(*<*)
theory Paper
imports "../Nominal/Test" "../Nominal/FSet" "LaTeXsugar"
begin
consts
fv :: "'a \<Rightarrow> 'b"
abs_set :: "'a \<Rightarrow> 'b \<Rightarrow> 'c"
alpha_bn :: "'a \<Rightarrow> 'a \<Rightarrow> bool"
abs_set2 :: "'a \<Rightarrow> perm \<Rightarrow> 'b \<Rightarrow> 'c"
Abs_dist :: "'a \<Rightarrow> 'b \<Rightarrow> 'c"
Abs_print :: "'a \<Rightarrow> 'b \<Rightarrow> 'c"
definition
"equal \<equiv> (op =)"
notation (latex output)
swap ("'(_ _')" [1000, 1000] 1000) and
fresh ("_ # _" [51, 51] 50) and
fresh_star ("_ #\<^sup>* _" [51, 51] 50) and
supp ("supp _" [78] 73) and
uminus ("-_" [78] 73) and
If ("if _ then _ else _" 10) and
alpha_gen ("_ \<approx>\<^raw:\raisebox{-1pt}{\makebox[0mm][l]{$\,_{\textit{set}}$}}>\<^bsup>_, _, _\<^esup> _") and
alpha_lst ("_ \<approx>\<^raw:\raisebox{-1pt}{\makebox[0mm][l]{$\,_{\textit{list}}$}}>\<^bsup>_, _, _\<^esup> _") and
alpha_res ("_ \<approx>\<^raw:\raisebox{-1pt}{\makebox[0mm][l]{$\,_{\textit{res}}$}}>\<^bsup>_, _, _\<^esup> _") and
abs_set ("_ \<approx>\<^raw:{$\,_{\textit{abs\_set}}$}> _") and
abs_set2 ("_ \<approx>\<^raw:\raisebox{-1pt}{\makebox[0mm][l]{$\,_{\textit{list}}$}}>\<^bsup>_\<^esup> _") and
fv ("fv'(_')" [100] 100) and
equal ("=") and
alpha_abs ("_ \<approx>\<^raw:{$\,_{\textit{abs\_set}}$}> _") and
Abs ("[_]\<^raw:$\!$>\<^bsub>set\<^esub>._" [20, 101] 999) and
Abs_lst ("[_]\<^raw:$\!$>\<^bsub>list\<^esub>._") and
Abs_dist ("[_]\<^raw:$\!$>\<^bsub>#list\<^esub>._") and
Abs_res ("[_]\<^raw:$\!$>\<^bsub>res\<^esub>._") and
Abs_print ("_\<^raw:$\!$>\<^bsub>set\<^esub>._") and
Cons ("_::_" [78,77] 73) and
supp_gen ("aux _" [1000] 10) and
alpha_bn ("_ \<approx>bn _")
(*>*)
section {* Introduction *}
text {*
%%% @{text "(1, (2, 3))"}
So far, Nominal Isabelle provides a mechanism for constructing
alpha-equated terms, for example
\begin{center}
@{text "t ::= x | t t | \<lambda>x. t"}
\end{center}
\noindent
where free and bound variables have names. For such alpha-equated 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{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,
respectively, of the form
%
\begin{equation}\label{tysch}
\begin{array}{l}
@{text "T ::= x | T \<rightarrow> T"}\hspace{5mm}
@{text "S ::= \<forall>{x\<^isub>1,\<dots>, x\<^isub>n}. T"}
\end{array}
\end{equation}
\noindent
and the @{text "\<forall>"}-quantification 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 the 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}
@{text "\<forall>{x, y}. x \<rightarrow> y \<approx>\<^isub>\<alpha> \<forall>{y, x}. y \<rightarrow> x"}
\end{equation}
\noindent
but assuming that @{text x}, @{text y} and @{text z} are distinct variables,
the following two should \emph{not} be alpha-equivalent
%
\begin{equation}\label{ex2}
@{text "\<forall>{x, y}. x \<rightarrow> y \<notapprox>\<^isub>\<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}
@{text "\<forall>{x}. x \<rightarrow> y \<approx>\<^isub>\<alpha> \<forall>{x, z}. x \<rightarrow> y"}
\end{equation}
\noindent
where @{text 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}
@{text "\<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}
@{text "\<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 @{text "\<LET>"}s containing patterns like
%
\begin{equation}\label{two}
@{text "\<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}
@{text "\<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 term-calculus.
By providing these general binding mechanisms, however, we have to work
around a problem that has been pointed out by Pottier \cite{Pottier06} and
Cheney \cite{Cheney05}: in @{text "\<LET>"}-constructs of the form
%
\begin{center}
@{text "\<LET> x\<^isub>1 = t\<^isub>1 \<AND> \<dots> \<AND> x\<^isub>n = t\<^isub>n \<IN> s \<END>"}
\end{center}
\noindent
which bind all the @{text "x\<^isub>i"} in @{text s}, we might not care
about the order in which the @{text "x\<^isub>i = t\<^isub>i"} are given,
but we do care about the information that there are as many @{text
"x\<^isub>i"} as there are @{text "t\<^isub>i"}. We lose this information if
we represent the @{text "\<LET>"}-constructor by something like
%
\begin{center}
@{text "\<LET> [x\<^isub>1,\<dots>,x\<^isub>n].s [t\<^isub>1,\<dots>,t\<^isub>n]"}
\end{center}
\noindent
where the notation @{text "[_]._"} indicates that the list of @{text "x\<^isub>i"}
becomes bound in @{text s}. In this representation the term
\mbox{@{text "\<LET> [x].s [t\<^isub>1, t\<^isub>2]"}} is a perfectly legal
instance, but the lengths of the two lists do not agree. 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 in 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}
@{text trm} & @{text "::="} & @{text "\<dots>"}\\
& @{text "|"} & @{text "\<LET> as::assn s::trm"}\hspace{4mm}
\isacommand{bind} @{text "bn(as)"} \isacommand{in} @{text "s"}\\[1mm]
@{text assn} & @{text "::="} & @{text "\<ANIL>"}\\
& @{text "|"} & @{text "\<ACONS> name trm assn"}
\end{tabular}
\end{center}
\noindent
where @{text assn} is an auxiliary type representing a list of assignments
and @{text bn} an auxiliary function identifying the variables to be bound
by the @{text "\<LET>"}. This function can be defined by recursion over @{text
assn} as follows
\begin{center}
@{text "bn(\<ANIL>) ="} @{term "{}"} \hspace{5mm}
@{text "bn(\<ACONS> x t as) = {x} \<union> bn(as)"}
\end{center}
\noindent
The scope of the binding is indicated by labels given to the types, for
example @{text "s::trm"}, and a binding clause, in this case
\isacommand{bind} @{text "bn(as)"} \isacommand{in} @{text "s"}. This binding
clause states to bind in @{text s} all the names the function @{text
"bn(as)"} 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 cope with all specifications that are
allowed by Ott. One reason is that Ott lets the user specify ``empty''
types like
\begin{center}
@{text "t ::= t t | \<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 \cite{Berghofer99}.
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, for example,
that the two type-schemes
\begin{center}
@{text "\<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 with 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 Isabelle/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 property 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}
@{text "fv(x) = {x}"}\hspace{10mm}
@{text "fv(t\<^isub>1 t\<^isub>2) = fv(t\<^isub>1) \<union> fv(t\<^isub>2)"}\\[1mm]
@{text "fv(\<lambda>x.t) = fv(t) - {x}"}
\end{center}
\noindent
then with the help of the quotient package we can obtain a function @{text "fv\<^sup>\<alpha>"}
operating on quotients, or alpha-equivalence classes of lambda-terms. This
lifted function is characterised by the equations
\begin{center}
@{text "fv\<^sup>\<alpha>(x) = {x}"}\hspace{10mm}
@{text "fv\<^sup>\<alpha>(t\<^isub>1 t\<^isub>2) = fv\<^sup>\<alpha>(t\<^isub>1) \<union> fv\<^sup>\<alpha>(t\<^isub>2)"}\\[1mm]
@{text "fv\<^sup>\<alpha>(\<lambda>x.t) = fv\<^sup>\<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
``raw'' definitions and theorems are respectful w.r.t.~alpha-equivalence.
For example, we will not be able to lift a bound-variable function. Although
this function can be defined for raw terms, it does not respect
alpha-equivalence and therefore cannot be lifted. To sum up, every lifting
of theorems to the quotient level needs proofs of some respectfulness
properties (see \cite{Homeier05}). In the paper we show that we are able to
automate these proofs and as a result can automatically establish a reasoning
infrastructure for alpha-equated terms.
The examples we have in mind where our reasoning infrastructure will be
helpful includes the term language of System @{text "F\<^isub>C"}, also
known as Core-Haskell (see Figure~\ref{corehas}). This term language
involves patterns that have lists of type-, coercion- and term-variables,
all of which are bound in @{text "\<CASE>"}-expressions. One
feature is that we do not know in advance how many variables need to
be bound. Another is that each bound variable comes with a kind or type
annotation. Representing such binders with single binders and reasoning
about them in a theorem prover would be a major pain. \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, for the moment
only manually, strong induction principles that
have the variable convention already built in.
\begin{figure}
\begin{boxedminipage}{\linewidth}
\begin{center}
\begin{tabular}{r@ {\hspace{1mm}}r@ {\hspace{2mm}}l}
\multicolumn{3}{@ {}l}{Type Kinds}\\
@{text "\<kappa>"} & @{text "::="} & @{text "\<star> | \<kappa>\<^isub>1 \<rightarrow> \<kappa>\<^isub>2"}\smallskip\\
\multicolumn{3}{@ {}l}{Coercion Kinds}\\
@{text "\<iota>"} & @{text "::="} & @{text "\<sigma>\<^isub>1 \<sim> \<sigma>\<^isub>2"}\smallskip\\
\multicolumn{3}{@ {}l}{Types}\\
@{text "\<sigma>"} & @{text "::="} & @{text "a | T | \<sigma>\<^isub>1 \<sigma>\<^isub>2 | S\<^isub>n"}$\;\overline{@{text "\<sigma>"}}$@{text "\<^sup>n"}
@{text "| \<forall>a:\<kappa>. \<sigma> | \<iota> \<Rightarrow> \<sigma>"}\smallskip\\
\multicolumn{3}{@ {}l}{Coercion Types}\\
@{text "\<gamma>"} & @{text "::="} & @{text "c | C | \<gamma>\<^isub>1 \<gamma>\<^isub>2 | S\<^isub>n"}$\;\overline{@{text "\<gamma>"}}$@{text "\<^sup>n"}
@{text "| \<forall>c:\<iota>. \<gamma> | \<iota> \<Rightarrow> \<gamma> "}\\
& @{text "|"} & @{text "refl \<sigma> | sym \<gamma> | \<gamma>\<^isub>1 \<circ> \<gamma>\<^isub>2 | \<gamma> @ \<sigma> | left \<gamma> | right \<gamma>"}\\
& @{text "|"} & @{text "\<gamma>\<^isub>1 \<sim> \<gamma>\<^isub>2 | rightc \<gamma> | leftc \<gamma> | \<gamma>\<^isub>1 \<triangleright> \<gamma>\<^isub>2"}\smallskip\\
\multicolumn{3}{@ {}l}{Terms}\\
@{text "e"} & @{text "::="} & @{text "x | K | \<Lambda>a:\<kappa>. e | \<Lambda>c:\<iota>. e | e \<sigma> | e \<gamma>"}\\
& @{text "|"} & @{text "\<lambda>x:\<sigma>. e | e\<^isub>1 e\<^isub>2 | \<LET> x:\<sigma> = e\<^isub>1 \<IN> e\<^isub>2"}\\
& @{text "|"} & @{text "\<CASE> e\<^isub>1 \<OF>"}$\;\overline{@{text "p \<rightarrow> e\<^isub>2"}}$ @{text "| e \<triangleright> \<gamma>"}\smallskip\\
\multicolumn{3}{@ {}l}{Patterns}\\
@{text "p"} & @{text "::="} & @{text "K"}$\;\overline{@{text "a:\<kappa>"}}\;\overline{@{text "c:\<iota>"}}\;\overline{@{text "x:\<sigma>"}}$\smallskip\\
\multicolumn{3}{@ {}l}{Constants}\\
& @{text C} & coercion constants\\
& @{text T} & value type constructors\\
& @{text "S\<^isub>n"} & n-ary type functions (which need to be fully applied)\\
& @{text K} & data constructors\smallskip\\
\multicolumn{3}{@ {}l}{Variables}\\
& @{text a} & type variables\\
& @{text c} & coercion variables\\
& @{text x} & term variables\\
\end{tabular}
\end{center}
\end{boxedminipage}
\caption{The term-language of System @{text "F\<^isub>C"}
\cite{CoreHaskell}, also often referred to as \emph{Core-Haskell}. In this
version of the term-language we made a modification by separating the
grammars for type kinds and coercion kinds, as well as for types and coercion
types. For this paper the interesting term-constructor is @{text "\<CASE>"},
which binds multiple type-, coercion- and term-variables.\label{corehas}}
\end{figure}
*}
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} (including proofs). We shall briefly review this work
to aid the description of what follows.
