necessary
%since it does not give the adversary any additional power.
%\end{mdframed}
\section{Games}
We now define four security games (anonymity, conservation, unforgeability and
income transparency) that are later used to define security properties for
our protocol. Similar to \cite{bellare2006code} we assume that the game and adversary
terminate in finite time, and thus random choices made by the challenger and
adversary can be taken from a finite sample space.
All games except income transpacency return $1$ to indicate that the adversary
has won and $0$ to indicate that the adversary has lost. The income
transparency game returns $0$ if the adversary has lost, and a positive
``laundering ratio'' if the adversary won.
\subsection{Anonymity}
Intuitively, an adversary~$\prt{A}$ who controls the exchange and merchants wins the
anonymity game if they have a nonnegligible advantage in correlating spending operations
with the withdrawal or refresh operations that created a coin used in the
spending operation.
Let $b$ be the bit that will determine the mapping between customers and spend
operations, which the adversary must guess.
We define a helper procedure
\begin{equation*}
\algo{Refresh}(\prt{E}(\V{sksE}), \prt{C}(\V{pkCustomer}, \V{pksE}, \V{coin}_0)) \mapsto \mathfrak{R}
\end{equation*}
that refreshes the whole remaining amount on $\V{coin}_0$ with repeated application of $\algo{RefreshRequest}$
and $\algo{RefreshPickup}$ using the smallest possible set of target denominations, and returns all protocol transcripts
in $\mathfrak{R}$.
\begin{figure}
\fbox{\begin{minipage}{\textwidth}
\small
\noindent $\mathit{Exp}_{\prt{A}}^{anon}(1^\lambda, 1^\kappa, b)$:
\vspace{0.5\topsep}
\begin{enumerate}
\setlength\itemsep{0em}
\item $(\V{sksE}, \V{pksE}, \V{skM}, \V{pkM}) \leftarrow {\prt{A}}()$
\item $(\V{pkCustomer}_0, \V{pkCustomer}_1, \V{transactionId}_0, \V{transactionId}_1, f) \leftarrow {\prt{A}}^{\oraSet{NoShare}}()$
\item Select distinct fresh coins
\begin{align*}
\V{coin}_0 &\in \V{wallet}[\V{pkCustomer}_0]\\
\V{coin}_1 &\in \V{wallet}[\V{pkCustomer}_1]
\end{align*}
Return $0$ if either $\V{pkCustomer}_0$ or $\V{pkCustomer}_1$ are not registered customers with sufficient fresh coins,
or an oracle has been called with the coin handle for $\V{coin}_0$ or $\V{coin}_1$.
\item For $i \in \{0,1\}$ run
\begin{align*}
&\V{dp_i} \leftarrow \algo{Spend}(\V{transactionId}_i, f, \V{coin}_{ib}, \V{pkM}) \\
&\algo{Deposit}(\prt{A}(), \prt{M}(\V{skM}, \V{pksE}, \V{dp}_i)) \\
&\mathfrak{R}_i \leftarrow \algo{Refresh}(\prt{A}(), \prt{C}(\V{pkCustomer}_i, \V{pksE}, \V{coin}_{ib}))
\end{align*}
\item $b' \leftarrow {\cal A}^{\oraSet{NoShare}}(\mathfrak{R}_0, \mathfrak{R}_1)$ \\
\item Return $0$ if $\ora{Spend}$ was used by the adversary on the coin handles
for $\V{coin}_0$ or $\V{coin}_1$ or $\ora{CorruptCustomer}$ was used on $\V{pkCustomer}_0$ or $\V{pkCustomer}_1$.
\item If $b = b'$ return $1$, otherwise return $0$.
\end{enumerate}
\end{minipage}}
\end{figure}
Note that unlike other anonymity games defined in the literature (such as
\cite{pointcheval2017cut}), our anonymity game always lets both customers spend
in order to avoid having to hide the missing coin in one customer's wallet
from the adversary.
\subsection{Conservation}
The adversary wins the conservation game if it can bring an honest customer in a
situation where the spendable financial value left in the user's wallet plus
the value spent for transactions known to the customer is less than the value
withdrawn by the same customer through by the exchange.
In practice, this property is necessary to guarantee that aborted or partially
completed withdrawals, payments or refreshes, as well as other (transient)
misbehavior from the exchange or merchant do not result in the customer losing
money or privacy.
\begin{figure}
\fbox{\begin{minipage}{\textwidth}
\small
\noindent $\mathit{Exp}_{\cal A}^{conserv}(1^\lambda, 1^\kappa)$:
\vspace{0.5\topsep}
\begin{enumerate}
\setlength\itemsep{0em}
\item $(\V{sksE}, \V{pksE}) \leftarrow \mathrm{ExchangeKeygen}(1^\lambda, 1^\kappa, M)$
\item $\V{pkCustomer} \leftarrow {\cal A}^{\oraSet{NoShare}}(\V{pksE})$
\item Return $0$ if $\V{pkCustomer}$ is a corrupted user.
\item \label{game:conserv:run} Run $\algo{WithdrawPickup}$ for each withdraw identifier $\V{wid}$
and $\algo{RefreshPickup}$ for each refresh identifier $\V{rid}$ that the user
has recorded in $\V{withdrawIds}$ and $\V{refreshIds}$. Run $\algo{Deposit}$
for all deposit permissions in $\V{acceptedContracts}$.
\item Let $v_{C}$ be the total financial value left on valid coins in $\V{wallet}[\V{pkCustomer}]$,
i.e. the denominated values minus the spend/refresh operations recorded in the exchange's database.
Let $v_{S}$ be the total financial value of contracts in $\V{acceptedContracts}[\V{pkCustomer}]$.
\item Return $1$ if $\V{withdrawn}[\V{pkCustomer}] > v_{C} + v_{S}$.
\end{enumerate}
\end{minipage}}
\end{figure}
Hence we ensure that:
\begin{itemize}
\item if a coin was spent, it was spent for a contract that the customer
knows about, i.e. in practice the customer could prove that they ``own'' what they
paid for,
\item if a coin was refreshed, the customer ``owns'' the resulting coins,
even if the operation was aborted, and
\item if the customer withdraws, they can always obtain a coin whenever the
exchange accounted for a withdrawal, even when protocol executions are
intermittently aborted.
\end{itemize}
We do not give the adversary access to the \ora{Share} oracle,
which naively lets them win the conservation game, because doing so
correctly requires tracking the corrupted customers more carefully,
which harms our exposition.
In practice, conservation only holds for customers that do not share
coins with parties that they do not fully trust.
