356 lines
20 KiB
TeX
356 lines
20 KiB
TeX
\section{Soundness}
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\newcommand{\CC}{\text{CC}}
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\begin{theorem}\label{testenv-theorem}
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Type inference produces a correctly typed program.
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\begin{description}
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\item[If] $\fjtypeinference(\mv{\Pi}, \texttt{class}\ \exptype{C}{\ol{X}
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\triangleleft \ol{N}} \triangleleft \type{N}\ \{ \overline{\type{T} \ f};\ \ol{M} \}) = \mtypeEnvironment{}'$
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\item[Then] $\texttt{class}\ \exptype{C}{\ol{X}
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\triangleleft \ol{N}} \triangleleft \type{N}\ \{ \overline{\type{T} \ f};\ \ol{M} \} \text{ok}$,
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with $\ol{M} = $
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\end{description}
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\end{theorem}
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\begin{lemma}{Soundness:}
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\unify{}'s type solutions for a constraint set generated by $\typeExpr{}$ are correct.
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\begin{description}
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\item[if] $\typeExpr{}(\mtypeEnvironment{}, \texttt{e}, \tv{a}) = C$
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, with $C = \set{ \overline{ \type{S} \lessdot \type{T} } \cup \overline{ \type{S'} \lessdot \type{T'} } }$
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and $fv(\overline{ \type{S} }) = \emptyset$, $fv(\overline{ \type{T} }) = \emptyset$
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and $\vdash \ol{L} : \mtypeEnvironment{}$
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and $\Gamma \subseteq \mtypeEnvironment{}$
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\item[given] a $(\Delta, \sigma)$ with $\Delta \vdash \overline{\sigma(\type{S}) <: \sigma(\type{T})}$
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and there exists a $\Delta'$ with $\Delta, \Delta' \vdash \overline{\CC{}(\sigma(\type{S'})) <: \sigma(\type{T'})}$
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%and $\Delta \vdash \sigma(\type{S}), \sigma(\type{T})\ \ok$
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\item[then] $\Delta|\Gamma \vdash \texttt{e} : \sigma(\tv{a})$
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\end{description}
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\end{lemma}
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Regular type placeholders represent type annotations.
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These are the only types a \wildFJ{} program needs to be correctly typed.
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The type placeholders flagged as wildcard placeholders are intermediate types
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used in let statements and as type parameters for generic method calls.
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%Unify needs to return S aswell and guarantee that the \Delta' environment are the wildcards
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% only used inside the constraint the wildcard variable occurs
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% should Unify also return the \Delta' environment? Otherwise the bounds of free wildcard variables are lost
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% Or is it possible to deduct the right \ol{S} directly from the types in the normal TPHs?
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\textit{Proof:}
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By structural induction over the expression $\texttt{e}$.
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\begin{description}
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\item[$\texttt{e}.\texttt{f}$] $\Delta|\Gamma \vdash \texttt{e} : \sigma(\tv{r})$ by assumption.
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$\Delta', \Delta\vdash \CC{}(\sigma(\tv{r})) <: \sigma(\exptype{C}{\overline{\wtv{a}}})$ by premise.
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Let $\sigma(\tv{r}) = \wcNtype{\Delta'}{N}$.
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%TODO: Proof this: The reason is that r is not flagged with a ? and we only add free variables here, that are in \Delta
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Then $\text{fv}(\wcNtype{\Delta'}{N}) \subseteq \text{dom}(\Delta)$
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and therefore $\text{fv}(\type{N}) \subseteq \text{dom}(\Delta, \Delta')$.
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This implies $\Delta, \Delta' \vdash \type{U}_i <: \type{S}$,
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because of the constraint $[\overline{\wtv{a}}/\ol{X}]\type{T} \lessdot \tv{a}$
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and $\textit{fields}(\sigma(\exptype{C}{\overline{\wtv{a}}})) = \sigma([\overline{\wtv{a}}/\ol{X}]\type{T})$
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and $\text{fv}(\type{U}_i) \subseteq \text{fv}(\type{N})$ by definition of $\textit{fields}$.
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% Set $\sigma(\tv{r}) = \type{T} = \wcNtype{\Delta'}{N}$.
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% Then $\Delta | \Gamma \vdash e : \type{T}$ by assumption. $\Delta \vdash \type{T} <: \wcNtype{\Delta'}{N}$ by S-Refl.
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% $\textit{fields}(\type{N}) = \overline{\type{U}\ f}$ by definition.
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% $\Delta, \Delta' \vdash \type{U}_i <: \type{S}$ by premise.
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$\Delta \vdash \type{S}, \wcNtype{\Delta'}{N} \ \ok$ %TODO
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%Easy, because unify only generates substitutions for normal type placeholders which are OK
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\item[$\texttt{e}.\texttt{m}(\ol{e})$]
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\begin{gather}
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\label{sp:1}
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\Delta, \Delta', \overline{\Delta} \vdash \ol{N} <: [\ol{S}/\ol{X}]\ol{U} \\
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\label{sp:2}
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\Delta, \Delta', \overline{\Delta} \vdash \ol{S} <: [\ol{S}/\ol{X}]\ol{U'} \\
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\Delta, \Delta', \overline{\Delta} \vdash \ol{S} \ \ok \\
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\label{sp:4}
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\Delta \vdash \ol{T} <: \ol{\wcNtype{\Delta}{N}} \\
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\end{gather}
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\ref{sp:1} is guaranteed by the constraints $\ol{r} \lessdot \ol{T}$.