Two central notions in the nominal logic work are sorted atoms and
sort-respecting permutations of atoms. We will use the letters @{text "a,
b, c, \<dots>"} to stand for atoms and @{text "p, q, \<dots>"} to stand for
permutations. The sorts of atoms can be used to represent different kinds of
variables, such as the term-, coercion- and type-variables in Core-Haskell.
It is assumed that there is an infinite supply of atoms for each
sort. However, in the interest of brevity, we shall restrict ourselves
in what follows to 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
%
\begin{center}
@{text "_ \<bullet> _ :: perm \<Rightarrow> \<beta> \<Rightarrow> \<beta>"}
\end{center}
\noindent
in which the generic type @{text "\<beta>"} stands for the type of the object
over 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 by induction over the type-hierarchy \cite{HuffmanUrban10};
for example permutations acting on products, lists, sets, functions and booleans is
given by:
%
\begin{equation}\label{permute}
\mbox{\begin{tabular}{@ {}cc@ {}}
\begin{tabular}{@ {}l@ {}}
@{thm permute_prod.simps[no_vars, THEN eq_reflection]}\\[2mm]
@{thm permute_list.simps(1)[no_vars, THEN eq_reflection]}\\
@{thm permute_list.simps(2)[no_vars, THEN eq_reflection]}\\
\end{tabular} &
\begin{tabular}{@ {}l@ {}}
@{thm permute_set_eq[no_vars, THEN eq_reflection]}\\
@{text "p \<bullet> f \<equiv> \<lambda>x. p \<bullet> (f (- p \<bullet> x))"}\\
@{thm permute_bool_def[no_vars, THEN eq_reflection]}\\
\end{tabular}
\end{tabular}}
\end{equation}
\noindent
Concrete permutations in Nominal Isabelle are built up from swappings,
written as \mbox{@{text "(a b)"}}, which are permutations that behave
as follows:
%
\begin{center}
@{text "(a b) = \<lambda>c. if a = c then b else if b = c then a else c"}
\end{center}
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:
%
\begin{equation}\label{suppdef}
@{thm supp_def[no_vars, THEN eq_reflection]}
\end{equation}
\noindent
There is also the derived notion for when an atom @{text a} is \emph{fresh}
for an @{text x}, defined as
%
\begin{center}
@{thm fresh_def[no_vars]}
\end{center}
\noindent
We 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}, leaves
@{text x} unchanged:
\begin{property}\label{swapfreshfresh}
@{thm[mode=IfThen] swap_fresh_fresh[no_vars]}
\end{property}
While often the support of an object can be relatively easily
described, for example for atoms, products, lists, function applications,
booleans and permutations as follows\\[-6mm]
%
\begin{eqnarray}
@{term "supp a"} & = & @{term "{a}"}\\
@{term "supp (x, y)"} & = & @{term "supp x \<union> supp y"}\\
@{term "supp []"} & = & @{term "{}"}\\
@{term "supp (x#xs)"} & = & @{term "supp x \<union> supp xs"}\\
@{text "supp (f x)"} & @{text "\<subseteq>"} & @{term "supp f \<union> supp x"}\label{suppfun}\\
@{term "supp b"} & = & @{term "{}"}\\
@{term "supp p"} & = & @{term "{a. p \<bullet> a \<noteq> a}"}
\end{eqnarray}
\noindent
in some cases it can be difficult to characterise the support precisely, and
only an approximation can be established (see \eqref{suppfun} above). Reasoning about
such approximations can be simplified with the notion \emph{supports}, defined
as follows:
\begin{defn}
A set @{text S} \emph{supports} @{text x} if for all atoms @{text a} and @{text b}
not in @{text S} we have @{term "(a \<rightleftharpoons> b) \<bullet> x = x"}.
\end{defn}
\noindent
The main point of @{text supports} is that we can establish the following
two properties.
\begin{property}\label{supportsprop}
{\it i)} @{thm[mode=IfThen] supp_is_subset[no_vars]}\\
{\it ii)} @{thm supp_supports[no_vars]}.
\end{property}
Another important notion in the nominal logic work is \emph{equivariance}.
For a function @{text f}, say of type @{text "\<alpha> \<Rightarrow> \<beta>"}, to be equivariant
it is required that every permutation leaves @{text f} unchanged, that is
%
\begin{equation}\label{equivariancedef}
@{term "\<forall>p. p \<bullet> f = f"}
\end{equation}
\noindent or equivalently that a permutation applied to the application
@{text "f x"} can be moved to the argument @{text x}. That means for equivariant
functions @{text f}, we have for all permutations @{text p}
%
\begin{equation}\label{equivariance}
@{text "p \<bullet> f = f"} \;\;\;\textit{if and only if}\;\;\;
@{text "p \<bullet> (f x) = f (p \<bullet> x)"}
\end{equation}
\noindent
From property \eqref{equivariancedef} and the definition of @{text supp}, we
can easily deduce that equivariant functions have empty support. There is
also a similar notion for equivariant relations, say @{text R}, namely the property
that
%
\begin{center}
@{text "x R y"} \;\;implies\;\; @{text "(p \<bullet> x) R (p \<bullet> y)"}
\end{center}
Finally, the nominal logic work provides us with general means to rename
binders. While in the older version of Nominal Isabelle, we used extensively
Property~\ref{swapfreshfresh} for renaming single binders, this property
proved unwieldy for dealing with multiple binders. For such binders the
following generalisations turned out to be easier to use.
\begin{property}\label{supppermeq}
@{thm[mode=IfThen] supp_perm_eq[no_vars]}
\end{property}
\begin{property}\label{avoiding}
For a finite set @{text as} and a finitely supported @{text x} with
@{term "as \<sharp>* x"} and also a finitely supported @{text c}, there
exists a permutation @{text p} such that @{term "(p \<bullet> as) \<sharp>* c"} and
@{term "supp x \<sharp>* p"}.
\end{property}
\noindent
The idea behind the second property is that given a finite set @{text as}
of binders (being bound, or fresh, in @{text x} is ensured by the
assumption @{term "as \<sharp>* x"}), then there exists a permutation @{text p} such that
the renamed binders @{term "p \<bullet> as"} avoid @{text c} (which can be arbitrarily chosen
as long as it is finitely supported) and also @{text "p"} does not affect anything
in the support of @{text x} (that is @{term "supp x \<sharp>* p"}). The last
fact and Property~\ref{supppermeq} allow us to ``rename'' just the binders
@{text as} in @{text x}, because @{term "p \<bullet> x = x"}.
Most properties given in this section are described in detail in \cite{HuffmanUrban10}
and of course all are formalised in Isabelle/HOL. In the next sections we will make
extensive use of these properties in order to define alpha-equivalence in
the presence of multiple binders.
*}
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 with deriving the reasoning infrastructure
manageable, we will wherever possible state definitions and perform proofs
on the user-level of Isabelle/HOL, as opposed to write custom ML-code that
generates them anew for each specification. To that end, we will consider
first pairs @{text "(as, x)"} of type @{text "(atom set) \<times> \<beta>"}. These pairs
are intended to represent the abstraction, or binding, of the set of atoms @{text
"as"} in the body @{text "x"}.
The first question we have to answer is when two pairs @{text "(as, x)"} and
@{text "(bs, y)"} are alpha-equivalent? (For 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 @{text "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)} @{text p} 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 equivalent terms. We also require that {\it iv)}
@{text p} makes the sets of abstracted atoms @{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}{@{term "(as, x) \<approx>gen R fv p (bs, y)"}\hspace{2mm}@{text "\<equiv>"}\hspace{30mm}}\\[1mm]
& @{term "fv(x) - as = fv(y) - bs"}\\
@{text "\<and>"} & @{term "(fv(x) - as) \<sharp>* p"}\\
@{text "\<and>"} & @{text "(p \<bullet> x) R y"}\\
@{text "\<and>"} & @{term "(p \<bullet> as) = bs"}\\
\end{array}
\end{equation}
\noindent
Note that this relation depends on the permutation @{text
"p"}. Alpha-equivalence between two pairs is then the relation where we
existentially quantify over this @{text "p"}. Also note that the relation is
dependent on a free-variable function @{text "fv"} and a relation @{text
"R"}. The reason for this extra generality is that we will use
$\approx_{\textit{set}}$ for both ``raw'' terms and alpha-equated terms. In
the latter case, $R$ will be replaced by equality @{text "="} and we
will prove that @{text "fv"} is equal to @{text "supp"}.
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}{@{term "(as, x) \<approx>lst R fv p (bs, y)"}\hspace{2mm}@{text "\<equiv>"}\hspace{30mm}}\\[1mm]
& @{term "fv(x) - (set as) = fv(y) - (set bs)"}\\
\wedge & @{term "(fv(x) - set as) \<sharp>* p"}\\
\wedge & @{text "(p \<bullet> x) R y"}\\
\wedge & @{term "(p \<bullet> as) = bs"}\\
\end{array}
\end{equation}
\noindent
where @{term set} is a function that coerces a list of atoms into a set of atoms.
Now the last clause ensures that the order of the binders matters (since @{text as}
and @{text bs} are lists of atoms).
If we do not want to make any difference between the order of binders \emph{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}{@{term "(as, x) \<approx>res R fv p (bs, y)"}\hspace{2mm}@{text "\<equiv>"}\hspace{30mm}}\\[1mm]
& @{term "fv(x) - as = fv(y) - bs"}\\
\wedge & @{term "(fv(x) - as) \<sharp>* p"}\\
\wedge & @{text "(p \<bullet> x) R y"}\\
\end{array}
\end{equation}
It might be useful to consider some examples for how these definitions of alpha-equivalence
pan out in practice.
For this consider the case of abstracting a set of variables over types (as in type-schemes).
We set @{text R} to be the usual equality @{text "="} and for @{text "fv(T)"} we define
\begin{center}
@{text "fv(x) = {x}"} \hspace{5mm} @{text "fv(T\<^isub>1 \<rightarrow> T\<^isub>2) = fv(T\<^isub>1) \<union> fv(T\<^isub>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 alpha-equivalent according to
$\approx_{\textit{set}}$ and $\approx_{\textit{res}}$ by taking @{text p} to
be the swapping @{term "(x \<rightleftharpoons> y)"}. In case of @{text "x \<noteq> y"}, then @{text
"([x, y], x \<rightarrow> y)"} $\not\approx_{\textit{list}}$ @{text "([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 example is @{text "({x}, x)"} $\approx_{\textit{res}}$
@{text "({x, y}, x)"} which holds by taking @{text p} to be the identity
permutation. However, if @{text "x \<noteq> y"}, then @{text "({x}, x)"}
$\not\approx_{\textit{set}}$ @{text "({x, y}, x)"} since there is no
permutation that makes the sets @{text "{x}"} and @{text "{x, y}"} equal
(similarly for $\approx_{\textit{list}}$). It can also relatively easily be
shown that all three notions of alpha-equivalence coincide, if we only
abstract a single atom.
In the rest of this section we are going to introduce three abstraction
types. For this we define
%
\begin{equation}
@{term "abs_set (as, x) (bs, x) \<equiv> \<exists>p. alpha_gen (as, x) equal supp p (bs, x)"}
\end{equation}
\noindent
(similarly for $\approx_{\textit{abs\_res}}$
and $\approx_{\textit{abs\_list}}$). We can show that these relations are equivalence
relations and equivariant.
\begin{lemma}\label{alphaeq}
The relations $\approx_{\textit{abs\_set}}$, $\approx_{\textit{abs\_list}}$
and $\approx_{\textit{abs\_res}}$ are equivalence relations, and if @{term
"abs_set (as, x) (bs, y)"} then also @{term "abs_set (p \<bullet> as, p \<bullet> x) (p \<bullet>
bs, p \<bullet> y)"} (similarly for the other two relations).