\subsection{Unforgeability}
Intuitively, adversarial customers win if they can obtain more valid coins than
they legitimately withdraw.
\begin{figure}
\fbox{\begin{minipage}{\textwidth}
\small
\noindent $\mathit{Exp}_{\cal A}^{forge}(1^\lambda, 1^\kappa)$:
\vspace{0.5\topsep}
\begin{enumerate}
\setlength\itemsep{0em}
\item $(skE, pkE) \leftarrow \mathrm{ExchangeKeygen}()$
\item $(C_0, \dots, C_\ell) \leftarrow \mathcal{A}^{\oraSet{All}}(pkExchange)$
\item Return $0$ if any $C_i$ is not of the form $(\V{skCoin}, \V{pkCoin}, \V{pkD}, \V{coinCert})$
or any $\V{coinCert}$ is not a valid signature by $\V{pkD}$ on the respective $\V{pkCoin}$.
\item Return $1$ if the sum of the unspent value of valid coins in $C_0
\dots, C_\ell$ exceeds the amount withdrawn by corrupted
customers, return $0$ otherwise.
\end{enumerate}
\end{minipage}}
\end{figure}
\subsection{Income Transparency}
Intuitively, the adversary wins if coins are in exclusive control of corrupted
customers, but the exchange has no record of withdrawal or spending for them.
This presumes that the adversary cannot delete from noncorrupted customer's
wallets, even though it can use oracles to force protocol interactions of
noncorrupted customers.
For practical ecash systems, income transparency disincentivizes the emergence
of ``black markets'' among mutually distrusting customers, where currency
circulates without the transactions being visible. This is in contrast to some
other proposed ecash systems and cryptocurrencies, where disintermediation is
an explicit goal. The Link protocol introduces the threat of losing exclusive
control of coins (despite having the option to refresh them) that were received
without being visible as income to the exchange.
\begin{figure}
\fbox{\begin{minipage}{\textwidth}
\small
\noindent $\mathit{Exp}_{\cal A}^{income}(1^\lambda, 1^\kappa)$:
\vspace{0.5\topsep}
\begin{enumerate}
\setlength\itemsep{0em}
\item $(skE, pkE) \leftarrow \mathrm{ExchangeKeygen}()$
\item $(\V{coin}_1, \dots, \V{coin}_\ell) \leftarrow \mathcal{A}^{\oraSet{All}}(pkExchange)$
(The $\V{coin}_i$ must be coins, including secret key and signature by the
denomination, for the adversary to win. However these coins need not be
present in any honest or corrupted customer's wallet.)
\item\label{game:income:spend} Augment the wallets of all noncorrupted customers with their
transitive closure using the \algo{Link} protocol.
Mark all remaining value on coins in wallets of noncorrupted customers as
spent in the exchange's database.
\item Let $L$ denote the sum of unspent value on valid coins in $(\V{coin}_1, \dots\, \V{coin}_\ell)$,
after accounting for the previous update of the exchange's database.
Also let $w'$ be the sum of coins withdrawn by corrupted customers.
Then $p := L  w'$ gives the adversary's untaxed income.
\item Let $w$ be the sum of coins withdrawn by noncorrupted customers, and
$s$ be the value marked as spent by noncorrupted customers, so that
$b := w  s$ gives the coins lost during refresh, that is the losses incurred attempting to hide income.
\item If $b+p \ne 0$, return $p \over b + p$, i.e. the laundering ratio for attempting to obtain untaxed income. Otherwise return $0$.
\end{enumerate}
\end{minipage}}
\end{figure}
\section{Security Definitions}\label{sec:securityproperties}
We now give security definitions based upon the games defined in the previous
section.
\begin{definition}[Anonymity]
We say that an ecash scheme satisfies \emph{anonymity} if the success
probability $\Prb{b \randsel \{0,1\}: \mathit{Exp}_{\cal A}^{anon}(1^\lambda,
1^\kappa, b) = 1}$ of the anonymity game is negligibly close to $2^{\lambda}$ for any
polynomial time adversary~$\mathcal{A}$.
\end{definition}
\begin{definition}[Conservation]
We say that an ecash scheme satisfies \emph{conservation} if
the success probability $\Prb{\mathit{Exp}_{\cal A}^{conserv}(1^\lambda, 1^\kappa) = 1}$
of the conservation game is negligible for any polynomial time adversary~$\mathcal{A}$.
\end{definition}
\begin{definition}[Unforgeability]
We say that an ecash scheme satisfies \emph{unforgeability} if the success
probability $\Prb{\mathit{Exp}_{\cal A}^{forge}(1^\lambda, 1^\kappa) = 1}$ of
the unforgeability game is negligible for any polynomial time adversary
$\mathcal{A}$.
\end{definition}
\begin{definition}[Weak Income Transparency]
We say that an ecash scheme satisfies \emph{weak income transparency}
if, for any polynomial time adversary~$\mathcal{A}$,
the return value of the the income transparency game satisfies
\begin{equation}\label{eq:incometransparencyexpectation}
E\left[\mathit{Exp}_{\cal A}^{income}(1^\lambda, 1^\kappa)\right] \le {1\over\kappa} \mathperiod
\end{equation}
In (\ref{eq:incometransparencyexpectation}), the expectation runs over
any probability space used by the adversary and challenger.
\end{definition}
For some instantiations, e.g. ones based on zero knowledge proofs, $\kappa$
might be a security parameter in the traditional sense. However for an ecash
scheme to be useful in practice, the adversary need not have only
negligible success probability in the income transparency game.
It suffices that the financial losses of the adversary in the game are a
deterrent, after all our purpose of the game is to characterize tax evasion.
\section{Instantiation}
We give an instantiation of our protocol syntax that is generic over
a blind signature scheme, a signature scheme, a combined signature scheme / key
exchange, a commitment scheme and a pseudorandom function family (PRF).
%\subsection{Generic Instantiation}
Let $\textsc{BlindSign}$ be a blind signature scheme with the following syntax, where the party $\mathcal{S}$
is the signer and $\mathcal{R}$ is the signature requester:
\begin{itemize}
\item $\algo{KeyGen}_{BS}(1^\lambda) \mapsto (\V{sk}, \V{pk})$ is the key generation algorithm
for the signer of the blind signature protocol.
\item $\algo{Blind}_{BS}(\mathcal{S}(\V{sk}), \mathcal{R}(\V{pk}, m)) \mapsto (\overline{m}, r)$ is a possibly interactive protocol
to blind a message $m$ that is to be signed later. The result is a blinded message $\overline{m}$ and
a residual $r$ that allows to unblind a blinded signature on $m$ made by $\V{sk}$.
\item $\algo{Sign}_{BS}(\mathcal{S}(\V{sk}), \mathcal{R}(\overline{m})) \mapsto
\overline{\sigma}$ is an algorithm to sign a blinded message $\overline{m}$.
The result $\overline{\sigma}$ is a blinded signature that must be unblinded
using the $r$ returned from the corresponding blinding operation before
verification. We restrict \algo{Sign} to be a twomove protocol, where the
requester sends the first message, and the signer responds.
\item $\algo{UnblindSig}_{BS}(r, m, \overline{\sigma}) \mapsto \sigma$
is an algorithm to unblind a blinded signature.
