(*<*)theory Paperimports "../Nominal/Test" "LaTeXsugar"beginnotation (latex output) swap ("'(_ _')" [1000, 1000] 1000) and fresh ("_ # _" [51, 51] 50) and fresh_star ("_ #* _" [51, 51] 50) and supp ("supp _" [78] 73) and uminus ("-_" [78] 73) and If ("if _ then _ else _" 10)(*>*)section {* Introduction *}text {* So far, Nominal Isabelle provides a mechanism for constructing alpha-equated terms, for example \begin{center} $t ::= x \mid t\;t \mid \lambda x. t$ \end{center} \noindent where free and bound variables have names. For such terms Nominal Isabelle derives automatically a reasoning infrastructure that has been used successfully in formalisations of an equivalence checking algorithm for LF \cite{UrbanCheneyBerghofer08}, Typed Scheme~\cite{TobinHochstadtFelleisen08}, several calculi for concurrency \cite{BengtsonParrow07,BengtsonParow09} and a strong normalisation result for cut-elimination in classical logic \cite{UrbanZhu08}. It has also been used by Pollack for formalisations in the locally-nameless approach to binding \cite{SatoPollack10}. However, Nominal Isabelle has fared less well in a formalisation of the algorithm W \cite{UrbanNipkow09}, where types and type-schemes are of the form % \begin{equation}\label{tysch} \begin{array}{l} T ::= x \mid T \rightarrow T \hspace{5mm} S ::= \forall \{x_1,\ldots, x_n\}. T \end{array} \end{equation} \noindent and the quantification $\forall$ binds a finite (possibly empty) set of type-variables. While it is possible to implement this kind of more general binders by iterating single binders, this leads to a rather clumsy formalisation of W. The need of iterating single binders is also one reason why Nominal Isabelle and similar theorem provers that only provide mechanisms for binding single variables have not fared extremely well with the more advanced tasks in the POPLmark challenge \cite{challenge05}, because also there one would like to bind multiple variables at once. Binding multiple variables has interesting properties that cannot be captured easily by iterating single binders. For example in case of type-schemes we do not want to make a distinction about the order of the bound variables. Therefore we would like to regard the following two type-schemes as alpha-equivalent % \begin{equation}\label{ex1} \forall \{x, y\}. x \rightarrow y \;\approx_\alpha\; \forall \{y, x\}. y \rightarrow x \end{equation} \noindent but assuming that $x$, $y$ and $z$ are distinct variables, the following two should \emph{not} be alpha-equivalent % \begin{equation}\label{ex2} \forall \{x, y\}. x \rightarrow y \;\not\approx_\alpha\; \forall \{z\}. z \rightarrow z \end{equation} \noindent Moreover, we like to regard type-schemes as alpha-equivalent, if they differ only on \emph{vacuous} binders, such as % \begin{equation}\label{ex3} \forall \{x\}. x \rightarrow y \;\approx_\alpha\; \forall \{x, z\}. x \rightarrow y \end{equation} \noindent where $z$ does not occur freely in the type. In this paper we will give a general binding mechanism and associated notion of alpha-equivalence that can be used to faithfully represent this kind of binding in Nominal Isabelle. The difficulty of finding the right notion for alpha-equivalence can be appreciated in this case by considering that the definition given by Leroy in \cite{Leroy92} is incorrect (it omits a side-condition). However, the notion of alpha-equivalence that is preserved by vacuous binders is not always wanted. For example in terms like % \begin{equation}\label{one} \LET x = 3 \AND y = 2 \IN x\,-\,y \END \end{equation} \noindent we might not care in which order the assignments $x = 3$ and $y = 2$ are given, but it would be unusual to regard \eqref{one} as alpha-equivalent with % \begin{center} $\LET x = 3 \AND y = 2 \AND z = loop \IN x\,-\,y \END$ \end{center} \noindent Therefore we will also provide a separate binding mechanism for cases in which the order of binders does not matter, but the ``cardinality'' of the binders has to agree. However, we found that this is still not sufficient for dealing with language constructs frequently occurring in programming language research. For example in $\mathtt{let}$s containing patterns % \begin{equation}\label{two} \LET (x, y) = (3, 2) \IN x\,-\,y \END \end{equation} \noindent we want to bind all variables from the pattern inside the body of the $\mathtt{let}$, but we also care about the order of these variables, since we do not want to regard \eqref{two} as alpha-equivalent with % \begin{center} $\LET (y, x) = (3, 2) \IN x\,- y\,\END$ \end{center} \noindent As a result, we provide three general binding mechanisms each of which binds multiple variables at once, and let the user chose which one is intended when formalising a programming language calculus. By providing these general binding mechanisms, however, we have to work around a problem that has been pointed out by Pottier in \cite{Pottier06} and Cheney in \cite{Cheney05}: in $\mathtt{let}$-constructs of the form % \begin{center} $\LET x_1 = t_1 \AND \ldots \AND x_n = t_n \IN s \END$ \end{center} \noindent which bind all the $x_i$ in $s$, we might not care about the order in which the $x_i = t_i$ are given, but we do care about the information that there are as many $x_i$ as there are $t_i$. We lose this information if we represent the $\mathtt{let}$-constructor by something like % \begin{center} $\LET [x_1,\ldots,x_n].s\;\; [t_1,\ldots,t_n]$ \end{center} \noindent where the notation $[\_\!\_].\_\!\_$ indicates that the $x_i$ become bound in $s$. In this representation the term \mbox{$\LET [x].s\;\;[t_1,t_2]$} would be a perfectly legal instance. To exclude such terms, additional predicates about well-formed terms are needed in order to ensure that the two lists are of equal length. This can result into very messy reasoning (see for example~\cite{BengtsonParow09}). To avoid this, we will allow type specifications for $\mathtt{let}$s as follows % \begin{center} \begin{tabular}{r@ {\hspace{2mm}}r@ {\hspace{2mm}}l} $trm$ & $::=$ & \ldots\\ & $\mid$ & $\mathtt{let}\;a\!::\!assn\;\;s\!::\!trm\quad\mathtt{bind}\;bn\,(a) \IN s$\\[1mm] $assn$ & $::=$ & $\mathtt{anil}$\\ & $\mid$ & $\mathtt{acons}\;\;name\;\;trm\;\;assn$ \end{tabular} \end{center} \noindent where $assn$ is an auxiliary type representing a list of assignments and $bn$ an auxiliary function identifying the variables to be bound by the $\mathtt{let}$. This function is defined by recursion over $assn$ as follows \begin{center} $bn\,(\mathtt{anil}) = \varnothing \qquad bn\,(\mathtt{acons}\;x\;t\;as) = \{x\} \cup bn\,(as)$ \end{center} \noindent The scope of the binding is indicated by labels given to the types, for example \mbox{$s\!::\!trm$}, and a binding clause, in this case $\mathtt{bind}\;bn\,(a) \IN s$, that states to bind in $s$ all the names the function call $bn\,(a)$ returns. This style of specifying terms and bindings is heavily inspired by the syntax of the Ott-tool \cite{ott-jfp}. However, we will not be able to deal with all specifications that are allowed by Ott. One reason is that Ott lets the user to specify ``empty'' types like \begin{center} $t ::= t\;t \mid \lambda x. t$ \end{center} \noindent where no clause for variables is given. Arguably, such specifications make some sense in the context of Coq's type theory (which Ott supports), but not at all in a HOL-based environment where every datatype must have a non-empty set-theoretic model. Another reason is that we establish the reasoning infrastructure for alpha-\emph{equated} terms. In contrast, Ott produces a reasoning infrastructure in Isabelle/HOL for \emph{non}-alpha-equated, or ``raw'', terms. While our alpha-equated terms and the raw terms produced by Ott use names for bound variables, there is a key difference: working with alpha-equated terms means that the two type-schemes (with $x$, $y$ and $z$ being distinct) \begin{center} $\forall \{x\}. x \rightarrow y \;=\; \forall \{x, z\}. x \rightarrow y$ \end{center} \noindent are not just alpha-equal, but actually \emph{equal}. As a result, we can only support specifications that make sense on the level of alpha-equated terms (offending specifications, which for example bind a variable according to a variable bound somewhere else, are not excluded by Ott, but we have to). Our insistence on reasoning with alpha-equated terms comes from the wealth of experience we gained with the older version of Nominal Isabelle: for non-trivial properties, reasoning about alpha-equated terms is much easier than reasoning with raw terms. The fundamental reason for this is that the HOL-logic underlying Nominal Isabelle allows us to replace ``equals-by-equals''. In contrast, replacing ``alpha-equals-by-alpha-equals'' in a representation based on raw terms requires a lot of extra reasoning work. Although in informal settings a reasoning infrastructure for alpha-equated terms is nearly always taken for granted, establishing it automatically in the Isabelle/HOL theorem prover is a rather non-trivial task. For every specification we will need to construct a type containing as elements the alpha-equated terms. To do so, we use the standard HOL-technique of defining a new type by identifying a non-empty subset of an existing type. The construction we perform in HOL can be illustrated by the following picture: \begin{center} \begin{tikzpicture} %\draw[step=2mm] (-4,-1) grid (4,1); \draw[very thick] (0.7,0.4) circle (4.25mm); \draw[rounded corners=1mm, very thick] ( 0.0,-0.8) rectangle ( 1.8, 0.9); \draw[rounded corners=1mm, very thick] (-1.95,0.85) rectangle (-2.85,-0.05); \draw (-2.0, 0.845) -- (0.7,0.845); \draw (-2.0,-0.045) -- (0.7,-0.045); \draw ( 0.7, 0.4) node {\begin{tabular}{@ {}c@ {}}$\alpha$-\\[-1mm]clas.\end{tabular}}; \draw (-2.4, 0.4) node {\begin{tabular}{@ {}c@ {}}$\alpha$-eq.\\[-1mm]terms\end{tabular}}; \draw (1.8, 0.48) node[right=-0.1mm] {\begin{tabular}{@ {}l@ {}}existing\\[-1mm] type\\ (sets of raw terms)\end{tabular}}; \draw (0.9, -0.35) node {\begin{tabular}{@ {}l@ {}}non-empty\\[-1mm]subset\end{tabular}}; \draw (-3.25, 0.55) node {\begin{tabular}{@ {}l@ {}}new\\[-1mm]type\end{tabular}}; \draw[<->, very thick] (-1.8, 0.3) -- (-0.1,0.3); \draw (-0.95, 0.3) node[above=0mm] {isomorphism}; \end{tikzpicture} \end{center} \noindent We take as the starting point a definition of raw terms (defined as a datatype in Isabelle/HOL); identify then the alpha-equivalence classes in the type of sets of raw terms according to our alpha-equivalence relation and finally define the new type as these alpha-equivalence classes (non-emptiness is satisfied whenever the raw terms are definable as datatype in Isabelle/HOL and the fact that our relation for alpha-equivalence is indeed an equivalence relation). The fact that we obtain an isomorphism between the new type and the non-empty subset shows that the new type is a faithful representation of alpha-equated terms. That is not the case for example for terms using the locally nameless representation of binders \cite{McKinnaPollack99}: in this representation there are ``junk'' terms that need to be excluded by reasoning about a well-formedness predicate. The problem with introducing a new type in Isabelle/HOL is that in order to be useful, a reasoning infrastructure needs to be ``lifted'' from the underlying subset to the new type. This is usually a tricky and arduous task. To ease it, we re-implemented in Isabelle/HOL the quotient package described by Homeier \cite{Homeier05} for the HOL4 system. This package allows us to lift definitions and theorems involving raw terms to definitions and theorems involving alpha-equated terms. For example if we define the free-variable function over raw lambda-terms \begin{center} $\fv(x) = \{x\}$\hspace{10mm} $\fv(t_1\;t_2) = \fv(t_1) \cup \fv(t_2)$\\[1mm] $\fv(\lambda x.t) = \fv(t) - \{x\}$ \end{center} \noindent then with not too great effort we obtain a function $\fv^\alpha$ operating on quotients, or alpha-equivalence classes of lambda-terms. This lifted function is characterised by the equations \begin{center} $\fv^\alpha(x) = \{x\}$\hspace{10mm} $\fv^\alpha(t_1\;t_2) = \fv^\alpha(t_1) \cup \fv^\alpha(t_2)$\\[1mm] $\fv^\alpha(\lambda x.t) = \fv^\alpha(t) - \{x\}$ \end{center} \noindent (Note that this means also the term-constructors for variables, applications and lambda are lifted to the quotient level.) This construction, of course, only works if alpha-equivalence is indeed an equivalence relation, and the lifted definitions and theorems are respectful w.r.t.