\end{lemma}
\begin{proof}
Reflexivity is by taking @{text "p"} to be @{text "0"}. For symmetry we have
a permutation @{text p} and for the proof obligation take @{term "-p"}. In case
of transitivity, we have two permutations @{text p} and @{text q}, and for the
proof obligation use @{text "q + p"}. All conditions are then by simple
calculations.
\end{proof}
\noindent
This lemma allows us to use our quotient package and introduce
new types @{text "\<beta> abs_set"}, @{text "\<beta> abs_res"} and @{text "\<beta> abs_list"}
representing alpha-equivalence classes of pairs of type
@{text "(atom set) \<times> \<beta>"} (in the first two cases) and of type @{text "(atom list) \<times> \<beta>"}
(in the third case).
The elements in these types will be, respectively, written as:
\begin{center}
@{term "Abs as x"} \hspace{5mm}
@{term "Abs_res as x"} \hspace{5mm}
@{term "Abs_lst as x"}
\end{center}
\noindent
indicating that a set (or list) of atoms @{text as} is abstracted in @{text x}. We will
call the types \emph{abstraction types} and their elements
\emph{abstractions}. The important property we need to derive is the support of
abstractions, namely:
\begin{thm}[Support of Abstractions]\label{suppabs}
Assuming @{text x} has finite support, then\\[-6mm]
\begin{center}
\begin{tabular}{l@ {\hspace{2mm}}c@ {\hspace{2mm}}l}
@{thm (lhs) supp_abs(1)[no_vars]} & $=$ & @{thm (rhs) supp_abs(1)[no_vars]}\\
@{thm (lhs) supp_abs(2)[no_vars]} & $=$ & @{thm (rhs) supp_abs(2)[no_vars]}\\
@{thm (lhs) supp_abs(3)[where bs="as", no_vars]} & $=$ & @{thm (rhs) supp_abs(3)[where bs="as", no_vars]}
\end{tabular}
\end{center}
\end{thm}
\noindent
Below we will show the first equation. The others
follow by similar arguments. By definition of the abstraction type @{text "abs_set"}
we have
%
\begin{equation}\label{abseqiff}
@{thm (lhs) abs_eq_iff(1)[where bs="as" and cs="bs", no_vars]} \;\;\text{if and only if}\;\;
@{thm (rhs) abs_eq_iff(1)[where bs="as" and cs="bs", no_vars]}
\end{equation}
\noindent
and also
%
\begin{equation}\label{absperm}
@{thm permute_Abs[no_vars]}
\end{equation}
\noindent
The second fact derives from the definition of permutations acting on pairs
\eqref{permute} and alpha-equivalence being equivariant
(see Lemma~\ref{alphaeq}). With these two facts at our disposal, we can show
the following lemma about swapping two atoms.
\begin{lemma}
@{thm[mode=IfThen] abs_swap1(1)[where bs="as", no_vars]}
\end{lemma}
\begin{proof}
This lemma is straightforward using \eqref{abseqiff} and observing that
the assumptions give us @{term "(a \<rightleftharpoons> b) \<bullet> (supp x - as) = (supp x - as)"}.
Moreover @{text supp} and set difference are equivariant (see \cite{HuffmanUrban10}).
\end{proof}
\noindent
Assuming that @{text "x"} has finite support, this lemma together
with \eqref{absperm} allows us to show
%
\begin{equation}\label{halfone}
@{thm abs_supports(1)[no_vars]}
\end{equation}
\noindent
which by Property~\ref{supportsprop} gives us ``one half'' of
Theorem~\ref{suppabs}. The ``other half'' is a bit more involved. To establish
it, we use a trick from \cite{Pitts04} and first define an auxiliary
function @{text aux}, taking an abstraction as argument:
%
\begin{center}
@{thm supp_gen.simps[THEN eq_reflection, no_vars]}
\end{center}
\noindent
Using the second equation in \eqref{equivariance}, we can show that
@{text "aux"} is equivariant (since @{term "p \<bullet> (supp x - as) =
(supp (p \<bullet> x)) - (p \<bullet> as)"}) and therefore has empty support.
This in turn means
%
\begin{center}
@{term "supp (supp_gen (Abs as x)) \<subseteq> supp (Abs as x)"}
\end{center}
\noindent
using \eqref{suppfun}. Assuming @{term "supp x - as"} is a finite set,
we further obtain
%
\begin{equation}\label{halftwo}
@{thm (concl) supp_abs_subset1(1)[no_vars]}
\end{equation}
\noindent
since for finite sets of atoms, @{text "bs"}, we have
@{thm (concl) supp_finite_atom_set[where S="bs", no_vars]}.
Finally, taking \eqref{halfone} and \eqref{halftwo} together establishes
Theorem~\ref{suppabs}.
The method of first considering abstractions of the
form @{term "Abs as x"} etc is motivated by the fact that
we can conveniently establish at the Isabelle/HOL level
properties about them. It would be
laborious to write custom ML-code that derives automatically such properties
for every term-constructor that binds some atoms. Also the generality of
the definitions for alpha-equivalence will help us in the next section.
*}
section {* Alpha-Equivalence and Free Variables\label{sec:alpha} *}
text {*
Our choice of syntax for specifications is influenced by the existing
datatype package of Isabelle/HOL \cite{Berghofer99} and by the syntax of the
Ott-tool \cite{ott-jfp}. For us a specification of a term-calculus is a
collection of (possibly mutual recursive) type declarations, say @{text
"ty\<AL>\<^isub>1, \<dots>, ty\<AL>\<^isub>n"}, and an associated collection of
binding functions, say @{text "bn\<AL>\<^isub>1, \<dots>, bn\<AL>\<^isub>m"}. The
syntax in Nominal Isabelle for such specifications is roughly as follows:
%
\begin{equation}\label{scheme}
\mbox{\begin{tabular}{@ {\hspace{-5mm}}p{1.8cm}l}
type \mbox{declaration part} &
$\begin{cases}
\mbox{\begin{tabular}{l}
\isacommand{nominal\_datatype} @{text "ty\<AL>\<^isub>1 = \<dots>"}\\
\isacommand{and} @{text "ty\<AL>\<^isub>2 = \<dots>"}\\
$\ldots$\\
\isacommand{and} @{text "ty\<AL>\<^isub>n = \<dots>"}\\
\end{tabular}}
\end{cases}$\\
binding \mbox{function part} &
$\begin{cases}
\mbox{\begin{tabular}{l}
\isacommand{binder} @{text "bn\<AL>\<^isub>1"} \isacommand{and} \ldots \isacommand{and} @{text "bn\<AL>\<^isub>m"}\\
\isacommand{where}\\
$\ldots$\\
\end{tabular}}
\end{cases}$\\
\end{tabular}}
\end{equation}
\noindent
Every type declaration @{text ty}$^\alpha_{1..n}$ consists of a collection of
term-constructors, each of which come with a list of labelled
types that stand for the types of the arguments of the term-constructor.
For example a term-constructor @{text "C\<^sup>\<alpha>"} might be specified with
\begin{center}
@{text "C\<^sup>\<alpha> label\<^isub>1::ty"}$'_1$ @{text "\<dots> label\<^isub>l::ty"}$'_l\;\;$ @{text "binding_clauses"}
\end{center}
\noindent
whereby some of the @{text ty}$'_{1..l}$ (or their components) can be contained
in the collection of @{text ty}$^\alpha_{1..n}$ declared in
\eqref{scheme}.
In this case we will call the corresponding argument a
\emph{recursive argument} of @{text "C\<^sup>\<alpha>"}.
%The types of such recursive
%arguments need to satisfy a ``positivity''
%restriction, which ensures that the type has a set-theoretic semantics
%\cite{Berghofer99}.
The labels
annotated on the types are optional. Their purpose is to be used in the
(possibly empty) list of \emph{binding clauses}, which indicate the binders
and their scope in a term-constructor. They come in three \emph{modes}:
\begin{center}
\begin{tabular}{l}
\isacommand{bind}\; {\it binders}\; \isacommand{in}\; {\it labels}\\
\isacommand{bind\_set}\; {\it binders}\; \isacommand{in}\; {\it labels}\\
\isacommand{bind\_res}\; {\it binders}\; \isacommand{in}\; {\it labels}\\
\end{tabular}
\end{center}
\noindent
The first mode is for binding lists of atoms (the order of binders matters);
the second is for sets of binders (the order does not matter, but the
cardinality does) and the last is for sets of binders (with vacuous binders
preserving alpha-equivalence). The ``\isacommand{in}-part'' of a binding
clause will be called \emph{bodies} (there can be more than one); the
``\isacommand{bind}-part'' will be called \emph{binders}. In contrast to
Ott, we allow multiple labels in binders and bodies. For example we allow
binding clauses of the form:
\begin{center}
\begin{tabular}{@ {}ll@ {}}
@{text "Foo\<^isub>1 x::name y::name t::lam s::lam"} &
\isacommand{bind} @{text "x y"} \isacommand{in} @{text "t s"}\\
@{text "Foo\<^isub>2 x::name y::name t::lam s::lam"} &
\isacommand{bind} @{text "x y"} \isacommand{in} @{text "t"},
\isacommand{bind} @{text "x y"} \isacommand{in} @{text "s"}\\
\end{tabular}
\end{center}
\noindent
Similarly for the other binding modes. Interestingly, in case of \isacommand{bind\_set}
and \isacommand{bind\_res} the binding clauses will make a difference to the semantics
of the specification (the corresponding alpha-equivalence will differ). We will
show this later with an example.
There are some restrictions we need to impose on binding clasues: First, a
body can only occur in \emph{one} binding clause of a term constructor. For
binders we distinguish between \emph{shallow} and \emph{deep} binders.
Shallow binders are just labels. The restriction we need to impose on them
is that in case of \isacommand{bind\_set} and \isacommand{bind\_res} the
labels must either refer to atom types or to sets of atom types; in case of
\isacommand{bind} the labels must refer to atom types or lists of atom
types. Two examples for the use of shallow binders are the specification of
lambda-terms, where a single name is bound, and type-schemes, where a finite
set of names is bound:
\begin{center}
\begin{tabular}{@ {}cc@ {}}
\begin{tabular}{@ {}l@ {\hspace{-1mm}}}
\isacommand{nominal\_datatype} @{text lam} =\\
\hspace{5mm}\phantom{$\mid$}~@{text "Var name"}\\
\hspace{5mm}$\mid$~@{text "App lam lam"}\\
\hspace{5mm}$\mid$~@{text "Lam x::name t::lam"}\\
\hspace{21mm}\isacommand{bind} @{text x} \isacommand{in} @{text t}\\
\end{tabular} &
\begin{tabular}{@ {}l@ {}}
\isacommand{nominal\_datatype}~@{text ty} =\\
\hspace{5mm}\phantom{$\mid$}~@{text "TVar name"}\\
\hspace{5mm}$\mid$~@{text "TFun ty ty"}\\
\isacommand{and}~@{text "tsc = All xs::(name fset) T::ty"}\\
\hspace{24mm}\isacommand{bind\_res} @{text xs} \isacommand{in} @{text T}\\
\end{tabular}
\end{tabular}
\end{center}
\noindent
Note that for @{text lam} it does not matter which binding mode we use. The
reason is that we bind only a single @{text name}. However, having
\isacommand{bind\_set} or \isacommand{bind} in the second case makes a
difference to the semantics of the specification. Note also that in
these specifications @{text "name"} refers to an atom type, and @{text
"fset"} to the type of finite sets.
A \emph{deep} binder uses an auxiliary binding function that ``picks'' out
the atoms in one argument of the term-constructor, which can be bound in
other arguments and also in the same argument (we will call such binders
\emph{recursive}, see below). The corresponding 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 a term-calculus containing @{text "Let"}s with
tuple patterns might be specified as:
%
\begin{equation}\label{letpat}
\mbox{%
\begin{tabular}{l}
\isacommand{nominal\_datatype} @{text trm} =\\
\hspace{5mm}\phantom{$\mid$}~@{term "Var name"}\\
\hspace{5mm}$\mid$~@{term "App trm trm"}\\
\hspace{5mm}$\mid$~@{text "Lam x::name t::trm"}
\;\;\isacommand{bind} @{text x} \isacommand{in} @{text t}\\
\hspace{5mm}$\mid$~@{text "Let p::pat trm t::trm"}
\;\;\isacommand{bind} @{text "bn(p)"} \isacommand{in} @{text t}\\
\isacommand{and} @{text pat} =\\
\hspace{5mm}\phantom{$\mid$}~@{text PNil}\\
\hspace{5mm}$\mid$~@{text "PVar name"}\\
\hspace{5mm}$\mid$~@{text "PTup pat pat"}\\
\isacommand{binder}~@{text "bn::pat \<Rightarrow> atom list"}\\
\isacommand{where}~@{text "bn(PNil) = []"}\\
\hspace{5mm}$\mid$~@{text "bn(PVar x) = [atom x]"}\\
\hspace{5mm}$\mid$~@{text "bn(PTup p\<^isub>1 p\<^isub>2) = bn(p\<^isub>1) @ bn(p\<^isub>2)"}\\
\end{tabular}}
\end{equation}
\noindent
In this specification the function @{text "bn"} determines which atoms of
@{text p} are bound in the argument @{text "t"}. Note that in the
second-last @{text bn}-clause the function @{text "atom"} coerces a name
into the generic atom type of Nominal Isabelle \cite{HuffmanUrban10}. This
allows us to treat binders of different atom type uniformly.