\item $\algo{Verify}_{BS}(\V{pk}, m, \sigma) \mapsto b$ is a algorithm to check the validity of a blind
signature. Returns $1$ if the signature $\sigma$ was valid for $m$ and $0$ otherwise.
\end{itemize}
Note that this syntax excludes some blind signature protocols, such as those
with interactive/probabilistic verification or those without a ``blinding
factor'', where the $\algo{Blind}_{BS}$ and $\algo{Sign}_{BS}$ and
$\algo{UnblindSig}_{BS}$ would be merged into one interactive signing protocol.
Such blind signature protocols have already been used to construct ecash
\cite{camenisch2005compact}.
We require the following two security properties for $\textsc{BlindSign}$:
\begin{itemize}
\item \emph{blindness}: Let $M$ be the set of all possible messages and $\overline{M}$ be the
set of all possible blinded messages. Then the distribution of
\[ \left\{ (m, \sigma, \overline{m}, \overline{\sigma}) \,\middle
\begin{array}{c}
m\, \randsel M, \\
\overline{m} \leftarrow \algo{Blind}_{BS}(\mathcal{S}(\V{sk}), m), \\
\overline{\sigma} \leftarrow \algo{Sign}_{BS}(\V{sk}, \overline{m}), \\
\sigma \leftarrow \algo{UnblindSig}_{BS}(r, m, \overline{\sigma})
\end{array}
\right\} \]
must be computationally
indistinguishable from
\[ \left\{ (m, \sigma, x, \sigma_x) \,\middle\,
\begin{array}{c}
m \randsel M, \\
\sigma \leftarrow \algo{UnblindSig}_{BS}(r, m, \algo{Sign}_{BS}(\V{sk}, \algo{Blind}_{BS}(\mathcal{S}(\V{sk}), m)) ) \\
x \randsel \overline{M}, \\
\sigma_x \leftarrow \algo{UnblindSig}_{BS}(r, x, \algo{Sign}_{BS}(\V{sk}, \algo{Blind}_{BS}(\mathcal{S}(\V{sk}), x)) )
\end{array}
\right\}. \]
\item \emph{unforgeability}: An adversary that requests $k$ signatures with $\algo{Sign}_{BS}$
is unable to produce $k+1$ valid signatures with nonnegligible probability.
\end{itemize}
For more generalized notions of the security of blind signatures, see e.g.
\cite{fischlin2009security,schroder2017security}.
Let $\textsc{CoinSignKx}$ be combination of a signature scheme and key exchange:
\begin{itemize}
\item $\algo{KeyGenSec}_{CSK}(1^\lambda) \mapsto \V{sk}$ is a secret key generation algorithm.
\item $\algo{KeyGenPub}_{CSK}(\V{sk}) \mapsto \V{pk}$ produces the corresponding public key.
\item $\algo{Sign}_{CSK}(\V{sk}, m) \mapsto \sigma$ produces a signature $\sigma$ over message $m$.
\item $\algo{Verify}_{CSK}(\V{pk}, m, \sigma) \mapsto b$ is a signature verification algorithm.
Returns $1$ if the signature $\sigma$ is a valid signature on $m$ by $\V{pk}$, and $0$ otherwise.
\item $\algo{Kx}_{CSK}(\V{sk}_1, \V{pk}_2) \mapsto x$ is a key exchange algorithm that computes
the shared secret $x$ from secret key $\V{sk}_1$ and public key $\V{pk}_2$.
\end{itemize}
We occasionally need these key generation algorithms separately, but
we usually combine them into $\algo{KeyGen}_{CSK}(1^\lambda) \mapsto (\V{sk}, \V{pk})$.
We require the following security properties to hold for $\textsc{CoinSignKx}$:
\begin{itemize}
\item \emph{unforgeability}: The signature scheme $(\algo{KeyGen}_{CSK}, \algo{Sign}_{CSK}, \algo{Verify}_{CSK})$
must satisfy existential unforgeability under chosen message attacks (EUFCMA).
\item \emph{honest key generation}:
Any probabilistic polynomialtime adversary has only negligible chance
to produce a public key $\V{pk}$ and a valid signature $\sigma$,
i.e.\ $\algo{Verify}_{CSK}(\V{pk}, m, \sigma) = 1$, without also producing
some secret key $\V{sk}$ such that $\V{pk} = \algo{KeyGenPub}_{CSK}(\V{sk})$.
\item \emph{key exchange completeness}:
Any probabilistic polynomialtime adversary has only negligible chance find
$(\V{sk}_x, \V{pk}_x) \leftarrow \algo{KeyGen}_{CSK}(1^\lambda)$ for $x=A,B$
for which the key exchange fails,
\begin{equation*}
\algo{Kex}_{CSK}(\V{sk}_A, \V{pk}_B) \neq \algo{Kex}_{CSK}(\V{sk}_B, \V{pk}_A).
\end{equation*}
\item \emph{key exchange security}: The output of $\algo{Kx}_{CSK}$ must be computationally
indistinguishable from a random shared secret of the same length, for inputs that have been
generated with $\algo{KeyGen}$.
\end{itemize}
Let $\textsc{Sign} = (\algo{KeyGen}_{S}, \algo{Sign}_{S}, \algo{Verify}_{S})$ be a signature
scheme that satisfies SUFCMA.
Let $(\algo{Setup}_C, H_{pk})$ be a computationally hiding and binding
commitment scheme, where $\algo{Setup}$ generates the public commitment key
$pk$ and $H_{pk} : \{0,1\}^* \rightarrow \{0,1\}^\lambda$ deterministically commits to a
bitstring.
Let $\V{PRF}$ be a pseudorandom function family.
Using these primitive, we now instantiate the syntax:
\begin{itemize}
\item $\algo{ExchangeKeygen}(1^{\lambda}, 1^{\kappa}, \mathfrak{D})$:
Generate the exchange's signing key pair $\V{skESign} \leftarrow \algo{KeyGen}_{S}(1^\lambda)$ and public
commitment key $\V{CK} \leftarrow Setup_C(1^\lambda)$.
For each element in the sequence $\mathfrak{D} = d_1,\dots,d_n$, generate
denomination key pair $(\V{skD}_i, \V{pkD}_i) \leftarrow \algo{KeyGen}_{BS}(1^\lambda)$.
\item $\algo{CustomerKeygen}(1^\lambda,1^\kappa)$:
Return key pair $\algo{KeyGen}_S(1^\lambda)$.
\item $\algo{MerchantKeygen}(1^\lambda,1^\kappa)$:
Return key pair $\algo{KeyGen}_S(1^\lambda)$.