~alpha-equivalence. Accordingly, we will not be able to lift a bound-variable function to alpha-equated terms (since it does not respect alpha-equivalence). To sum up, every lifting of theorems to the quotient level needs proofs of some respectfulness properties. In the paper we show that we are able to automate these proofs and therefore can establish a reasoning infrastructure for alpha-equated terms.\medskip \noindent {\bf Contributions:} We provide new definitions for when terms involving multiple binders are alpha-equivalent. These definitions are inspired by earlier work of Pitts \cite{Pitts04}. By means of automatic proofs, we establish a reasoning infrastructure for alpha-equated terms, including properties about support, freshness and equality conditions for alpha-equated terms. We are also able to derive, at the moment only manually, strong induction principles that have the variable convention already built in.*}section {* A Short Review of the Nominal Logic Work *}text {* At its core, Nominal Isabelle is an adaption of the nominal logic work by Pitts \cite{Pitts03}. This adaptation for Isabelle/HOL is described in \cite{HuffmanUrban10}, which we review here briefly to aid the description of what follows. Two central notions in the nominal logic work are sorted atoms and sort-respecting permutations of atoms. The sorts can be used to represent different kinds of variables, such as term- and type-variables in Core-Haskell, and it is assumed that there is an infinite supply of atoms for each sort. However, in order to simplify the description, we shall assume in what follows that there is only one sort of atoms. Permutations are bijective functions from atoms to atoms that are the identity everywhere except on a finite number of atoms. There is a two-place permutation operation written % @{text[display,indent=5] "_ \<bullet> _ :: perm \<Rightarrow> \<beta> \<Rightarrow> \<beta>"} \noindent in which the generic type @{text "\<beta>"} stands for the type of the object on which the permutation acts. In Nominal Isabelle, the identity permutation is written as @{term "0::perm"}, the composition of two permutations @{term p} and @{term q} as \mbox{@{term "p + q"}} and the inverse permutation of @{term p} as @{text "- p"}. The permutation operation is defined for products, lists, sets, functions, booleans etc (see \cite{HuffmanUrban10}). Concrete permutations are build up from swappings, written as @{text "(a b)"}, which are permutations that behave as follows: % @{text[display,indent=5] "(a b) = \<lambda>c. if a = c then b else if b = c then a else c"} The most original aspect of the nominal logic work of Pitts is a general definition for the notion of ``the set of free variables of an object @{text "x"}''. This notion, written @{term "supp x"}, is general in the sense that it applies not only to lambda-terms (alpha-equated or not), but also to lists, products, sets and even functions. The definition depends only on the permutation operation and on the notion of equality defined for the type of @{text x}, namely: % @{thm[display,indent=5] supp_def[no_vars, THEN eq_reflection]} \noindent There is also the derived notion for when an atom @{text a} is \emph{fresh} for an @{text x}, defined as % @{thm[display,indent=5] fresh_def[no_vars]} \noindent We also use for sets of atoms the abbreviation @{thm (lhs) fresh_star_def[no_vars]} defined as @{thm (rhs) fresh_star_def[no_vars]}. A striking consequence of these definitions is that we can prove without knowing anything about the structure of @{term x} that swapping two fresh atoms, say @{text a} and @{text b}, leave @{text x} unchanged. \begin{property} @{thm[mode=IfThen] swap_fresh_fresh[no_vars]} \end{property} \noindent For a proof see \cite{HuffmanUrban10}. \begin{property} @{thm[mode=IfThen] at_set_avoiding[no_vars]} \end{property}*}section {* General Binders\label{sec:binders} *}text {* In Nominal Isabelle, the user is expected to write down a specification of a term-calculus and then a reasoning infrastructure is automatically derived from this specification (remember that Nominal Isabelle is a definitional extension of Isabelle/HOL, which does not introduce any new axioms). In order to keep our work manageable, we will wherever possible state definitions and perform proofs inside Isabelle, as