As said above, for deep binders we allow binding clauses such as
%
\begin{center}
\begin{tabular}{ll}
@{text "Bar p::pat t::trm"} &
\isacommand{bind} @{text "bn(p)"} \isacommand{in} @{text "p t"}\\
\end{tabular}
\end{center}
\noindent
where the argument of the deep binder is also in the body. We call such
binders \emph{recursive}. To see the purpose of such recursive binders,
compare ``plain'' @{text "Let"}s and @{text "Let_rec"}s in the following
specification:
\begin{equation}\label{letrecs}
\mbox{%
\begin{tabular}{@ {}l@ {}}
\isacommand{nominal\_datatype}~@{text "trm ="}\\
\hspace{5mm}\phantom{$\mid$}\ldots\\
\hspace{5mm}$\mid$~@{text "Let as::assn t::trm"}
\;\;\isacommand{bind} @{text "bn(as)"} \isacommand{in} @{text t}\\
\hspace{5mm}$\mid$~@{text "Let_rec as::assn t::trm"}
\;\;\isacommand{bind} @{text "bn(as)"} \isacommand{in} @{text "as t"}\\
\isacommand{and} @{text "ass"} =\\
\hspace{5mm}\phantom{$\mid$}~@{text "ANil"}\\
\hspace{5mm}$\mid$~@{text "ACons name trm assn"}\\
\isacommand{binder} @{text "bn::assn \<Rightarrow> atom list"}\\
\isacommand{where}~@{text "bn(ANil) = []"}\\
\hspace{5mm}$\mid$~@{text "bn(ACons a t as) = [atom a] @ bn(as)"}\\
\end{tabular}}
\end{equation}
\noindent
The difference is that with @{text Let} we only want to bind the atoms @{text
"bn(as)"} in the term @{text t}, but with @{text "Let_rec"} we also want to bind the atoms
inside the assignment. This difference has consequences for the free-variable
function and alpha-equivalence relation, which we are going to define below.
For this definition, we have to impose some restrictions on deep binders. First,
we cannot have more than one binding function for one binder. So we exclude:
\begin{center}
\begin{tabular}{ll}
@{text "Baz p::pat t::trm"} &
\isacommand{bind} @{text "bn\<^isub>1(p) bn\<^isub>2(p)"} \isacommand{in} @{text t}\\
\end{tabular}
\end{center}
\noindent
The reason why we must exclude such specifications is that they cannot be
represented by the general binders described in Section~\ref{sec:binders}.
We also need to restrict the form of the binding functions. As will shortly
become clear, we cannot return an atom in a binding function that is also
bound in the corresponding term-constructor. That means in the example
\eqref{letpat} that the term-constructors @{text PVar} and @{text PTup} may
not have a binding clause (all arguments are used to define @{text "bn"}).
In contrast, in case of \eqref{letrecs} the term-constructor @{text ACons}
may have binding clause involving the argument @{text t} (the only one that
is \emph{not} used in the definition of the binding function). In the version of
Nominal Isabelle described here, we also adopted the restriction from the
Ott-tool that binding functions can only return: the empty set or empty list
(as in case @{text PNil}), a singleton set or singleton list containing an
atom (case @{text PVar}), or unions of atom sets or appended atom lists
(case @{text PTup}). This restriction will simplify some automatic proofs
later on.
In order to simplify our definitions, we shall assume specifications
of term-calculi are \emph{completed}. By this we mean that
for every argument of a term-constructor that is \emph{not}
already part of a binding clause given by the user, we add implicitly a special \emph{empty} binding
clause, written \isacommand{bind}~@{term "{}"}~\isacommand{in}~@{text "labels"}. In case
of the lambda-calculus, the completion produces
\begin{center}
\begin{tabular}{@ {}l@ {\hspace{-1mm}}}
\isacommand{nominal\_datatype} @{text lam} =\\
\hspace{5mm}\phantom{$\mid$}~@{text "Var x::name"}
\;\;\isacommand{bind}~@{term "{}"}~\isacommand{in}~@{text "x"}\\
\hspace{5mm}$\mid$~@{text "App t\<^isub>1::lam t\<^isub>2::lam"}
\;\;\isacommand{bind}~@{term "{}"}~\isacommand{in}~@{text "t\<^isub>1 t\<^isub>2"}\\
\hspace{5mm}$\mid$~@{text "Lam x::name t::lam"}
\;\;\isacommand{bind}~@{text x} \isacommand{in} @{text t}\\
\end{tabular}
\end{center}
\noindent
The point of completion is that we can make definitions over the binding
clauses and be sure to have captured all arguments of a term constructor.
Having dealt with all syntax matters, the problem now is how we can turn
specifications into actual type definitions in Isabelle/HOL and then
establish a reasoning infrastructure for them. As
Pottier and Cheney pointed out \cite{Pottier06,Cheney05}, just
re-arranging the arguments of
term-constructors so that binders and their bodies are next to each other will
result in inadequate representations. Therefore we will first
extract datatype definitions from the specification and then define
explicitly an alpha-equivalence relation over them.
The datatype definition can be obtained by stripping off the
binding clauses and the labels from the types. We also have to invent
new names for the types @{text "ty\<^sup>\<alpha>"} and term-constructors @{text "C\<^sup>\<alpha>"}
given by the user. In our implementation we just use the affix ``@{text "_raw"}''.
But for the purpose of this paper, we use the superscript @{text "_\<^sup>\<alpha>"} to indicate
that a notion is defined over alpha-equivalence classes and leave it out
for the corresponding notion defined on the ``raw'' level. So for example
we have
\begin{center}
@{text "ty\<^sup>\<alpha> \<mapsto> ty"} \hspace{2mm}and\hspace{2mm} @{text "C\<^sup>\<alpha> \<mapsto> C"}
\end{center}
\noindent
where @{term ty} is the type used in the quotient construction for
@{text "ty\<^sup>\<alpha>"} and @{text "C"} is the term-constructor on the ``raw'' type @{text "ty"}.
The resulting datatype definition is legal in Isabelle/HOL provided the datatypes are
non-empty and the types in the constructors only occur in positive
position (see \cite{Berghofer99} for an indepth description of the datatype package
in Isabelle/HOL). We then define the user-specified binding
functions, called @{term "bn"}, by primitive recursion over the corresponding
raw datatype. We can also easily define permutation operations by
primitive recursion so that for each term constructor @{text "C ty\<^isub>1 \<dots> ty\<^isub>n"}
we have that
%
\begin{equation}\label{ceqvt}
@{text "p \<bullet> (C z\<^isub>1 \<dots> z\<^isub>n) = C (p \<bullet> z\<^isub>1) \<dots> (p \<bullet> z\<^isub>n)"}
\end{equation}
The first non-trivial step we have to perform is the generation of free-variable
functions from the specifications. For atomic types we define the auxilary
free variable functions:
\begin{center}
\begin{tabular}{r@ {\hspace{2mm}}c@ {\hspace{2mm}}l}
@{text "fv_atom a"} & @{text "="} & @{text "atom a"}\\
@{text "fv_atom_set as"} & @{text "="} & @{text "atoms as"}\\
@{text "fv_atom_list as"} & @{text "="} & @{text "atoms (set as)"}\\
\end{tabular}
\end{center}
\noindent
Like the coercion function @{text atom}, @{text "atoms"} coerces
the set of atoms to a set of the generic atom type.
It is defined as @{text "atoms as \<equiv> {atom a | a \<in> as}"}.
Given the raw types @{text "ty\<^isub>1 \<dots> ty\<^isub>n"} we define now the
free-variable functions
\begin{center}
@{text "fv_ty\<^isub>1, \<dots>, fv_ty\<^isub>n"}
\end{center}
\noindent
We define them together with auxiliary free-variable functions for
the binding functions. Given binding functions
@{text "bn\<^isub>1 \<dots> bn\<^isub>m"} we define
%
\begin{center}
@{text "fv_bn\<^isub>1, \<dots>, fv_bn\<^isub>m"}
\end{center}
\noindent
The reason for this setup is that in a deep binder not all atoms have to be
bound, as we saw in the example with ``plain'' @{text Let}s. We need therefore a function
that calculates those unbound atoms.
While the idea behind these free-variable functions is clear (they just
collect all atoms that are not bound), because of our rather complicated
binding mechanisms their definitions are somewhat involved. Given a
term-constructor @{text "C"} of type @{text ty} and some associated binding
clauses @{text bcs}, the result of the @{text "fv_ty (C z\<^isub>1 \<dots> z\<^isub>n)"} will be
the union of the values, @{text v}, calculated by the rules for empty, shallow and
deep binding clauses:
%
\begin{equation}\label{deepbinderA}
\mbox{%
\begin{tabular}{c@ {\hspace{2mm}}p{7.4cm}}
\multicolumn{2}{@ {}l}{\isacommand{bind}~@{term "{}"}~\isacommand{in}~@{text "y"}:}\\
$\bullet$ & @{term "v = fv_ty y"} provided @{text "y"} is of type @{text "ty"}\smallskip\\
%
\multicolumn{2}{@ {}l}{\isacommand{bind}~@{text "x\<^isub>1\<dots>x\<^isub>n"}~\isacommand{in}~@{text "y\<^isub>1\<dots>y\<^isub>m"}:}\\
$\bullet$ & @{text "v = (fv_ty\<^isub>1 y\<^isub>1 \<union> .. \<union> fv_ty\<^isub>m y\<^isub>m)"}\\
& \hspace{15mm}@{text "- (fv_aty\<^isub>1 x\<^isub>1 \<union> .. \<union> fv_aty\<^isub>n x\<^isub>n)"}\\
& provided the bodies @{text "y"}$_{1..m}$ are of type @{text "ty"}$_{1..m}$ and the
binders @{text "x"}$_{1..n}$ of atomic types @{text "aty"}$_{1..n}$\smallskip\\
\end{tabular}}
\end{equation}
\begin{equation}\label{deepbinderB}
\mbox{%
\begin{tabular}{c@ {\hspace{2mm}}p{7.4cm}}
\multicolumn{2}{@ {}l}{\isacommand{bind}~@{text "bn x"}~\isacommand{in}~@{text "y\<^isub>1\<dots>y\<^isub>m"}:}\\
$\bullet$ & @{text "v = (fv_ty\<^isub>1 y\<^isub>1 \<union> .. \<union> fv_ty\<^isub>m y\<^isub>m) - set (bn x) \<union> (fv_bn x)"}\\
$\bullet$ & @{text "v = (fv_ty\<^isub>1 y\<^isub>1 \<union> .. \<union> fv_ty\<^isub>m y\<^isub>m) - set (bn x)"}\\
& provided the @{text "y"}$_{1..m}$ are of type @{text "ty"}$_{1..m}$; the first
clause applies for @{text x} being a non-recursive deep binder (that is
@{text "x \<noteq> y"}$_{1..m}$), the second for a recursive deep binder
\end{tabular}}
\end{equation}
\noindent
Similarly for the other binding modes. Note that in a non-recursive deep
binder, we have to add all atoms that are left unbound by the binding
function @{text "bn"}. For this we define the function @{text "fv_bn"}. Assume
for the constructor @{text "C"} the binding function is of the form @{text
"bn (C z\<^isub>1 \<dots> z\<^isub>n) = rhs"}. We again define a value
@{text v} which is exactly as in \eqref{deepbinderA}/\eqref{deepbinderB} for shallow and deep
binding clauses, except for empty binding clauses are defined as follows:
%
\begin{equation}\label{bnemptybinder}
\mbox{%
\begin{tabular}{c@ {\hspace{2mm}}p{7cm}}
\multicolumn{2}{@ {}l}{\isacommand{bind}~@{term "{}"}~\isacommand{in}~@{text "y"}:}\\
$\bullet$ & @{term "v = fv_ty y"} provided @{text "y"} does not occur in @{text "rhs"}
and the free-variable function for the type of @{text "y"} is @{text "fv_ty"}\\
$\bullet$ & @{term "v = fv_bn y"} provided @{text "y"} occurs in @{text "rhs"}
with a recursive call @{text "bn y"}\\
$\bullet$ & @{term "v = {}"} provided @{text "y"} occurs in @{text "rhs"},
but without a recursive call\\
\end{tabular}}
\end{equation}
\noindent
The reason why we only have to treat the empty binding clauses specially in
the definition of @{text "fv_bn"} is that binding functions can only use arguments
that do not occur in binding clauses. Otherwise the @{text "bn"} function cannot
be lifted to alpha-equated terms.