\item $\algo{WithdrawRequest}(\prt{E}(\V{sksE}, \V{pkCustomer}), \prt{C}(\V{skCustomer}, \V{pksE}, \V{pkD}))$:
Let $\V{skD}$ be the exchange's denomination secret key corresponding to $\V{pkD}$.
\begin{enumerate}
\item \prt{C} generates coin key pair $(\V{skCoin}, \V{pkCoin}) \leftarrow \algo{KeyGen}_{CSK}(1^\lambda)$
\item \prt{C} runs $(\overline{m}, r) \leftarrow \algo{Blind}_{CSK}(\mathcal{E}(\V{skCoin}), \mathcal{C}(m))$ with the exchange
\end{enumerate}
The withdraw identifier is then
\begin{equation*}
\V{wid} := (\V{skCoin}, \V{pkCoin}, \overline{m}, r)
\end{equation*}
\item $\algo{WithdrawPickup}(\prt{E}(\V{sksE}, \V{pkCustomer}), \prt{C}(\V{skCustomer}, \V{pksE}, \V{wid}))$:
The customer looks up $\V{skCoin}$, $\V{pkCoin}$, \V{pkD} $\overline{m}$
and $r$ via the withdraw identifier $\V{wid}$.
\begin{enumerate}
\item \prt{C} runs $\overline{\sigma} \leftarrow \algo{Sign}_{BS}(\mathcal{E}(\V{skD}), \mathcal{C}(\overline{m}))$ with the exchange
\item \prt{C} unblinds the signature $\sigma \leftarrow \algo{UnblindSig}_{BS}(\overline{\sigma}, r, \overline{m})$
and stores the coin $(\V{skCoin}, \V{pkCoin}, \V{pkD}, \sigma)$ in their wallet.
\end{enumerate}
\item $\algo{Spend}(\V{transactionId}, f, \V{coin}, \V{pkMerchant})$:
The deposit permission is computed as $\V{depositPermission} = (\V{pkCoin},
\sigma_D, m)$ where $m = (\V{transactionId}, f, \V{pkMerchant})$ and $\sigma_D
= \algo{Sign}_{CSK}(skCoin, m)$.
\item $\algo{Deposit}(\prt{E}(\V{sksE}, \V{pkMerchant}), \prt{M}(\V{skMerchant}, \V{pksE}, \V{depositPermission}))$:
Exchange checks the signature $\sigma_D$ in the deposit permission, records
contribution $f$ in database $\V{deposited}$ of deposited coins.
\item $\algo{RefreshRequest}(\prt{E}(\V{sksE}, \V{pkCustomer}), \prt{C}(\V{skCustomer}, \V{pksE}, \V{coin}_0, \V{pkD}_u))$:
Let $\V{skD}_u$ be the secret key corresponding to $\V{pkD}_u$.
We write \[ \algo{Blind}^*_{BS}(\mathcal{S}(\V{sk}), \mathcal{R}(R,
\V{sk}_C, \V{pk}, m)) \mapsto (\overline{m}, r, \mathcal{T}_{B*}) \] for a
modified version of $\algo{Blind}_{BS}$ where the signature requester
$\mathcal{R}$ takes all randomness from the sequence
$\left(\V{PRF}(R,\texttt{"blind"}\Vert n)\right)_{n>0}$, the messages from the exchange are
recorded in transcript $\mathcal{T}_{B*}$, and all messages sent by the
requester are signed with $\V{sk}_C$.
Furthermore we write \[ \algo{KeyGen}^*_{CSK}(R, 1^\lambda) \mapsto
(\V{sk}, \V{pk}) \] for a modified version of the key generation algorithm
that takes randomness from the sequence $\left(\V{PRF}(R,\texttt{"key"}\Vert
n)\right)_{n>0}$.
For each $i\in \{1,\dots,\kappa \}$, the customer
\begin{enumerate}
\item generates seed $s_i \randsel \{1, \dots, 1^\lambda\}$
\item generates transfer key pair $(t_i, T_i) \leftarrow \algo{KeyGen}^*_{CSK}(s_i, 1^\lambda)$
\item computes transfer secret $x_i \leftarrow \algo{Kx}(t_i, \V{pkCoin}_0)$
\item computes coin key pair $(\V{skCoin}_i, \V{pkCoin}_i) \leftarrow
\algo{KeyGen}^*_{CSK}(x_i, 1^\lambda)$
\item and executes the modified blinding protocol
\[
(\overline{m}_i, r_i, \mathcal{T}_{(B*,i)}) \leftarrow
\algo{Blind}^*_{BS}(\mathcal{E}(\V{skD_u}), \mathcal{C}(x_i, \V{skCoin}_0, \V{pkD}_u, \V{skCoin}_i))
\]
with the exchange.
\end{enumerate}
The customer stores the refresh identifier
\begin{equation}
\V{rid} := (\V{coin}_0, \V{pkD}_u, \V{nonce}, \{ s_i \}, \{ \overline{m}_i \}, \{r_i\}, \{\mathcal{T}_{(B*,i)}\} ).
\end{equation}
Now, the customer's wallet sends the commitment $\pi_1 = (\V{pkCoin}_0, \V{pkD}_u, h_C)$ together with signature $\V{sig}_1
\leftarrow \algo{Sign}_{CSK}(\V{skCoin}_0, \pi_1)$ to the exchange, where
\begin{align*}
h_T &:= H_{pk}(T_1, \dots, T_\kappa)\\
h_{\overline{m}} &:= G_{pk}(\overline{m}_1, \dots, \overline{m}_\kappa)\\
h_C &:= H_{pk}((h_T \Vert h_{\overline{m}})
\end{align*}
The exchange checks the signature $\V{sig}_1$, and aborts if invalid. Otherwise,
depending on the commitment:
\begin{enumerate}
\item if the exchange did not see $\pi_1$ before, it marks $\V{pkCoin}_0$
as spent for $D(\V{pkD}_u)$, chooses a uniform random $0 \le \gamma < \kappa$, stores it,
and sends this choice in a signed message $(\gamma, \V{sig}_2)$ to the customer,
where $\V{sig}_2 \leftarrow \algo{Sign}_{S}(\V{skESig}, \gamma)$.
\item otherwise, the exchange sends back the same $\pi_2$ as it sent for the last
equivalent $\pi_1$.
\end{enumerate}
\item $\algo{RefreshPickup}(\prt{E}(\V{sksE}, \V{pkCustomer}), \prt{C}(\V{skCustomer}, \V{pksE}, \V{rid}) \rightarrow \mathcal{T}$:
The customer's wallet looks up the refresh identifier $\V{rid}$ and recomputes the transfer key pairs,
transfer secrets and new coin key pairs. The customer sends the reveal message
\begin{equation*}
\pi_3 = T_\gamma, \overline{m}_\gamma,
(s_1, \dots, s_{\gamma1}, s_{\gamma+1}, \dots, s_\kappa)
\end{equation*}
and signature
\begin{equation*}
\V{sig}_{3'} \leftarrow \algo{Sign}_{CSK}(\V{skCoin}_0, (\V{pkCoin}_0,
\V{pkD}_u, \mathcal{T}_{(B*,\gamma)}, T_\gamma, \overline{m}_\gamma))
\end{equation*} to the exchange.