opposed to write custom ML-code that generates them anew for each specification. To that end, we will consider pairs @{text "(as, x)"} of type @{text "(atom set) \<times> \<beta>"}. These pairs are intended to represent the abstraction, or binding, of the set @{text "as"} in the body @{text "x"}. The first question we have to answer is when the pairs $(as, x)$ and $(bs, y)$ are alpha-equivalent? (At the moment we are interested in the notion of alpha-equivalence that is \emph{not} preserved by adding vacuous binders.) To answer this, we identify four conditions: {\it i)} given a free-variable function $\fv$ of type \mbox{@{text "\<beta> \<Rightarrow> atom set"}}, then @{text x} and @{text y} need to have the same set of free variables; moreover there must be a permutation @{text p} such that {\it ii)} it leaves the free variables of @{text x} and @{text y} unchanged, but {\it iii)} ``moves'' their bound names so that we obtain modulo a relation, say \mbox{@{text "_ R _"}}, two equal terms. We also require {\it iv)} that @{text p} makes the abstracted sets @{text as} and @{text bs} equal. The requirements {\it i)} to {\it iv)} can be stated formally as follows: % \begin{equation}\label{alphaset} \begin{array}{@ {\hspace{10mm}}r@ {\hspace{2mm}}l} \multicolumn{2}{l}{(as, x) \approx\hspace{0.05mm}_{set}^{\fv, R, p} (bs, y) \;\dn\hspace{30mm}\;}\\[1mm] & @{text "fv(x) - as = fv(y) - bs"}\\ \wedge & @{text "(fv(x) - as) #* p"}\\ \wedge & @{text "(p \<bullet> x) R y"}\\ \wedge & @{text "(p \<bullet> as) = bs"}\\ \end{array} \end{equation} \noindent Note that this relation is dependent on $p$. Alpha-equivalence is then the relation where we existentially quantify over this $p$. Also note that the relation is dependent on a free-variable function $\fv$ and a relation $R$. The reason for this extra generality is that we will use $\approx_{set}$ for both ``raw'' terms and alpha-equated terms. In the latter case, $R$ will be replaced by equality $(op =)$ and we are going to prove that $\fv$ will be equal to the support of $x$ and $y$. The definition in \eqref{alphaset} does not make any distinction between the order of abstracted variables. If we want this, then we can define alpha-equivalence for pairs of the form \mbox{@{text "(as, x)"}} with type @{text "(atom list) \<times> \<beta>"} as follows % \begin{equation}\label{alphalist} \begin{array}{@ {\hspace{10mm}}r@ {\hspace{2mm}}l} \multicolumn{2}{l}{(as, x) \approx\hspace{0.05mm}_{list}^{\fv, R, p} (bs, y) \;\dn\hspace{30mm}\;}\\[1mm] & @{text "fv(x) - (set as) = fv(y) - (set bs)"}\\ \wedge & @{text "(fv(x) - set as) #* p"}\\ \wedge & @{text "(p \<bullet> x) R y"}\\ \wedge & @{text "(p \<bullet> as) = bs"}\\ \end{array} \end{equation} \noindent where $set$ is the function that coerces a list of atoms into a set of atoms. If we do not want to make any difference between the order of binders and also allow vacuous binders, then we keep sets of binders, but drop the fourth condition in \eqref{alphaset}: % \begin{equation}\label{alphares} \begin{array}{@ {\hspace{10mm}}r@ {\hspace{2mm}}l} \multicolumn{2}{l}{(as, x) \approx\hspace{0.05mm}_{res}^{\fv, R, p} (bs, y) \;\dn\hspace{30mm}\;}\\[1mm] & @{text "fv(x) - as = fv(y) - bs"}\\ \wedge & @{text "(fv(x) - as) #* p"}\\ \wedge & @{text "(p \<bullet> x) R y"}\\ \end{array} \end{equation} \begin{exmple}\rm It might be useful to consider some examples for how these definitions pan out in practise. For this consider the case of abstracting a set of variables over types (as in type-schemes). We set $R$ to be the equality and for $\fv(T)$ we define \begin{center} $\fv(x) = \{x\} \qquad \fv(T_1 \rightarrow T_2) = \fv(T_1) \cup \fv(T_2)$ \end{center} \noindent Now recall the examples shown in \eqref{ex1}, \eqref{ex2} and \eqref{ex3}. It can be easily checked that @{text "({x, y}, x \<rightarrow> y)"} and @{text "({y, x}, y \<rightarrow> x)"} are equal according to $\approx_{set}$ and $\approx_{res}$ by taking $p$ to be the swapping @{term "(x \<rightleftharpoons> y)"}. In case of @{text "x \<noteq> y"}, then $([x, y], x \rightarrow y) \not\approx_{list} ([y,x], x \rightarrow y)$ since there is no permutation that makes the lists @{text "[x, y]"} and @{text "[y, x]"} equal, and also leaves the type \mbox{@{text "x \<rightarrow> y"}} unchanged. Another examples is $(\{x\}, x) \approx_{res} (\{x,y\}, x)$ which holds by taking $p$ to be the identity permutation. However, if @{text "x \<noteq> y"}, then $(\{x\}, x) \not\approx_{set} (\{x,y\}, x)$ since there is no permutation that makes the sets $\{x\}$ and $\{x,y\}$ equal (similarly for $\approx_{list}$). \end{exmple} \noindent Let $\star$ range over $\{set, res, list\}$. We prove next under which conditions the $\approx\hspace{0.05mm}_\star^{\fv, R, p}$ are equivalence relations and equivariant: \begin{lemma} {\it i)} Given the fact that $x\;R\;x$ holds, then $(as, x) \approx\hspace{0.05mm}^{\fv, R, 0}_\star (as, x)$. {\it ii)} Given that @{text "(p \<bullet> x) R y"} implies @{text "(-p \<bullet> y) R x"}, then $(as, x) \approx\hspace{0.05mm}^{\fv, R, p}_\star (bs, y)$ implies $(bs, y) \approx\hspace{0.05mm}^{\fv, R, - p}_\star (as, x)$. {\it iii)} Given that @{text "(p \<bullet> x) R y"} and @{text "(q \<bullet> y) R z"} implies @{text "((q + p) \<bullet> x) R z"}, then $(as, x) \approx\hspace{0.05mm}^{\fv, R, p}_\star (bs, y)$ and $(bs, y) \approx\hspace{0.05mm}^{\fv, R, q}_\star (cs, z)$ implies $(as, x) \approx\hspace{0.05mm}^{\fv, R, q + p}_\star (cs, z)$. Given @{text "(q \<bullet> x) R y"} implies @{text "(p \<bullet> (q \<bullet> x)) R (p \<bullet> y)"} and @{text "p \<bullet> (fv x) = fv (p \<bullet> x)"} then @{text "p \<bullet> (fv y) = fv (p \<bullet> y)"}, then $(as, x) \approx\hspace{0.05mm}^{\fv, R, q}_\star (bs, y)$ implies $(p \;\isasymbullet\; as, p \;\isasymbullet\; x) \approx\hspace{0.05mm}^{\fv, R, q}_\star (p \;\isasymbullet\; bs, p \;\isasymbullet\; y)$. \end{lemma} \begin{proof} All properties are by unfolding the definitions and simple calculations. \end{proof} \begin{lemma} $supp ([as]set. x) = supp x - as$ \end{lemma}*}section {* Alpha-Equivalence and Free Variables *}text {* Our specifications for term-calculi are heavily inspired by the syntax of the Ott-tool \cite{ott-jfp}. A specification is a collection of (possibly mutual recursive) type declarations, say $ty^\alpha_1$, $ty^\alpha_2$, \ldots $ty^\alpha_n$, and an associated collection of binding function declarations, say $bn^\alpha_1$, \ldots $bn^\alpha_m$. The syntax in Nominal Isabelle for such specifications is rougly as follows: % \begin{equation}\label{scheme} \mbox{\begin{tabular}{@ {\hspace{-9mm}}p{1.8cm}l} type \mbox{declaration part} & $\begin{cases} \mbox{\begin{tabular}{l} \isacommand{nominal\_datatype} $ty^\alpha_1 = \ldots$\\ \isacommand{and} $ty^\alpha_2 = \ldots$\\ $\ldots$\\ \isacommand{and} $ty^\alpha_n = \ldots$\\ \end{tabular}} \end{cases}$\\ binding \mbox{function part} & $\begin{cases} \mbox{\begin{tabular}{l} \isacommand{with} $bn^\alpha_1$ \isacommand{and} \ldots \isacommand{and} $bn^\alpha_m$ $\ldots$\\ \isacommand{where}\\ $\ldots$\\ \end{tabular}} \end{cases}$\\ \end{tabular}} \end{equation} \noindent Every type declaration $ty^\alpha_i$ consists of a collection of term-constructors, each of which comes with a list of labelled types that stand for the types of the arguments of the term-constructor. For example for a term-constructor $C^\alpha$ we might have \begin{center} $C^\alpha\;label_1\!::\!ty'_1\;\ldots label_l\!::\!ty'_l\;\;\textit{binding\_clauses}$ \end{center} \noindent whereby some of the $ty'_k$ (or their type components) are contained in the set of $ty^\alpha_i$ declared in \eqref{scheme}. In this case we will call the corresponding argument a \emph{recursive argument}. The labels annotated on the types are optional and can be used in the (possibly empty) list of \emph{binding clauses}. These clauses indicate the binders and the scope of the binders in a term-constructor. They come in three \emph{modes} \begin{center} \begin{tabular}{l} \isacommand{bind}\; {\it binders}\; \isacommand{in}\; {\it label}\\ \isacommand{bind\_set}\; {\it binders}\; \isacommand{in}\; {\it label}\\ \isacommand{bind\_res}\; {\it binders}\; \isacommand{in}\; {\it label}\\ \end{tabular} \end{center} \noindent The first mode is for binding lists of atoms (the order matters); the second is for sets of binders (the order does not matter, but