To see how these definitions work in practice, let us reconsider the term-constructors
@{text "Let"} and @{text "Let_rec"} from the example shown in \eqref{letrecs}.
For this specification we need to define three free-variable functions, namely
@{text "fv\<^bsub>trm\<^esub>"}, @{text "fv\<^bsub>assn\<^esub>"} and @{text "fv\<^bsub>bn\<^esub>"}. They are as follows:
%
\begin{center}
\begin{tabular}{@ {}l@ {\hspace{1mm}}c@ {\hspace{1mm}}l@ {}}
@{text "fv\<^bsub>trm\<^esub> (Let as t)"} & @{text "="} & @{text "(fv\<^bsub>trm\<^esub> t - set (bn as)) \<union> fv\<^bsub>bn\<^esub> as"}\\
@{text "fv\<^bsub>trm\<^esub> (Let_rec as t)"} & @{text "="} & @{text "(fv\<^bsub>assn\<^esub> as \<union> fv\<^bsub>trm\<^esub> t) - set (bn as)"}\\[1mm]
@{text "fv\<^bsub>assn\<^esub> (ANil)"} & @{text "="} & @{term "{}"}\\
@{text "fv\<^bsub>assn\<^esub> (ACons a t as)"} & @{text "="} & @{text "{atom a} \<union> (fv\<^bsub>trm\<^esub> t) \<union> (fv\<^bsub>assn\<^esub> as)"}\\[1mm]
@{text "fv\<^bsub>bn\<^esub> (ANil)"} & @{text "="} & @{term "{}"}\\
@{text "fv\<^bsub>bn\<^esub> (ACons a t as)"} & @{text "="} & @{text "(fv\<^bsub>trm\<^esub> t) \<union> (fv\<^bsub>bn\<^esub> as)"}
\end{tabular}
\end{center}
\noindent
Since there are only (implicit) empty binding clauses for the term-constructors @{text ANil}
and @{text "ACons"}, the corresponding free-variable function @{text
"fv\<^bsub>assn\<^esub>"} returns all atoms occuring in an assignment. The
binding only takes place in @{text Let} and @{text "Let_rec"}. In the @{text
"Let"}-clause we want to bind all atoms given by @{text "set (bn as)"} in
@{text t}. Therefore we have to subtract @{text "set (bn as)"} from @{text
"fv\<^bsub>trm\<^esub> t"}. However, we also need to add all atoms that are
free in @{text "as"}. This is what the purpose of the function @{text
"fv\<^bsub>bn\<^esub>"} is. In contrast, in @{text "Let_rec"} we have a
recursive binder where we want to also bind all occurrences of the atoms
in @{text "set (bn as)"} inside @{text "as"}. Therefore we have to subtract @{text
"set (bn as)"} from @{text "fv\<^bsub>trm\<^esub> t"}, as well as from
@{text "fv\<^bsub>assn\<^esub> as"}. An interesting point in this example is
that an assignment ``alone'' does not have any bound variables. Only in the
context of a @{text Let} or @{text "Let_rec"} will some atoms become bound.
This is a phenomenon
that has also been pointed out in \cite{ott-jfp}. For us it is crucial, because
we would not be able to lift a @{text "bn"}-function that is defined over
arguments that are either binders or bodies. In that case @{text "bn"} would
not respect alpha-equivalence. We can also see that
%
\begin{equation}\label{bnprop}
@{text "fv_ty x = bn x \<union> fv_bn x"}.
\end{equation}
\noindent
holds for any @{text "bn"}-function defined for the type @{text "ty"}.
Next we define alpha-equivalence relations for the types @{text "ty\<^isub>1, \<dots>, ty\<^isub>n"}. We call them
@{text "\<approx>ty\<^isub>1, \<dots>, \<approx>ty\<^isub>n"}. Like with the free-variable functions,
we also need to define auxiliary alpha-equivalence relations for the binding functions.
Say we have binding functions @{text "bn\<^isub>1, \<dots>, bn\<^isub>m"}, then we also define @{text "\<approx>bn\<^isub>1, \<dots>, \<approx>bn\<^isub>m"}.
To simplify our definitions we will use the following abbreviations for
relations and free-variable acting on products.
%
\begin{center}
\begin{tabular}{r@ {\hspace{2mm}}c@ {\hspace{2mm}}l}
@{text "(x\<^isub>1, y\<^isub>1) (R\<^isub>1 \<otimes> R\<^isub>2) (x\<^isub>2, y\<^isub>2)"} & @{text "\<equiv>"} & @{text "x\<^isub>1 R\<^isub>1 y\<^isub>1 \<and> x\<^isub>2 R\<^isub>2 y\<^isub>2"}\\
@{text "(fv\<^isub>1 \<oplus> fv\<^isub>2) (x, y)"} & @{text "\<equiv>"} & @{text "fv\<^isub>1 x \<union> fv\<^isub>2 y"}\\
\end{tabular}
\end{center}
The relations for alpha-equivalence are inductively defined predicates, whose clauses have
conclusions of the form
%
\begin{center}
@{text "C x\<^isub>1 \<dots> x\<^isub>n \<approx>ty C y\<^isub>1 \<dots> y\<^isub>n"}
\end{center}
\noindent
For what follows, let us assume @{text C} is of type @{text ty}. The task
is to specify what the premises of these clauses are. Again for this we
analyse the binding clauses and produce a corresponding premise.
*}
(*<*)
consts alpha_ty ::'a
consts alpha_trm ::'a
consts fv_trm :: 'a
consts alpha_trm2 ::'a
consts fv_trm2 :: 'a
notation (latex output)
alpha_ty ("\<approx>ty") and
alpha_trm ("\<approx>\<^bsub>trm\<^esub>") and
fv_trm ("fv\<^bsub>trm\<^esub>") and
alpha_trm2 ("\<approx>\<^bsub>assn\<^esub> \<otimes> \<approx>\<^bsub>trm\<^esub>") and
fv_trm2 ("fv\<^bsub>assn\<^esub> \<oplus> fv\<^bsub>trm\<^esub>")
(*>*)
text {*
\begin{equation}\label{alpha}
\mbox{%
\begin{tabular}{c@ {\hspace{2mm}}p{7cm}}
\multicolumn{2}{@ {}l}{Empty binding clauses of the form
\isacommand{bind\_set}~@{term "{}"}~\isacommand{in}~@{text "x\<^isub>i"}:}\\
$\bullet$ & @{text "x\<^isub>i \<approx>ty y\<^isub>i"} provided @{text "x\<^isub>i"} and @{text "y\<^isub>i"}
are recursive arguments of @{text "C"}\\
$\bullet$ & @{term "x\<^isub>i = y\<^isub>i"} provided @{text "x\<^isub>i"} and @{text "y\<^isub>i"} are
non-recursive arguments\smallskip\\
\multicolumn{2}{@ {}l}{Shallow binders of the form
\isacommand{bind\_set}~@{text "x\<^isub>1\<dots>x\<^isub>n"}~\isacommand{in}~@{text "x'\<^isub>1\<dots>x'\<^isub>m"}:}\\
$\bullet$ & Assume the bodies @{text "x'\<^isub>1\<dots>x'\<^isub>m"} are of type @{text "ty\<^isub>1\<dots>ty\<^isub>m"},
@{text "R"} is @{text "\<approx>ty\<^isub>1 \<otimes> ... \<otimes> \<approx>ty\<^isub>m"} and @{text fv} is
@{text "fv_ty\<^isub>1 \<oplus> ... \<oplus> fv_ty\<^isub>m"}, then
\begin{center}
@{term "\<exists>p. (x\<^isub>1 \<union> \<xi> \<union> x\<^isub>n, (x'\<^isub>1,\<xi>,x'\<^isub>m)) \<approx>gen R fv p (y\<^isub>1 \<union> \<xi> \<union> y\<^isub>n, (y'\<^isub>1,\<xi>,y'\<^isub>m))"}
\end{center}\\
\multicolumn{2}{@ {}l}{Deep binders of the form
\isacommand{bind\_set}~@{text "bn x"}~\isacommand{in}~@{text "x'\<^isub>1\<dots>x'\<^isub>m"}:}\\
$\bullet$ & Assume the bodies @{text "x'\<^isub>1\<dots>x'\<^isub>m"} are of type @{text "ty\<^isub>1\<dots>ty\<^isub>m"},
@{text "R"} is @{text "\<approx>ty\<^isub>1 \<otimes> ... \<otimes> \<approx>ty\<^isub>m"} and @{text fv} is
@{text "fv_ty\<^isub>1 \<oplus> ... \<oplus> fv_ty\<^isub>m"}, then for recursive deep binders
\begin{center}
@{term "\<exists>p. (bn x, (x'\<^isub>1,\<xi>,x'\<^isub>m)) \<approx>gen R fv p (bn y, (y'\<^isub>1,\<xi>,y'\<^isub>m))"}
\end{center}\\
$\bullet$ & for non-recursive binders we generate in addition @{text "x \<approx>bn y"}\\
\end{tabular}}
\end{equation}
\noindent
Similarly for the other binding modes.
From this definition it is clear why we have to impose the restriction
of excluding overlapping deep binders, as these would need to be translated into separate
abstractions.
The alpha-equivalence relations @{text "\<approx>bn\<^isub>j"} for binding functions
are similar. We again have conclusions of the form \mbox{@{text "C x\<^isub>1 \<dots> x\<^isub>n \<approx>bn C y\<^isub>1 \<dots> y\<^isub>n"}}
and need to generate appropriate premises. We generate first premises according to the first three
rules given above. Only the ``left-over'' pairs @{text "(x\<^isub>i, y\<^isub>i)"} need to be treated
differently. They depend on whether @{text "x\<^isub>i"} occurs in @{text "rhs"} of the
clause @{text "bn (C x\<^isub>1 \<dots> x\<^isub>n) = rhs"}:
\begin{center}
\begin{tabular}{c@ {\hspace{2mm}}p{7cm}}
$\bullet$ & @{text "x\<^isub>i \<approx>ty y\<^isub>i"} provided @{text "x\<^isub>i"} does not occur in @{text rhs}
and the type of @{text "x\<^isub>i"} is @{text ty} and @{text "x\<^isub>i"} is a recursive argument
in the term-constructor\\
$\bullet$ & @{text "x\<^isub>i = y\<^isub>i"} provided @{text "x\<^isub>i"} does not occur in @{text rhs}
and @{text "x\<^isub>i"} is not a recursive argument in the term-constructor\\
$\bullet$ & @{text "x\<^isub>i \<approx>bn y\<^isub>i"} provided @{text "x\<^isub>i"} occurs in @{text rhs}
with the recursive call @{text "bn x\<^isub>i"}\\
$\bullet$ & none provided @{text "x\<^isub>i"} occurs in @{text rhs} but it is not
in a recursive call involving a @{text "bn"}
\end{tabular}
\end{center}
Again lets take a look at a concrete example for these definitions. For \eqref{letrecs}
we have three relations, namely $\approx_{\textit{trm}}$, $\approx_{\textit{assn}}$ and
$\approx_{\textit{bn}}$, with the clauses as follows:
\begin{center}
\begin{tabular}{@ {}c @ {}}
\infer{@{text "Let as t \<approx>\<^bsub>trm\<^esub> Let as' t'"}}
{@{text "as \<approx>\<^bsub>bn\<^esub> as'"} & @{term "\<exists>p. (bn as, t) \<approx>lst alpha_trm fv_trm p (bn as', t')"}}\smallskip\\
\infer{@{text "Let_rec as t \<approx>\<^bsub>trm\<^esub> Let_rec as' t'"}}
{@{term "\<exists>p. (bn as, (as, t)) \<approx>lst alpha_trm2 fv_trm2 p (bn as', (as', t'))"}}
\end{tabular}
\end{center}
\begin{center}
\begin{tabular}{@ {}c @ {}}
\infer{@{text "ANil \<approx>\<^bsub>assn\<^esub> ANil"}}{}\\
\infer{@{text "ACons a t as \<approx>\<^bsub>assn\<^esub> ACons a' t' as"}}
{@{text "a = a'"} & @{text "t \<approx>\<^bsub>trm\<^esub> t'"} & @{text "as \<approx>\<^bsub>assn\<^esub> as'"}}
\end{tabular}
\end{center}
\begin{center}
\begin{tabular}{@ {}c @ {}}
\infer{@{text "ANil \<approx>\<^bsub>bn\<^esub> ANil"}}{}\\
\infer{@{text "ACons a t as \<approx>\<^bsub>bn\<^esub> ACons a' t' as"}}
{@{text "t \<approx>\<^bsub>trm\<^esub> t'"} & @{text "as \<approx>\<^bsub>bn\<^esub> as'"}}
\end{tabular}
\end{center}
\noindent
Note the difference between $\approx_{\textit{assn}}$ and
$\approx_{\textit{bn}}$: the latter only ``tracks'' alpha-equivalence of
the components in an assignment that are \emph{not} bound. Therefore we have
a $\approx_{\textit{bn}}$-premise in the clause for @{text "Let"} (which is
a non-recursive binder). The underlying reason is that the terms inside an assignment are not meant
to be ``under'' the binder. Such a premise is \emph{not} needed in @{text "Let_rec"},
because there everything from the assignment is under the binder.