The exchange checks the signature $\V{sig}_{3'}$ and then computes for $i \ne \gamma$:
\begin{align*}
(t_i', T_i') &\leftarrow \algo{KeyGen}^*_{CSK}(s_i, 1^\lambda)\\
x_i' &\leftarrow \algo{Kx}(t_i, \V{pkCoin}_0)\\
(\V{skCoin}_i', \V{pkCoin}_i') &\leftarrow
\algo{KeyGen}^*_{CSK}(x_i', 1^\lambda) \\
h_T' &:= H(T'_1, \dots, T_{\gamma1}, T_\gamma, T_{\gamma+1}', \dots, T_\kappa')
\end{align*}
and simulates the blinding protocol with recorded transcripts (without signing each message,
as indicated by the dot ($\cdot$) instead of a signing secret key), obtaining
\begin{align*}
(\overline{m}_i', r_i', \mathcal{T}_i) &\leftarrow
\algo{Blind}^*_{BS}(\mathcal{S}(\V{skD}_u), \mathcal{R}(H(x_i'), \cdot, \V{pkD}_u, \V{skCoin}_i))\\
\end{align*}
and finally
\begin{align*}
h_{\overline{m}}' &:= H(\overline{m}_1', \dots, \overline{m}_\gamma, \dots, \overline{m}_\kappa')\\
h_C &:= H(h_T' \Vert h_{\overline{m}}').
\end{align*}
For each $i \ne \gamma$, the exchange computes
\begin{align*}
\overline{\sigma}_i' &\leftarrow \algo{Sign}(\mathcal{E}(\V{skD}_u), \mathcal{E}(\overline{m}_i'))\\
\sigma_i' &\leftarrow \algo{UnblindSig}(r_i', \V{pkCoin}_i', \overline{\sigma}_i')\\
b_i &\leftarrow \algo{Verify}_{BS}(\V{pkD}, \V{skCoin}_i', \sigma_i')
\end{align*}
Now the exchange checks if $h_C = h_C'$ and if all $b_i = 1$ for $i \ne \gamma$.
If one of the checks fails, the exchange aborts the protocol.
Otherwise, the exchange sends a message back to $\prt{C}$ that the commitment verification succeeded.
As a last step, the customer obtains the signature $\sigma_\gamma$ on the new coin's public key $\V{pkCoin}_u$ with
\begin{align*}
\overline{\sigma}_\gamma &\leftarrow \algo{Sign}(\mathcal{E}(\V{skD}_u), \mathcal{C}(\overline{m}_\gamma))\\
\sigma_\gamma &\leftarrow \algo{UnblindSig}(r_\gamma, \V{pkCoin}_\gamma, \overline{\sigma}_\gamma).
\end{align*}
Thus the new, unlinkable coin is $\V{coin}_u = (\V{skCoin}_\gamma, \V{pkCoin}_\gamma, \V{pkD}_u, \sigma_\gamma)$.
\item $\algo{Link}(\prt{E}(\V{sksE}), \prt{C}(\V{skCustomer}, \V{pksE}, \V{coin}_0))$:
The customer sends the public key $\V{pkCoin}_0$ of $\V{coin}_0$ to the exchange.
For each completed refresh on $\V{pkCoin}_0$ recorded in the exchange's
database, the exchange sends the following data back to the customer: the
signed commit message $(\V{sig}_1, \pi_1)$, the transfer public key
$T_\gamma$, the signature $\V{sig}_{3'}$, the blinded signature $\overline{\sigma}_\gamma$, and the
transcript $\mathcal{T}_{(B*,\gamma)}$ of the customer's and exchange's messages
during the \algo{Blind} protocol execution.
The following logic is repeated by the customer for each response:
\begin{enumerate}
\item Verify the signatures (both from $\V{pkESig}$ and $\V{pkCoin}_0$) on the transcript $\mathcal{T}_\kappa$,
abort otherwise.
\item Verify that $\V{sig}_1$ is a valid signature on $\pi_1$ by $\V{pkCoin}_0$, abort otherwise.
\item Recompute the transfer secret and the new coin's key pair as
\begin{align*}
x_\gamma &\leftarrow \algo{Kx}_{CSK}(\V{skCoin}_0, T_\gamma)\\
(\V{skCoin}_\gamma, \V{pkCoin}_\gamma) &\leftarrow \algo{KeyGen}_{CSK}^*(x_\gamma, 1^\lambda).
\end{align*}
\item Simulate the blinding protocol with the message transcript received from the exchange to obtain
$(\overline{m}_\gamma, r_\gamma)$.
\item Check that $\algo{Verify}_{CSK}(\V{pkCoin}_0,
\V{pkD}_u, \V{skCoin}_0,(\mathcal{T}_{(B*,\gamma)}, \overline{m}_\gamma), \V{sig}_{3'})$
indicates a valid signature, abort otherwise.
\item Unblind the signature to obtain $\sigma_\gamma \leftarrow \algo{UnblindSig}(r_\gamma, \V{pkCoin}_\gamma, \overline{\sigma}_\gamma)$
\item (Re)add the coin $(\V{skCoin}_\gamma, \V{pkCoin}_\gamma, \V{pkD}_u, \sigma_\gamma)$ to the customer's wallet.
\end{enumerate}
\end{itemize}
%In Taler's refresh, we avoid key exchange failures entirely because the
%Edwards addition law is complete abelian group operation on the curve,
%and the signature scheme verifies that the point lies on the curve.
%% https://safecurves.cr.yp.to/refs.html#2007/bernsteinnewelliptic
%% https://safecurves.cr.yp.to/complete.html
%We warn however that Weierstrass curves have incomplete addition formulas
%that permit an adversarial merchant to pick transfer keys that yields failures.
%There are further implementation mistakes that might enable collaborative
%key exchange failures, like if the exchange does not enforce the transfer
%private key being a multiple of the cofactor.
%
%In this vein, almost all postquantum key exchanges suffer from key exchange
%failures that permit invalid key attacks against nonephemeral keys.
%All these schemes support only one ephemeral party by revealing the
%ephemeral party's private key to the nonephemeral party,
% ala the FujisakiOkamoto transform~\cite{fujisakiokamoto} or similar.
%We cannot reveal the old coin's private key to the exchange when
%verifying the transfer private keys though, which
% complicates verifying honest key generation of the old coin's key.
\bibliographystyle{splncs04}
\bibliography{ref}
\appendix
\section{Proofs}
%\begin{mdframed}
% Currently the proofs don't have any explicit tightess bounds.
% Because we don't know where to ``inject'' the value that we get from the challenger when carrying out
% a reduction, we need to randomly guess in which coin/signature we should ``hijack'' our challenge value.