cardinality does) and the last is for sets of binders (with vacuous binders preserving alpha-equivalence). In addition we distinguish between \emph{shallow} binders and \emph{deep} binders. Shallow binders are of the form \isacommand{bind}\; {\it label}\; \isacommand{in}\; {\it another\_label} (similar for the other two modes). The restriction we impose on shallow binders is that the {\it label} must either refer to a type that is an atom type or to a type that is a finite set or list of an atom type. For example the specifications of lambda-terms, where a single name is bound, and type-schemes, where a finite set of names is bound, use shallow binders (the type \emph{name} is an atom type): \begin{center} \begin{tabular}{@ {}cc@ {}} \begin{tabular}{@ {}l@ {\hspace{-1mm}}} \isacommand{nominal\_datatype} {\it lam} =\\ \hspace{5mm}\phantom{$\mid$} Var\;{\it name}\\ \hspace{5mm}$\mid$ App\;{\it lam}\;{\it lam}\\ \hspace{5mm}$\mid$ Lam\;{\it x::name}\;{\it t::lam}\\ \hspace{21mm}\isacommand{bind} {\it x} \isacommand{in} {\it t}\\ \end{tabular} & \begin{tabular}{@ {}l@ {}} \isacommand{nominal\_datatype} {\it ty} =\\ \hspace{5mm}\phantom{$\mid$} TVar\;{\it name}\\ \hspace{5mm}$\mid$ TFun\;{\it ty}\;{\it ty}\\ \isacommand{and} {\it tsc} = All\;{\it xs::(name fset)}\;{\it T::ty}\\ \hspace{24mm}\isacommand{bind\_res} {\it xs} \isacommand{in} {\it T}\\ \end{tabular} \end{tabular} \end{center} \noindent If we have shallow binders that ``share'' a body, for instance $t$ in the term-constructor Foo$_0$ \begin{center} \begin{tabular}{ll} \it {\rm Foo}$_0$ x::name y::name t::lam & \it \isacommand{bind}\;x\;\isacommand{in}\;t,\; \isacommand{bind}\;y\;\isacommand{in}\;t \end{tabular} \end{center} \noindent then we have to make sure the modes of the binders agree. We cannot have in the first binding clause the mode \isacommand{bind} and in the second \isacommand{bind\_set}. A \emph{deep} binder uses an auxiliary binding function that ``picks'' out the atoms in one argument of the term-constructor that can be bound in other arguments and also in the same argument (we will call such binders \emph{recursive}). The binding functions are expected to return either a set of atoms (for \isacommand{bind\_set} and \isacommand{bind\_res}) or a list of atoms (for \isacommand{bind}). They can be defined by primitive recursion over the corresponding type; the equations must be given in the binding function part of the scheme shown in \eqref{scheme}. For example for a calculus containing lets with tuple patterns, you might declare \begin{center} \begin{tabular}{l} \isacommand{nominal\_datatype} {\it trm} =\\ \hspace{5mm}\phantom{$\mid$} Var\;{\it name}\\ \hspace{5mm}$\mid$ App\;{\it trm}\;{\it trm}\\ \hspace{5mm}$\mid$ Lam\;{\it x::name}\;{\it t::trm} \;\;\isacommand{bind} {\it x} \isacommand{in} {\it t}\\ \hspace{5mm}$\mid$ Let\;{\it p::pat}\;{\it trm}\; {\it t::trm} \;\;\isacommand{bind\_set} {\it bn(p)} \isacommand{in} {\it t}\\ \isacommand{and} {\it pat} =\\ \hspace{5mm}\phantom{$\mid$} PNo\\ \hspace{5mm}$\mid$ PVr\;{\it name}\\ \hspace{5mm}$\mid$ PPr\;{\it pat}\;{\it pat}\\ \isacommand{with} {\it bn::pat $\Rightarrow$ atom set}\\ \isacommand{where} $bn(\textrm{PNil}) = \varnothing$\\ \hspace{5mm}$\mid$ $bn(\textrm{PVar}\;x) = \{atom\; x\}$\\ \hspace{5mm}$\mid$ $bn(\textrm{PPrd}\;p_1\;p_2) = bn(p_1) \cup bn(p_2)$\\ \end{tabular} \end{center} \noindent In this specification the function $atom$ coerces a name into the generic atom type of Nominal Isabelle. This allows us to treat binders of different atom type uniformly. As will shortly become clear, we cannot return an atom in a binding function that also is bound in the term-constructor. In the present version of Nominal Isabelle, we adopted the restriction the Ott-tool imposes on the binding functions, namely a binding function can only return the empty set (case PNil), a singleton set containing an atom (case PVar) or unions of atom sets (case PPrd). Moreover, as with shallow binders, deep binders with shared body need to have the same binding mode. Finally, the most drastic restriction we have to impose on deep binders is that we cannot have ``overlapping'' deep binders. Consider for example the term-constructors: \begin{center} \begin{tabular}{ll} \it {\rm Foo}$_1$ p::pat q::pat t::trm & \it \isacommand{bind}\;bn(p)\;\isacommand{in}\;t,\; \isacommand{bind}\;bn(q)\;\isacommand{in}\;t\\ \it {\rm Foo}$_2$ x::name p::pat t::trm & \it \it \isacommand{bind}\;x\;\isacommand{in}\;t,\; \isacommand{bind}\;bn(p)\;\isacommand{in}\;t \end{tabular} \end{center} \noindent In the first case we bind all atoms from the pattern $p$ in $t$ and also all atoms from $q$ in $t$. As a result we have no way to determine whether the binder came from the binding function in $p$ or $q$. Similarly in the second case: the binder $bn(p)$ overlaps with the shallow binder $x$. We must exclude such specifiactions, as we will not be able to represent them using the general binders described in Section \ref{sec:binders}. However the following two term-constructors are allowed: \begin{center} \begin{tabular}{ll} \it {\rm Bar}$_1$ p::pat t::trm s::trm & \it \isacommand{bind}\;bn(p)\;\isacommand{in}\;t,\; \isacommand{bind}\;bn(p)\;\isacommand{in}\;s\\ \it {\rm Bar}$_2$ p::pat t::trm & \it \isacommand{bind}\;bn(p)\;\isacommand{in}\;p,\; \isacommand{bind}\;bn(p)\;\isacommand{in}\;t\\ \end{tabular} \end{center} \noindent since there is no overlap of binders. Now the question is how we can turn specifications into actual type definitions in Isabelle/HOL and then establish a reasoning infrastructure for them? Because of the problem Pottier and Cheney pointed out, we cannot in general re-arrange arguments of term-constructors so that binders and their scopes next to each other, an then use the type constructors @{text "abs_set"}, @{text "abs_res"} and @{text "abs_list"}. Therefore we will first extract datatype definitions from the specification and then define an alpha-equiavlence relation over them. The datatype definition can be obtained by just stripping of the binding clauses and the labels on the types. We also have to invent new names for the types, $ty^\alpha$ and term-constructors $C^\alpha$ given by user. We just use an affix: \begin{center} $ty^\alpha \mapsto ty\_raw \qquad C^\alpha \mapsto C\_raw$ \end{center} \noindent This definition can be made, provided the usual conditions hold: the types must be non-empty and the types in the term-constructors need to be, what is called, positive position (see \cite{}). We then take the binding functions and define them by primitive recursion over the raw datatypes. binding function. The first non-trivial step is to read off from the specification free-variable functions. There are two kinds: free-variable functions corresponding to types, written $\fv\_ty$, and free-variable functions corresponding to binding functions, written $\fv\_bn$. They have to be defined at the same time since there can be interdependencies. Given a term-constructor $C ty_1 \ldots ty_n$ and some binding clauses, the function $\fv (C x_1 \ldots x_n)$ will be the union of the values generated for each argument, say $x_i$, as follows: \begin{center} \begin{tabular}{cp{8cm}} $\bullet$ & if it is a shallow binder, then $\varnothing$\\ $\bullet$ & if it is a deep binder, then $\fv_bn x_i$\\ $\bullet$ & if \end{tabular} \end{center}*}text {* Restrictions \begin{itemize} \item non-emptiness \item positive datatype definitions \item finitely supported abstractions \item respectfulness of the bn-functions\bigskip \item binders can only have a ``single scope'' \item all bindings must have the same mode \end{itemize}*}section {* Examples *}section {* Adequacy *}section {* Related Work *}text {* Ott is better with list dot specifications; subgrammars untyped; *}section {* Conclusion *}text {* Complication when the single scopedness restriction is lifted (two overlapping permutations)*}text {* TODO: function definitions: \medskip \noindent {\bf Acknowledgements:} We are very grateful to Andrew Pitts for many discussions about Nominal Isabelle. We thank Peter Sewell for making the informal notes \cite{SewellBestiary} available to us and also for patiently explaining some of the finer points about the abstract definitions and about the implementation of the Ott-tool. Lookup: Merlin paper by James Cheney; Mark Shinwell PhD Future work: distinct list abstraction*}(*<*)end(*>*)