*}
section {* Establishing the Reasoning Infrastructure *}
text {*
Having made all necessary definitions for raw terms, we can start
introducing the reasoning infrastructure for the alpha-equated types the
user originally specified. We sketch in this section the facts we need for establishing
this reasoning infrastructure. First we have to show that the
alpha-equivalence relations defined in the previous section are indeed
equivalence relations.
\begin{lemma}\label{equiv}
Given the raw types @{text "ty\<^isub>1, \<dots>, ty\<^isub>n"} and binding functions
@{text "bn\<^isub>1, \<dots>, bn\<^isub>m"}, the relations @{text "\<approx>ty\<^isub>1, \<dots>, \<approx>ty\<^isub>n"} and
@{text "\<approx>bn\<^isub>1 \<dots> \<approx>bn\<^isub>m"} are equivalence relations and equivariant.
\end{lemma}
\begin{proof}
The proof is by mutual induction over the definitions. The non-trivial
cases involve premises built up by $\approx_{\textit{set}}$,
$\approx_{\textit{res}}$ and $\approx_{\textit{list}}$. They
can be dealt with as in Lemma~\ref{alphaeq}.
\end{proof}
\noindent
We can feed this lemma into our quotient package and obtain new types @{text
"ty\<AL>\<^bsub>1..n\<^esub>"} representing alpha-equated terms of types @{text "ty\<^bsub>1..n\<^esub>"}. We also obtain
definitions for the term-constructors @{text
"C"}$^\alpha_{1..m}$ from the raw term-constructors @{text
"C"}$_{1..m}$, and similar definitions for the free-variable functions @{text
"fv_ty"}$^\alpha_{1..n}$ and the binding functions @{text
"bn"}$^\alpha_{1..k}$. However, these definitions are not really useful to the
user, since they are given in terms of the isomorphisms we obtained by
creating new types in Isabelle/HOL (recall the picture shown in the
Introduction).
The first useful property to the user is the fact that term-constructors are
distinct, that is
%
\begin{equation}\label{distinctalpha}
\mbox{@{text "C"}$^\alpha_i$@{text "x\<^isub>1 \<dots> x\<^isub>n"} @{text "\<noteq>"}
@{text "C"}$^\alpha_j$@{text "y\<^isub>1 \<dots> y\<^isub>m"} holds for @{text "i \<noteq> j"}.}
\end{equation}
\noindent
By definition of alpha-equivalence we can show that
for the raw term-constructors
%
\begin{equation}\label{distinctraw}
\mbox{@{text "C\<^isub>i x\<^isub>1 \<dots> x\<^isub>n"}\;$\not\approx$@{text ty}\;@{text "C\<^isub>j y\<^isub>1 \<dots> y\<^isub>m"} holds for @{text "i \<noteq> j"}.}
\end{equation}
\noindent
In order to generate \eqref{distinctalpha} from \eqref{distinctraw}, the quotient
package needs to know that the term-constructors @{text "C\<^isub>i"} and @{text "C\<^isub>j"}
are \emph{respectful} w.r.t.~the alpha-equivalence relations (see \cite{Homeier05}).
Assuming @{text "C\<^isub>i"} is of type @{text "ty"} with argument types
@{text "ty\<^isub>1, \<dots>, ty\<^isub>n"}, then respectfulness amounts to showing that
\begin{center}
@{text "C\<^isub>i x\<^isub>1 \<dots> x\<^isub>n \<approx>ty C\<^isub>i y\<^isub>1 \<dots> y\<^isub>n"}
\end{center}
\noindent
are alpha-equivalent under the assumption that @{text "x\<^isub>i \<approx>ty\<^isub>i y\<^isub>i"} holds for all recursive arguments
and @{text "x\<^isub>i = y\<^isub>i"} holds for all non-recursive arguments of @{text "C\<^isub>i"}. We can prove this by
analysing the definition of @{text "\<approx>ty"}. For this proof to succeed we have to establish
the following auxiliary fact about binding functions. Given a binding function @{text bn\<^isub>i} defined
for the type @{text ty\<^isub>i}, we have that
%
\begin{center}
@{text "x \<approx>ty\<^isub>i y"} implies @{text "x \<approx>bn\<^isub>i y"}
\end{center}
\noindent
This can be established by induction on the definition of @{text "\<approx>ty\<^isub>i"}.
Having established respectfulness for every raw term-constructor, the
quotient package is able to automatically deduce \eqref{distinctalpha} from \eqref{distinctraw}.
In fact we can from now on lift facts from the raw level to the alpha-equated level
as long as they contain raw term-constructors only. The
induction principles derived by the datatype package in Isabelle/HOL for the types @{text
"ty\<^bsub>1..n\<^esub>"} fall into this category. So we can also add to our infrastructure
induction principles that establish
\begin{center}
@{text "P\<^bsub>ty\<AL>\<^esub>\<^isub>1 y\<^isub>1 \<dots> P\<^bsub>ty\<AL>\<^esub>\<^isub>n y\<^isub>n "}
\end{center}
\noindent
for all @{text "y\<^isub>i"} whereby the variables @{text "P\<^bsub>ty\<AL>\<^esub>\<^isub>i"} stand for properties
defined over the types @{text "ty\<AL>\<^isub>1 \<dots> ty\<AL>\<^isub>n"}. The premises of
these induction principles look
as follows
\begin{center}
@{text "\<forall>x\<^isub>1\<dots>x\<^isub>n. P\<^bsub>ty\<AL>\<^esub>\<^isub>i x\<^isub>i \<and> \<dots> \<and> P\<^bsub>ty\<AL>\<^esub>\<^isub>j x\<^isub>j \<Rightarrow> P\<^bsub>ty\<AL>\<^esub> (C\<^sup>\<alpha> x\<^isub>1 \<dots> x\<^isub>n)"}
\end{center}
\noindent
where the @{text "x\<^isub>i, \<dots>, x\<^isub>j"} are the recursive arguments of @{text "C\<^sup>\<alpha>"}.
Next we lift the permutation operations defined in \eqref{ceqvt} for
the raw term-constructors @{text "C"}. These facts contain, in addition
to the term-constructors, also permutation operations. In order to make the
lifting go through,
we have to know that the permutation operations are respectful
w.r.t.~alpha-equivalence. This amounts to showing that the
alpha-equivalence relations are equivariant, which we already established
in Lemma~\ref{equiv}. As a result we can establish the equations
\begin{equation}\label{calphaeqvt}
@{text "p \<bullet> (C\<^sup>\<alpha> x\<^isub>1 \<dots> x\<^isub>n) = C\<^sup>\<alpha> (p \<bullet> x\<^isub>1) \<dots> (p \<bullet> x\<^isub>n)"}
\end{equation}
\noindent
for our infrastructure. In a similar fashion we can lift the equations
characterising the free-variable functions @{text "fn_ty\<AL>\<^isub>j"} and @{text "fv_bn\<AL>\<^isub>k"}, and the
binding functions @{text "bn\<AL>\<^isub>k"}. The necessary respectfulness lemmas for these
lifting are the properties:
%
\begin{equation}\label{fnresp}
\mbox{%
\begin{tabular}{l}
@{text "x \<approx>ty\<^isub>j y"} implies @{text "fv_ty\<^isub>j x = fv_ty\<^isub>j y"}\\
@{text "x \<approx>ty\<^isub>k y"} implies @{text "fv_bn\<^isub>k x = fv_bn\<^isub>k y"}\\
@{text "x \<approx>ty\<^isub>k y"} implies @{text "bn\<^isub>k x = bn\<^isub>k y"}
\end{tabular}}
\end{equation}
\noindent
which can be established by induction on @{text "\<approx>ty"}. Whereas the first
property is always true by the way we defined the free-variable
functions for types, the second and third do \emph{not} hold in general. There is, in principle,
the possibility that the user defines @{text "bn\<^isub>k"} so that it returns an already bound
variable. Then the third property is just not true. However, our
restrictions imposed on the binding functions
exclude this possibility.
Having the facts \eqref{fnresp} at our disposal, we can lift the
definitions that characterise when two terms of the form
\begin{center}
@{text "C x\<^isub>1 \<dots> x\<^isub>n \<approx>ty C y\<^isub>1 \<dots> y\<^isub>n"}
\end{center}
\noindent
are alpha-equivalent. This gives us conditions when the corresponding
alpha-equated terms are \emph{equal}, namely
\begin{center}
@{text "C\<^sup>\<alpha> x\<^isub>1 \<dots> x\<^isub>n = C\<^sup>\<alpha> y\<^isub>1 \<dots> y\<^isub>n"}
\end{center}
\noindent
We call these conditions as \emph{quasi-injectivity}. Except for one function, which
we have to lift in the next section, we completed
the lifting part of establishing the reasoning infrastructure.
By working now completely on the alpha-equated level, we can first show that
the free-variable functions and binding functions
are equivariant, namely
\begin{center}
\begin{tabular}{rcl}
@{text "p \<bullet> (fv_ty\<^sup>\<alpha> x)"} & $=$ & @{text "fv_ty\<^sup>\<alpha> (p \<bullet> x)"}\\
@{text "p \<bullet> (fv_bn\<^sup>\<alpha> x)"} & $=$ & @{text "fv_bn\<^sup>\<alpha> (p \<bullet> x)"}\\
@{text "p \<bullet> (bn\<^sup>\<alpha> x)"} & $=$ & @{text "bn\<^sup>\<alpha> (p \<bullet> x)"}
\end{tabular}
\end{center}
\noindent
These properties can be established by structural induction over the @{text x}
(using the induction principles we lifted above for the types @{text "ty\<AL>\<^bsub>1..n\<^esub>"}).
Until now we have not yet derived anything about the support of the
alpha-equated terms. This however will be necessary in order to derive
the strong induction principles in the next section.
Using the equivariance properties in \eqref{ceqvt} we can
show for every term-constructor @{text "C\<^sup>\<alpha>"} that
\begin{center}
@{text "(supp x\<^isub>1 \<union> \<dots> supp x\<^isub>n) supports (C\<^sup>\<alpha> x\<^isub>1 \<dots> x\<^isub>n)"}
\end{center}
\noindent
holds. This together with Property~\ref{supportsprop} allows us to show
\begin{center}
@{text "finite (supp x\<^isub>i)"}
\end{center}
\noindent
by a structural induction over @{text "x\<^isub>1, \<dots>, x\<^isub>n"} (whereby @{text "x\<^isub>i"} ranges over the types
@{text "ty\<AL>\<^isub>1 \<dots> ty\<AL>\<^isub>n"}). Lastly, we can show that the support of elements in
@{text "ty\<AL>\<^bsub>1..n\<^esub>"} coincides with @{text "fv_ty\<AL>\<^bsub>1..n\<^esub>"}.
\begin{lemma}
For every @{text "x\<^isub>i"} of type @{text "ty\<AL>\<^bsub>1..n\<^esub>"}, we have that
@{text "supp x\<^isub>i = fv_ty\<AL>\<^isub>i x\<^isub>i"} holds.