% Thus for the proofs to work fully formally, we need to bound the total number of oracle invocations,
% and our exact bound for the tightness of the reduction depends on this limit.
%\end{mdframed}
We now give proofs for the security properties defined in Section \ref{sec:securityproperties}
with the generic instantiation.
\subsection{Anonymity}
\begin{theorem}
Our instantiation satisfies anonymity.
\end{theorem}
\begin{proof}
We give a proof via a sequence of games $\mathbb{G}_0(b), \mathbb{G}_1(b),
\mathbb{G}_2(b)$, where $\mathbb{G}_0(b)$ is the original anonymity game
$\mathit{Exp}_{\cal A}^{anon}(1^\lambda, 1^\kappa, b)$. We show the
adversary can distinguish between subsequent games with only negligible
probability. Let $\epsilon_{HC}$ and $\epsilon_{KX}$ be the advantage of an
adversary for finding hash collisions and for breaking the security of the
key exchange respectively.
We define $\mathbb{G}_1$ by replacing the link oracle \ora{Link} with a
modified version that behaves the same as \ora{Link}, unless the adversary
responds with link data that did not occur in the transcript of a successful
refresh operation, but despite of that still passes the customer's
verification. In that case, the game is aborted instead.
Observe that in case this failure event happens, the adversary must have forged a
signature on $\V{sig}_{3'}$ on values not signed by the customer, yielding
an existential forgery. Thus $\left \Prb{\mathbb{G}_0 = 1}  \Prb{\mathbb{G}_1 = 1}
\right$ is negligible.
In $\mathbb{G}_2$, the refresh oracle is modified so that the customer
responds with value drawn from a uniform random distribution $D_1$ for the the
$\gamma$th commitment instead of using the key exchange function. We must
handle the fact that $\gamma$ is chosen by the adversary after seeing the
commitments, so the challenger first makes a guess $\gamma^*$ and replaces
only the $\gamma^*$th commitment with a uniform random value. If the
$\gamma$ chosen by the adversary does not match $\gamma^*$, then the
challenger rewinds \prt{A} to the point where the refresh oracle was called.
Note that we only replace the one commitment that
will not be opened to the adversary later.
Since $\kappa \ll \lambda$ and the security property of $\algo{Kx}$
guarantees that the adversary cannot distinguish the result of a key exchange
from randomness, the runtime complexity of the challenger still stays
polynomial in $\lambda$. An adversary that could with high probability
choose a $\gamma$ that would cause a rewind, could also distinguish
randomness from the output of $\algo{Kx}$.
%\mycomment{Tighness bound also missing}
We now show that $\left \Prb{\mathbb{G}_1 = 1}  \Prb{\mathbb{G}_2 = 1}
\right \le \epsilon_{KX}$ by defining a distinguishing game $\mathbb{G}_{1
\sim 2}$ for the key exchange as follows:
\bigskip
\noindent $\mathbb{G}_{1 \sim 2}(b)$:
\vspace{0.5\topsep}
\begin{enumerate}
\setlength\itemsep{0em}
\item If $b=0$, set
\[
D_0 := \{ (A, B, \V{Kex}(a, B)) \mid (a, A) \leftarrow \V{KeyGen}(1^\lambda),(b, B) \leftarrow \V{KeyGen}(1^\lambda) \}.
\]
Otherwise, set
\[
D_1 := \{ (A, B, C) \mid (a, A) \leftarrow \V{KeyGen}(1^\lambda),
(b, B) \leftarrow \V{KeyGen}(1^\lambda),
C \randsel \{1,\dots,2^\lambda\} \}.
\]
\item Return $\mathit{Exp'}_{\cal A}^{anon}(b, D_b)$
(Modified anonymity game where the $\gamma$th commitment in the
refresh oracle is drawn uniformly from $D_b$ (using rewinding). Technically, we need to
draw from $D_b$ on withdraw for the coin secret key, write it to a table, look it up on refresh and
use the matching tuple, so that with $b=0$ we perfectly simulate $\mathbb{G}_1$.)
\end{enumerate}
Depending on the coin flip $b$, we either simulate
$\mathbb{G}_1$ or $\mathbb{G}_2$ perfectly for our adversary~$\mathcal{A}$
against $\mathbb{G}_1$. At the same time $b$ determines whether \prt{A}
receives the result of the key exchange or real randomness. Thus $\left
\Prb{\mathbb{G}_1 = 1}  \Prb{\mathbb{G}_2 = 1} \right = \epsilon_{KX}$ is
exactly the advantage of $\mathbb{G}_{1 \sim 2}$.
We observe in $\mathbb{G}_2$ that as $x_\gamma$ is uniform random and not
learned by the adversary, the generation of $(\V{skCoin}_\gamma,
\V{pkCoin}_\gamma)$ and the execution of the blinding protocol is equivalent (under the PRF assumption)
to using the randomized algorithms
$\algo{KeyGen}_{CSK}$ and $\algo{Blind}_{BS}$.
By the blindness of the $\textsc{BlindSign}$ scheme, the adversary is not
able to distinguish blinded values from randomness. Thus, the adversary is
unable to correlate a $\algo{Sign}_{BS}$ operation in refresh or withdraw
with the unblinded value observed during $\algo{Deposit}$.
We conclude the success probability for $\mathbb{G}_2$ must be $1/2$ and
hence the success probability for $\mathit{Exp}_{\cal A}^{anon}(1^\lambda,
\kappa, b)$ is at most $1/2 + \epsilon(\lambda)$, where $\epsilon$ is a
negligible function.
\end{proof}
% RSA ratios vs CDH in BLS below
\subsection{Conservation}
\begin{theorem}
Assuming existential unforgeability of signatures (EUFCMA), our instantiation satisfies conservation.
\end{theorem}
\begin{proof}
% FIXME: argue that reduction is tight when you have malleability
In honest executions, we have $\V{withdrawn}[\V{pkCustomer}] = v_C + v_S$, i.e.
the coins withdrawn add up to the coins still available and the coins spent
for known transactions.
In order to win the conservation game, the adversary must increase
$\V{withdrawn}[\V{pkCustomer}]$ or decrease $v_C$ or $v_S$. An adversary can
abort withdraw operations, thus causing $\V{withdrawn}[\V{pkCustomer}]$ to increase,
while the customer does not obtain any coins. However, in step
\ref{game:conserv:run}, the customer obtains coins from interrupted withdraw
operations. Similarly for the refresh protocol, aborted \algo{RefreshRequest} / \algo{RefreshPickup}
operations that result in a coin's remaining value being reduced are completed
in step \ref{game:conserv:run}.
Thus only remaining option for the adversary to decrease $v_C$ or $v_S$ is
with the $\ora{RefreshPickup}$ and $\ora{Deposit}$ oracles respectively.
Since the exchange verifies signatures made by the secret key of the coin
that is being spent/refreshed, the adversary must forge this signature or have
access to the coin's secret key. As we do not give the adversary access to
the sharing oracle, it does not have direct access to any of the honest
customer's coin secret keys.
Thus the adversary must either compute the coin's secret key from observing
the coin's public key (e.g. during a partial deposit operation), or forge
signatures directly. Both possibilities allow us to carry out a reduction
against the unforgeability property of the $\textsc{CoinSignKx}$ scheme, by
injecting the challenger's public key into one of the coins.
\end{proof}
\subsection{Unforgeability}
\begin{theorem}
Our instantiation satisfies {unforgeability}.
% by probabilistic polynomially time adversaries.
\end{theorem}
\begin{proof}
The adversary must have produced at least one coin that was not blindly
signed by the exchange. %TODO: Way too fasty here, resurect the chain
In order to carry out a reduction from this adversary to a blind signature
forgery, we inject the challenger's public key into one randomly chosen
denomination. Since we do not have access to the corresponding secret key
of the challenger, signing operations for this denomination are replaced
with calls to the challenger's signing oracle in \ora{WithdrawPickup} and
\ora{RefreshPickup}. For $n$ denominations, an adversary against the
unforgeability game would produce a blind signature forgery with probability $1/n$.
\end{proof}
%TODO: RSAKTI
\subsection{Income Transparency}
\begin{theorem}
Our instantiation satisfies {weak income transparency}.
\end{theorem}
\begin{proof}
In our refresh operation, the commitment phase sends only the hash
of blinded coins and transfer public keys to reduce bandwidth.
We therefore first convert our adversary into an adversary for a
variant protocol in which these commitments contain the full values:
We rewind the adversary to try two distinct $\gamma \in 1,\dots,\kappa$
during each refresh operation, so that we obtain all values.
We need only try two choices because the adversary reveals all but
one planchet in each run. We now witness a hash collision if the
transfer secret the adversary reveals does not yield the correct coins.
If Taler satisfies unforgeability then this variant protocol does so too,
because an adversary against the protocol with commitment to full planchets
can trivially be replaced by an adversary against the protocol with hash
commitments.
We consider the directed forest on coins induced by the refresh protocol.
It follows from unforgeability that any coin must originate from some
customer's withdraw in this graph.
We may assume that all $\V{coin}_1, \dots, \V{coin}_l$ originate from
noncorrupted users, for some $l \leq \ell$. % So $\ell \leq w + X$.
For any $i \leq l$, there is a final refresh operation $R_i$ in which
a noncorrupted user could obtain the coin $C'$ consumed in the refresh
via the linking protocol, but no noncorrupted user could obtain the
coin provided by the refresh, as otherwise $\V{coin}_i$ gets marked as
spent in step step \ref{game:income:spend}.
Set $F := \{ R_i \mid i \leq l \}$. %TODO: Not ellegant, clean up below.
During each $R_i \in F$, our adversary must have submitted a blinded
coin and transfer public key for which the linking protocol fails to
produce the resulting coin correctly, otherwise the coin would have
been spent in step \ref{game:income:spend}. In this case, we consider
several nonexclusive cases
\begin{enumerate}
\item the execution of the refresh protocol is incomplete,
\item the commitment for the $\gamma$th blinded coin and transfer
public key was wrong,
\item a commitment for a blinded coin and transfer public key other
than the $\gamma$th was wrong,
\end{enumerate}
We show these to be exhaustive by assuming their converses all hold:
As the commitment is signed, our our honest key generation assumption
of $\textsc{CoinSignKx}$ applies to the coin public key.
We assumed the $\gamma$th transfer public key is honest too, so
our key exchange completeness assumption of $\textsc{CoinSignKx}$
yields $t C' \neq c' T$ where $T = t G$ is the transfer key,
so the customer obtains the correct transfer secret.
We assumed the refresh concluded and all submissions besides the
$\gamma$th were honest, so the exchange correctly reveals the signed
blinded coin. We assumed the $\gamma$th blinded coin is correct too,
so customer now recompute the new coin correctly, violating $R_i \in F$.
We shall prove
\begin{equation}\label{eq:incometransparencyproof}
\Exp{{p \over b + p} \middle F \neq \emptyset} = {1\over\kappa}
\end{equation}
where the expectation runs over
any probability space used by the adversary and challenger.
We shall optimize our adversary in ways that maximize $p \over b + p$ by
proving the optimised adversary exists. We may then restrict our analysis
to these optimised adversaries, which suffices for our game result, but
note that reductions would not permit this trick.
%TODO: Improve??
As a reminder, if a refresh operation is initiated using a false commitment
that is detected by the exchange, then the new coin cannot be obtained, and
contributes to the lost coins $b := w  s$ instead of the winnings $p := L 
w'$. We also note $b + p$ gives the value of
refreshes attempted with false commitments. As these are nonnegative, $p \over b
+ p$ is undefined if and only if $p = 0$ and $b = 0$, which happens if and
only if the adversary does not use false commitments, i.e. $F = \emptyset$.
We may now assume for optimality that $\mathcal{A}$ submits a false
commitment for at most one choice of $\gamma$ in any $R_i \in F$, as
otherwise the refresh always fails. Furthermore for an optimal adversary we
can exclude refreshes in $F$ that are incomplete, but that would be possible
to complete successfully, as completing such a refresh would only increase the
adversaries winnings.
We emphasize that an adversary that loses an $R_i$ loses the coin that would
have resulted from it completely, while an optimal adversary who wins an
$R_i$ should not gamble again. Indeed, an adversary has no reason to touch
its winnings from an $R_i$.
% There is no way to influence $p$ or $b$ through withdrawals or spends
% by corrupted users of course. In principle, one could decrease $b$ by
% sharing from a corrupted user to a noncorrupted users,
% but we may assume this does not occur either, again by optimality.
For any $R_i$, there are $\kappa$ game runs identical up through the
commitment phase of $R_i$ and exhibit different outcomes based on the
challenger's random choice of $\gamma$.
If $v_i$ is the financial value of the coin resulting from refresh operation
$R_i$ then one of the possible runs adds $v_i$ to $p$, while the remaining
$\kappa1$ runs add $v_i$ to $b$.
We define $p_i$ and $b_i$ to be these contributions summed over the $\kappa$ possible
runs, i.e.
\begin{align*}
p_i &:= v_i\\
b_i &= (\kappa  1)v_i
\end{align*}
The adversary will succeed in $1/\kappa$ runs ($p_i=v$) and looses in
$(\kappa1)/\kappa$ runs ($p_i=0$). Hence:
\begin{align*}
\Exp{{p \over b + p} \middle F \neq \emptyset}
&= \frac{1}{F} \sum_{R_i\in F} {p_i \over b_i + p_i} \\
&= \frac{1}{\kappa F} \sum_{R_i\in F} {v_i \over 0 + v_i} + \frac{\kappa1}{\kappa F} \sum_{R_i \in F} {0 \over v_i + 0} \\
&= {1\over\kappa},
\end{align*}
which yields the equality (\ref{eq:incometransparencyproof}).