\end{lemma}
\begin{proof}
The proof proceeds by structural induction over the @{text "x\<^isub>i"}. In each case
we unfold the definition of @{text "supp"}, move the swapping inside the
term-constructors and the use then quasi-injectivity lemmas in order to complete the
proof. For the abstraction cases we use the facts derived in Theorem~\ref{suppabs}.
\end{proof}
\noindent
To sum up, we can established automatically a reasoning infrastructure
for the types @{text "ty\<AL>\<^isub>1, \<dots>, ty\<AL>\<^isub>n"}
by first lifting definitions from the raw level to the quotient level and
then establish facts about these lifted definitions. All necessary proofs
are generated automatically by custom ML-code. This code can deal with
specifications like the one shown in Figure~\ref{nominalcorehas} for Core-Haskell.
\begin{figure}[t!]
\begin{boxedminipage}{\linewidth}
\small
\begin{tabular}{l}
\isacommand{atom\_decl}~@{text "var"}\\
\isacommand{atom\_decl}~@{text "cvar"}\\
\isacommand{atom\_decl}~@{text "tvar"}\\[1mm]
\isacommand{nominal\_datatype}~@{text "tkind ="}\\
\phantom{$|$}~@{text "KStar"}~$|$~@{text "KFun tkind tkind"}\\
\isacommand{and}~@{text "ckind ="}\\
\phantom{$|$}~@{text "CKSim ty ty"}\\
\isacommand{and}~@{text "ty ="}\\
\phantom{$|$}~@{text "TVar tvar"}~$|$~@{text "T string"}~$|$~@{text "TApp ty ty"}\\
$|$~@{text "TFun string ty_list"}~%
$|$~@{text "TAll tv::tvar tkind ty::ty"} \isacommand{bind}~@{text "tv"}~\isacommand{in}~@{text ty}\\
$|$~@{text "TArr ckind ty"}\\
\isacommand{and}~@{text "ty_lst ="}\\
\phantom{$|$}~@{text "TNil"}~$|$~@{text "TCons ty ty_lst"}\\
\isacommand{and}~@{text "cty ="}\\
\phantom{$|$}~@{text "CVar cvar"}~%
$|$~@{text "C string"}~$|$~@{text "CApp cty cty"}~$|$~@{text "CFun string co_lst"}\\
$|$~@{text "CAll cv::cvar ckind cty::cty"} \isacommand{bind}~@{text "cv"}~\isacommand{in}~@{text cty}\\
$|$~@{text "CArr ckind cty"}~$|$~@{text "CRefl ty"}~$|$~@{text "CSym cty"}~$|$~@{text "CCirc cty cty"}\\
$|$~@{text "CAt cty ty"}~$|$~@{text "CLeft cty"}~$|$~@{text "CRight cty"}~$|$~@{text "CSim cty cty"}\\
$|$~@{text "CRightc cty"}~$|$~@{text "CLeftc cty"}~$|$~@{text "Coerce cty cty"}\\
\isacommand{and}~@{text "co_lst ="}\\
\phantom{$|$}@{text "CNil"}~$|$~@{text "CCons cty co_lst"}\\
\isacommand{and}~@{text "trm ="}\\
\phantom{$|$}~@{text "Var var"}~$|$~@{text "K string"}\\
$|$~@{text "LAM_ty tv::tvar tkind t::trm"} \isacommand{bind}~@{text "tv"}~\isacommand{in}~@{text t}\\
$|$~@{text "LAM_cty cv::cvar ckind t::trm"} \isacommand{bind}~@{text "cv"}~\isacommand{in}~@{text t}\\
$|$~@{text "App_ty trm ty"}~$|$~@{text "App_cty trm cty"}~$|$~@{text "App trm trm"}\\
$|$~@{text "Lam v::var ty t::trm"} \isacommand{bind}~@{text "v"}~\isacommand{in}~@{text t}\\
$|$~@{text "Let x::var ty trm t::trm"} \isacommand{bind}~@{text x}~\isacommand{in}~@{text t}\\
$|$~@{text "Case trm assoc_lst"}~$|$~@{text "Cast trm co"}\\
\isacommand{and}~@{text "assoc_lst ="}\\
\phantom{$|$}~@{text ANil}~%
$|$~@{text "ACons p::pat t::trm assoc_lst"} \isacommand{bind}~@{text "bv p"}~\isacommand{in}~@{text t}\\
\isacommand{and}~@{text "pat ="}\\
\phantom{$|$}~@{text "Kpat string tvtk_lst tvck_lst vt_lst"}\\
\isacommand{and}~@{text "vt_lst ="}\\
\phantom{$|$}~@{text VTNil}~$|$~@{text "VTCons var ty vt_lst"}\\
\isacommand{and}~@{text "tvtk_lst ="}\\
\phantom{$|$}~@{text TVTKNil}~$|$~@{text "TVTKCons tvar tkind tvtk_lst"}\\
\isacommand{and}~@{text "tvck_lst ="}\\
\phantom{$|$}~@{text TVCKNil}~$|$ @{text "TVCKCons cvar ckind tvck_lst"}\\
\isacommand{binder}\\
@{text "bv :: pat \<Rightarrow> atom list"}~\isacommand{and}~%
@{text "bv1 :: vt_lst \<Rightarrow> atom list"}~\isacommand{and}\\
@{text "bv2 :: tvtk_lst \<Rightarrow> atom list"}~\isacommand{and}~%
@{text "bv3 :: tvck_lst \<Rightarrow> atom list"}\\
\isacommand{where}\\
\phantom{$|$}~@{text "bv (K s tvts tvcs vs) = (bv3 tvts) @ (bv2 tvcs) @ (bv1 vs)"}\\
$|$~@{text "bv1 VTNil = []"}\\
$|$~@{text "bv1 (VTCons x ty tl) = (atom x)::(bv1 tl)"}\\
$|$~@{text "bv2 TVTKNil = []"}\\
$|$~@{text "bv2 (TVTKCons a ty tl) = (atom a)::(bv2 tl)"}\\
$|$~@{text "bv3 TVCKNil = []"}\\
$|$~@{text "bv3 (TVCKCons c cty tl) = (atom c)::(bv3 tl)"}\\
\end{tabular}
\end{boxedminipage}
\caption{The nominal datatype declaration for Core-Haskell. For the moment we
do not support nested types; therefore we explicitly have to unfold the
lists @{text "co_lst"}, @{text "assoc_lst"} and so on. This will be improved
in a future version of Nominal Isabelle. Apart from that, the
declaration follows closely the original in Figure~\ref{corehas}. The
point of our work is that having made such a declaration in Nominal Isabelle,
one obtains automatically a reasoning infrastructure for Core-Haskell.
\label{nominalcorehas}}
\end{figure}
*}
section {* Strong Induction Principles *}
text {*
In the previous section we were able to provide induction principles that
allow us to perform structural inductions over alpha-equated terms.
We call such induction principles \emph{weak}, because in case of a term-constructor @{text "C\<^sup>\<alpha> x\<^isub>1 \<dots> x\<^isub>n"},
the induction hypothesis requires us to establish the implication
%
\begin{equation}\label{weakprem}
@{text "\<forall>x\<^isub>1\<dots>x\<^isub>n. P\<^isub>i x\<^isub>i \<and> \<dots> \<and> P\<^isub>j x\<^isub>j \<Rightarrow> P (C\<^sup>\<alpha> x\<^isub>1 \<dots> x\<^isub>n)"}
\end{equation}
\noindent
where the @{text "x\<^isub>i, \<dots>, x\<^isub>j"} are the recursive arguments of @{text "C\<^sup>\<alpha>"}.
The problem with this implication is that in general it is not easy to establish it.
The reason is that we cannot make any assumption about the binders that might be in @{text "C\<^sup>\<alpha>"}
(for example we cannot assume the variable convention for them).
In \cite{UrbanTasson05} we introduced a method for automatically
strengthening weak induction principles for terms containing single
binders. These stronger induction principles allow the user to make additional
assumptions about binders.
These additional assumptions amount to a formal
version of the informal variable convention for binders. A natural question is
whether we can also strengthen the weak induction principles involving
the general binders presented here. We will indeed be able to so, but for this we need an
additional notion for permuting deep binders.
Given a binding function @{text "bn"} we define an auxiliary permutation
operation @{text "_ \<bullet>\<^bsub>bn\<^esub> _"} which permutes only bound arguments in a deep binder.
Assuming a clause of @{text bn} is defined as @{text "bn (C x\<^isub>1 \<dots> x\<^isub>n) = rhs"}, then
we define %
\begin{center}
@{text "p \<bullet>\<^bsub>bn\<^esub> (C x\<^isub>1 \<dots> x\<^isub>n) \<equiv> C y\<^isub>1 \<dots> y\<^isub>n"}
\end{center}
\noindent
with @{text "y\<^isub>i"} determined as follows:
%
\begin{center}
\begin{tabular}{c@ {\hspace{2mm}}p{7cm}}
$\bullet$ & @{text "y\<^isub>i \<equiv> x\<^isub>i"} provided @{text "x\<^isub>i"} does not occur in @{text "rhs"}\\
$\bullet$ & @{text "y\<^isub>i \<equiv> p \<bullet>\<^bsub>bn'\<^esub> x\<^isub>i"} provided @{text "bn' x\<^isub>i"} is in @{text "rhs"}\\
$\bullet$ & @{text "y\<^isub>i \<equiv> p \<bullet> x\<^isub>i"} otherwise
\end{tabular}
\end{center}
\noindent
Using again the quotient package we can lift the @{text "_ \<bullet>\<^bsub>bn\<^esub> _"} function to
alpha-equated terms. We can then prove the following two facts
\begin{lemma}\label{permutebn}
Given a binding function @{text "bn\<^sup>\<alpha>"} then for all @{text p}
{\it i)} @{text "p \<bullet> (bn\<^sup>\<alpha> x) = bn\<^sup>\<alpha> (p \<bullet>\<^bsub>bn\<^esub>\<^sup>\<alpha> x)"} and {\it ii)}
@{text "fv_bn\<^isup>\<alpha> x = fv_bn\<^isup>\<alpha> (p \<bullet>\<^bsub>bn\<^esub>\<^sup>\<alpha> x)"}.
\end{lemma}
\begin{proof}
By induction on @{text x}. The equations follow by simple unfolding
of the definitions.
\end{proof}
\noindent
The first property states that a permutation applied to a binding function is
equivalent to first permuting the binders and then calculating the bound
variables. The second amounts to the fact that permuting the binders has no
effect on the free-variable function. The main point of this permutation
function, however, is that if we have a permutation that is fresh
for the support of an object @{text x}, then we can use this permutation
to rename the binders in @{text x}, without ``changing'' @{text x}. In case of the
@{text "Let"} term-constructor from the example shown
\eqref{letpat} this means for a permutation @{text "r"}:
%
\begin{equation}\label{renaming}
\begin{array}{l}
\mbox{if @{term "supp (Abs_lst (bn p) t\<^isub>2) \<sharp>* r"}}\\
\qquad\mbox{then @{text "Let p t\<^isub>1 t\<^isub>2 = Let (r \<bullet>\<^bsub>bnpat\<^esub> p) t\<^isub>1 (r \<bullet> t\<^isub>2)"}}
\end{array}
\end{equation}
\noindent
This fact will be crucial when establishing the strong induction principles.
In our running example about @{text "Let"}, they state that instead
of establishing the implication
\begin{center}
@{text "\<forall>p t\<^isub>1 t\<^isub>2. P\<^bsub>pat\<^esub> p \<and> P\<^bsub>trm\<^esub> t\<^isub>1 \<and> P\<^bsub>trm\<^esub> t\<^isub>2 \<Rightarrow> P\<^bsub>trm\<^esub> (Let p t\<^isub>1 t\<^isub>2)"}
\end{center}
\noindent
it is sufficient to establish the following implication
%
\begin{equation}\label{strong}
\mbox{\begin{tabular}{l}
@{text "\<forall>p t\<^isub>1 t\<^isub>2 c."}\\
\hspace{5mm}@{text "set (bn p) #\<^sup>* c \<and>"}\\
\hspace{5mm}@{text "(\<forall>d. P\<^bsub>pat\<^esub> d p) \<and> (\<forall>d. P\<^bsub>trm\<^esub> d t\<^isub>1) \<and> (\<forall>d. P\<^bsub>trm\<^esub> d t\<^isub>2)"}\\
\hspace{15mm}@{text "\<Rightarrow> P\<^bsub>trm\<^esub> c (Let p t\<^isub>1 t\<^isub>2)"}
\end{tabular}}
\end{equation}
\noindent
While this implication contains an additional argument, namely @{text c}, and
also additional universal quantifications, it is usually easier to establish.