As for $F = \emptyset$, the return value of the game must be $0$, we conclude
\begin{equation*}
E\left[\mathit{Exp}_{\cal A}^{income}(1^\lambda, 1^\kappa)\right] \le {1\over\kappa}.
\end{equation*}
\end{proof}
% https://math.stackexchange.com/questions/852890/expectationofrandomvariablesratio
%%% $L  w' \over (L  w') + (w  s)$
% $E(b/p) + 1 = E(b/p + p/p) = E((b + p)/p) = \kappa$
% $E(b/p) = \kappa1$
% $E(p/b) = {1 \over \kappa1}$
%As it turns out, there is a simple hashbased solutions that provides
%postquantum anonymity without additional assumptions though:
%% because the coin holder is encrypting to themselves:
%We extend the coin private key $c$ by a secret $m$ and extend the
%coin signing key $C$ to be a pair $(C,R)$ in which $R$ is the root
%of a Merkle tree whose $i$th leave is $H(m,i)$.
%In a refresh, the wallet first constructs the planchets from
%$H(t C', H(m,i))$ and commits to the index $i$ along with with each
%transfer public key $T$, and later when revealing $t$ also reveals
%$H(m,i)$ and the proof that it lives in the Merkle tree.
%In this scheme, our Merkle tree should be large enough to accommodate
%some fixed number of refreshes per coin, possibly just one, while
%our wallet must avoid any fragility in committing its $i$ choices to disk.
%\section{Standard Definitions}
%\begin{definition}[Onemore forgery]
%For any integer $\ell$, an $(\ell, \ell + 1)$forgery comes from
%a probabilistic polynomial time Turing machine $\mathcal{A}$ that can
%compute, after $\ell$ interactions with the signer $\Sigma$, $\ell + 1$ signatures with nonnegligible
% probability. The ``onemore forgery'' is an $(\ell, \ell + 1)$forgery for some
%integer $\ell$.
%\end{definition}
%
%Taken from \cite{pointcheval1996provably}. This definition applies to blind signature schemes in general.
%Intuition: EUFCMA is not strong enough for blind signatures,
%since attacker can choose the message (without going through a hash function before signing).
%
%\begin{definition}[RSAKTI]
% Game (security parameter $k$, $m : \mathbb{N} \rightarrow \mathbb{N}$,
% $\mathcal{A}$ adversary with access to RSA inversion oracle \ora{Inv}):
% \begin{enumerate}
% \item $(N, e, d) \leftarrow \mathsf{KeyGen}(k)$
% \item $y_i \leftarrow_R \mathbb{Z}_N^*$ for $i \in 1, \dots, m(k) + 1$
% \item $(x_1, \dots, x_{m(k) + 1}) \leftarrow \mathcal{A}^{\ora{Inv}}(N, e, k, y_1, \dots, y_{m(k) + 1})$
% \item $\mathcal{A}$ wins if it made at most $m(k)$ oracle queries and $x_i^e \equiv y_i \pmod N$
% \end{enumerate}
%\end{definition}
%
%From \cite{bellare2003onemore}. Assumption to prove security of RSA blind signatures. Equivalent to RSACTI.
\section{Concrete Instantitation}
We now give a concrete instantiation that is used in our implementation
for the schemes \textsc{BlindSign}, \textsc{CoinSignKx} and \textsc{Sign}.
For \textsc{BlindSign}, we use RSAFDH blind signatures~\cite{chaum1983blind}
with blinding factors produced by FDPRF constructed from a hash
function with range the full RSA domain $[0,pq)$. An adversary who
defeats the blindness property can recognise blinding factors,
violating this PRF assumption. For the unforgeability property,
we require the RSAKTI assumption as discussed in \cite{bellare2003onemore}.
%TODO: Jeff always cited multoiiple papers for RSAKTI, but forgets why.
As the blinding step in RSA blind signatures is noninteractive, storage
and verification of the transcript is omitted in refresh and link.
We envision BLS blind signatures working equally well.
We instantiate \textsc{CoinSignKx} with signatures and key exchange operations
on elliptic curves in Edwards form, where the same key is used for signatures
and the DiffieHellman key exchange operations. In practice, we use Ed25519
signatures \cite{bernstein2012high} and scalar multiplication for the Edwards
curve, which gives $\lambda=128$. % Curve25519 \cite{bernstein2006curve25519}
We caution that some other elliptic curve key exchange implementations might
not satisfy the robustness property that we require, due to the lack of
complete addition laws.
% and that both robustness and honest key generation become tricky when ..
For \textsc{Sign}, we use ellipticcurve signatures, concretely Ed25519.
\section{Strong Income Transparency}
We characterized {\em weak income transparency} in
Section~\ref{sec:securityproperties}. We will now discuss the notion
of {\em strong income transparency}, as it might be useful for other
instantiations that provide more absolute guarantees.
\begin{definition}[Strong Income Transparency]
We say that an ecash scheme satisfies \emph{strong income transparency} if
the success probability $\Prb{\mathit{Exp}_{\cal A}^{income}(1^\lambda, 1^\kappa) \ne 0}$
for the income transparency game is negligible for any polynomial time adversary~$\mathcal{A}$.
\end{definition}
The adversary is said to win one execution of the strong income
transparency game if the game's return value is nonzero, i.e. there
was at least one successful attempt to obtain untaxed income.
Our protocol does not fulfill strong income transparency since in
practice $\kappa$ must be a small cutandchoose parameter, as the
complexity of our cutandchoose protocol grows linearly with
$\kappa$. Instead we showed that our protocol satisfies weak income
transparency, which is a statement about the adversary's financial
loss when winning the game instead of the winning probability. Here,
the returnoninvestment (represented by the game's return value) is
bounded by $1/\kappa$.
\section{Endorsed ECash and Fair Exchange}
%\subsection{Other Properties}
% FIXME: move this to design or implementation
%\subsubsection{Fair Exchange}
% FIXME: we should mention "atomic swap" here
The Endorsed ECash system by Camenisch et al.~\cite{camenisch2007endorsed}
allows for fair exchange of (online or
offline) ecash against digital goods. The online version does not protect the
customer against loss of anonymity from linkability of aborted fair exchanges.
The refresh protocol can be used for fair exchange of online ecash against
digital goods, without any loss of anonymity due to linkability of aborted
transactions, with the following small extension: The customer asks the
exchange to \emph{lock coins to a merchant} until a timeout. Until the timeout
occurs, the exchange provides the merchant with a guarantee that these coins can
only be spent with this specific merchant, or not at all. The
fair exchange exchanges the merchant's digital goods against the customer's
deposit permissions for the locked coins. On aborted fair exchanges,
the customer refreshes to obtain unlinkable coins.
\end{document}