The reason is that we can make the freshness
assumption @{text "set (bn\<^sup>\<alpha> p) #\<^sup>* c"}, whereby @{text c} can be arbitrarily
chosen by the user as long as it has finite support.
Let us now show how we derive the strong induction principles from the
weak ones. In case of the @{text "Let"}-example we derive by the weak
induction the following two properties
%
\begin{equation}\label{hyps}
@{text "\<forall>q c. P\<^bsub>trm\<^esub> c (q \<bullet> t)"} \hspace{4mm}
@{text "\<forall>q\<^isub>1 q\<^isub>2 c. P\<^bsub>pat\<^esub> (q\<^isub>1 \<bullet>\<^bsub>bn\<^esub> (q\<^isub>2 \<bullet> p))"}
\end{equation}
\noindent
For the @{text Let} term-constructor we therefore have to establish @{text "P\<^bsub>trm\<^esub> c (q \<bullet> Let p t\<^isub>1 t\<^isub>2)"}
assuming \eqref{hyps} as induction hypotheses (the first for @{text t\<^isub>1} and @{text t\<^isub>2}).
By Property~\ref{avoiding} we
obtain a permutation @{text "r"} such that
%
\begin{equation}\label{rprops}
@{term "(r \<bullet> set (bn (q \<bullet> p))) \<sharp>* c "}\hspace{4mm}
@{term "supp (Abs_lst (bn (q \<bullet> p)) (q \<bullet> t\<^isub>2)) \<sharp>* r"}
\end{equation}
\noindent
hold. The latter fact and \eqref{renaming} give us
\begin{center}
\begin{tabular}{l}
@{text "Let (q \<bullet> p) (q \<bullet> t\<^isub>1) (q \<bullet> t\<^isub>2) ="} \\
\hspace{15mm}@{text "Let (r \<bullet>\<^bsub>bn\<^esub> (q \<bullet> p)) (q \<bullet> t\<^isub>1) (r \<bullet> (q \<bullet> t\<^isub>2))"}
\end{tabular}
\end{center}
\noindent
So instead of proving @{text "P\<^bsub>trm\<^esub> c (q \<bullet> Let p t\<^isub>1 t\<^isub>2)"}, we can equally
establish
\begin{center}
@{text "P\<^bsub>trm\<^esub> c (Let (r \<bullet>\<^bsub>bn\<^esub> (q \<bullet> p)) (q \<bullet> t\<^isub>1) (r \<bullet> (q \<bullet> t\<^isub>2)))"}
\end{center}
\noindent
To do so, we will use the implication \eqref{strong} of the strong induction
principle, which requires us to discharge
the following four proof obligations:
\begin{center}
\begin{tabular}{rl}
{\it i)} & @{text "set (bn (r \<bullet>\<^bsub>bn\<^esub> (q \<bullet> p))) #\<^sup>* c"}\\
{\it ii)} & @{text "\<forall>d. P\<^bsub>pat\<^esub> d (r \<bullet>\<^bsub>bn\<^esub> (q \<bullet> p))"}\\
{\it iii)} & @{text "\<forall>d. P\<^bsub>trm\<^esub> d (q \<bullet> t\<^isub>1)"}\\
{\it iv)} & @{text "\<forall>d. P\<^bsub>trm\<^esub> d (r \<bullet> (q \<bullet> t\<^isub>2))"}\\
\end{tabular}
\end{center}
\noindent
The first follows from \eqref{rprops} and Lemma~\ref{permutebn}.{\it i)}; the
others from the induction hypotheses in \eqref{hyps} (in the fourth case
we have to use the fact that @{term "(r \<bullet> (q \<bullet> t\<^isub>2)) = (r + q) \<bullet> t\<^isub>2"}).
Taking now the identity permutation @{text 0} for the permutations in \eqref{hyps},
we can establish our original goals, namely @{text "P\<^bsub>trm\<^esub> c t"} and \mbox{@{text "P\<^bsub>pat\<^esub> c p"}}.
This completes the proof showing that the strong induction principle derives from
the weak induction principle. For the moment we can derive the difficult cases of
the strong induction principles only by hand, but they are very schematic
so that with little effort we can even derive them for
Core-Haskell given in Figure~\ref{nominalcorehas}.
*}
section {* Related Work *}
text {*
To our knowledge, the earliest usage of general binders in a theorem prover
is described in \cite{NaraschewskiNipkow99} about a formalisation of the
algorithm W. This formalisation implements binding in type schemes using a
de-Bruijn indices representation. Since type schemes of W contain only a single
binder, different indices do not refer to different binders (as in the usual
de-Bruijn representation), but to different bound variables. A similar idea
has been recently explored for general binders in the locally nameless
approach to binding \cite{chargueraud09}. There, de-Bruijn indices consist
of two numbers, one referring to the place where a variable is bound and the
other to which variable is bound. The reasoning infrastructure for both
representations of bindings comes for free in theorem provers like Isabelle/HOL or
Coq, as the corresponding term-calculi can be implemented as ``normal''
datatypes. However, in both approaches it seems difficult to achieve our
fine-grained control over the ``semantics'' of bindings (i.e.~whether the
order of binders should matter, or vacuous binders should be taken into
account). To do so, one would require additional predicates that filter out
unwanted terms. Our guess is that such predicates result in rather
intricate formal reasoning.
Another representation technique for binding is higher-order abstract syntax
(HOAS), which for example is implemented in the Twelf system. This representation
technique supports very elegantly many aspects of \emph{single} binding, and
impressive work is in progress that uses HOAS for mechanising the metatheory
of SML~\cite{LeeCraryHarper07}. We are, however, not aware how multiple
binders of SML are represented in this work. Judging from the submitted
Twelf-solution for the POPLmark challenge, HOAS cannot easily deal with
binding constructs where the number of bound variables is not fixed. For
example in the second part of this challenge, @{text "Let"}s involve
patterns that bind multiple variables at once. In such situations, HOAS
representations have to resort to the iterated-single-binders-approach with
all the unwanted consequences when reasoning about the resulting terms.
Two formalisations involving general binders have also been performed in older
versions of Nominal Isabelle (one about Psi-calculi and one about alpgorithm W
\cite{BengtsonParow09, UrbanNipkow09}). Both
use the approach based on iterated single binders. Our experience with
the latter formalisation has been disappointing. The major pain arose from
the need to ``unbind'' variables. This can be done in one step with our
general binders, but needs a cumbersome
iteration with single binders. The resulting formal reasoning turned out to
be rather unpleasant. The hope is that the extension presented in this paper
is a substantial improvement.
The most closely related work to the one presented here is the Ott-tool
\cite{ott-jfp} and the C$\alpha$ml language \cite{Pottier06}. Ott is a nifty
front-end for creating \LaTeX{} documents from specifications of
term-calculi involving general binders. For a subset of the specifications,
Ott can also generate theorem prover code using a raw representation of
terms, and in Coq also a locally nameless representation. The developers of
this tool have also put forward (on paper) a definition for
alpha-equivalence of terms that can be specified in Ott. This definition is
rather different from ours, not using any nominal techniques. To our
knowledge there is also no concrete mathematical result concerning this
notion of alpha-equivalence. A definition for the notion of free variables
in a term are work in progress in Ott.
Although we were heavily inspired by their syntax,
their definition of alpha-equivalence is unsuitable for our extension of
Nominal Isabelle. First, it is far too complicated to be a basis for
automated proofs implemented on the ML-level of Isabelle/HOL. Second, it
covers cases of binders depending on other binders, which just do not make
sense for our alpha-equated terms. Third, it allows empty types that have no
meaning in a HOL-based theorem prover. We also had to generalise slightly their
binding clauses. In Ott you specify binding clauses with a single body; we
allow more than one. We have to do this, because this makes a difference
for our notion of alpha-equivalence in case of \isacommand{bind\_set} and
\isacommand{bind\_res}. This makes
\begin{center}
\begin{tabular}{@ {}ll@ {}}
@{text "Foo\<^isub>1 xs::name fset x::name y::name"} &
\isacommand{bind\_set} @{text "xs"} \isacommand{in} @{text "x y"}\\
@{text "Foo\<^isub>2 xs::name fset x::name y::name"} &
\isacommand{bind\_set} @{text "xs"} \isacommand{in} @{text "x"},
\isacommand{bind\_set} @{text "xs"} \isacommand{in} @{text "y"}\\
\end{tabular}
\end{center}
\noindent
behave differently. To see this, consider the following equations
\begin{center}
\begin{tabular}{c}
@{term "Abs_print [a,b] (a, b) = Abs_print [a,b] (b, a)"}\\
@{term "Abs_print [a,b] (a, b) = Abs_print [a,b] (a, b)"}\\
\end{tabular}
\end{center}
\noindent
The interesting point is that they do not imply
\begin{center}
\begin{tabular}{c}
@{term "Abs_print [a,b] ((a, b), (a, b)) = Abs_print [a,b] ((a, b), (b, a))"}\\
\end{tabular}
\end{center}
Because of how we set up our definitions, we had to impose restrictions,
like a single binding function for a deep binder, that are not present in Ott. Our
expectation is that we can still cover many interesting term-calculi from
programming language research, for example Core-Haskell. For providing support
of neat features in Ott, such as subgrammars, the existing datatype
infrastructure in Isabelle/HOL is unfortunately not powerful enough. On the
other hand, we are not aware that Ott can make the distinction between our
three different binding modes.
Pottier presents in \cite{Pottier06} a language, called C$\alpha$ml, for
representing terms with general binders inside OCaml. This language is
implemented as a front-end that can be translated to OCaml with a help of
a library. He presents a type-system in which the scope of general binders
can be indicated with some special constructs, written @{text "inner"} and
@{text "outer"}. With this, it seems, our and his specifications can be
inter-translated, but we have not proved this. However, we believe the
binding specifications in the style of Ott are a more natural way for
representing terms involving bindings. Pottier gives a definition for
alpha-equivalence, which is similar to our binding mod \isacommand{bind}.
Although he uses also a permutation in case of abstractions, his
definition is rather different from ours. He proves that his notion
of alpha-equivalence is indeed a equivalence relation, but a complete
reasoning infrastructure is well beyond the purposes of his language.
*}
section {* Conclusion *}
text {*
We have presented an extension of Nominal Isabelle for deriving
automatically a convenient reasoning infrastructure that can deal with
general binders, that is term-constructors binding multiple variables at
once. This extension has been tried out with the Core-Haskell
term-calculus. For such general binders, we can also extend
earlier work that derives strong induction principles which have the usual
variable convention already built in. For the moment we can do so only with manual help,
but future work will automate them completely. The code underlying the presented
extension will become part of the Isabelle distribution, but for the moment
it can be downloaded from the Mercurial repository linked at
\href{http://isabelle.in.tum.de/nominal/download}
{http://isabelle.in.tum.de/nominal/download}.
We have left out a discussion about how functions can be defined over
alpha-equated terms that involve general binders. In earlier versions of Nominal
Isabelle \cite{UrbanBerghofer06} this turned out to be a thorny issue. We
hope to do better this time by using the function package that has recently
been implemented in Isabelle/HOL and also by restricting function
definitions to equivariant functions (for such functions it is possible to
provide more automation).
There are some restrictions we imposed in this paper, that we would like to lift in
future work. One is the exclusion of nested datatype definitions. Nested
datatype definitions allow one to specify, for instance, the function kinds
in Core-Haskell as @{text "TFun string (ty list)"} instead of the unfolded
version @{text "TFun string ty_list"} in Figure~\ref{nominalcorehas}. For
them we need a more clever implementation than we have at the moment.
More interesting line of investigation is whether we can go beyond the
simple-minded form for binding functions that we adopted from Ott. At the moment, binding
functions can only return the empty set, a singleton atom set or unions
of atom sets (similarly for lists). It remains to be seen whether
properties like \eqref{bnprop} provide us with means to support more interesting
binding functions.
We have also not yet played with other binding modes. For example we can
imagine that there is need for a binding mode \isacommand{bind\_\#list} with
an associated abstraction of the form
%
\begin{center}
@{term "Abs_dist as x"}
\end{center}
\noindent
where instead of lists, we abstract lists of distinct elements.
Once we feel confident about such binding modes, our implementation
can be easily extended to accommodate them.
\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 of the work on the Ott-tool. We
also thank Stephanie Weirich for suggesting to separate the subgrammars
of kinds and types in our Core-Haskell example.
*}
(*<*)
end
(*>*)