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$\ol{\tv{r}} \lessdot \ol{T}$ says that there is a $\Delta'$ with
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$\Delta, \Delta' \vdash CC(\sigma(\tv{r})) <: \type{T}$.
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If $\overline{\sigma(\tv{r}) = \wcNtype{\Delta}{N}}$
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then \ref{sp:1} due to lemma \ref{lemma:unifyCC}.
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% X.List<X> <. List<a?>
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% set \sigma(r) = X.N
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% If there exists a subtype r <: T, then r <: X.N, N <: T with X.N == r
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% If CC(X.N) <: T, then N <: T
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% R <: X.N due to S-Refl.
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% \Delta, X \vdash N <: T due to assumption and S-Exists.
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% What if X.N <: Y (where is this Y coming from?)
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% %TODO:
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% We have to show, that all constraints using the same wtv's C = X1.T1 <. X2.T2 ,...
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% with a? \in T1, T2, ....
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% then \Delta' \subseteq X1, X2, ... for \Delta, \Delta \vdash sigma(T1) <: sigma(X2.T2), ....
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% when only wildcards in the top level domain from the constraints involving the a? variables are involved
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% then they are all contained in $\Delta, \Delta', \overline{Delta}$, which makes S ok true! (does it?)
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% and due to r <. T, the judgements N <: [S/X]U and T_r <: \Delta.N are true!
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%TODO for Unify:
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We have to make sure, that wildcards are correctly saved in the Wildcard environment
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% If $\ol{\tv{r}} \lessdot \ol{T}$ how can we say that \Delta' only uses
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\end{description}
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\begin{lemma}
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\begin{description}
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\item[If] $(\sigma, \Delta) = \unify{}( C )$
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\item[Then] $\forall x \in \set{\type{S} \mid \type{S} \in C, \text{fv}(\type{S}) = \emptyset }: \text{fv}(\sigma(\type{S})) \subseteq \text{dom}(\Delta)$
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\end{description}
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\end{lemma}
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\textit{Proof:} by induction over the \unify{} algorithm.
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A unifier $\sigma(\tv{a}) = \type{T}$, where the type variable $\tv{a}$ is not flagged as a wildcard will always hold
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$\text{fv}(\type{T}) \subseteq \text{dom}(\Delta)$.
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%UNIFY fails when there are free variables on the right side of a a =. T constraint
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\begin{lemma}
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\unify{} only produces well-formed type substitutions.
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% Could this be done via the GenSigma rule? just check if substitutions are well-formed?
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% L <. U
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% not possible because well-formedness depends on the input, which is not checked (should we check it? is that possible?)
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%only type variables have to be OK
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% TODO: Problem: how to proof \ol{S} ok for T-Call.
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% a <. List<x?> --> here the type substituted for x? must be ok. Proof that by changing Unify to not return type insertions for this type and assume there is a correct substitution for x
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\end{lemma}
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% %This lemma can be used to proof Normalize rule!
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% \begin{lemma}\label{lemma:wildcardReplacement}
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% Wildcards with the same upper and lower bound can be replaced by their bounds without breaking subtype relations.
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% \begin{description}
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% \item[If] $\Delta \cup \set{\wildcard{X}{\type{U}}{\type{L}}} \vdash \type{T} <: \type{S}$
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% \item[and] $\type{U} = \type{L}$ and $\text{fv}(\type{U}) = \emptyset$
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% \item[Then] $\Delta \vdash [\type{U}/\type{X}]\type{T} <: [\type{U}/\type{X}]\type{S}$
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% \end{description}
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% \end{lemma}
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% \textit{Proof:} %TODO
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% %By structural induction over the subtype relation
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% %S-Refl: by assumption L <: L implies [\type{U}/\type{X}]L
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% \begin{lemma}\label{lemma:noAdditionalFV}
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% Type solution $\sigma$ does not add additional free variables.
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% \begin{description}
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% \item[If] $(\sigma, \Delta) = \unify{}( \overline{ \type{S} \lessdot \type{T} } \cup C)$
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% \item[and] $fv(\overline{ \type{S} }) = \emptyset$, $fv(\overline{ \type{T} }) = \emptyset$
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% \item[Then] $fv(\sigma(\overline{ \type{S} })) = \emptyset, fv(\sigma(\overline{ \type{T} })) = \emptyset$
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% \end{description}
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% \end{lemma}
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% \begin{lemma}
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% % https://math.libretexts.org/Bookshelves/Mathematical_Logic_and_Proof/Book%3A_Mathematical_Reasoning__Writing_and_Proof_(Sundstrom)/07%3A_Equivalence_Relations/7.03%3A_Equivalence_Classes
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% $\doteq$ is an equivalence relation on a constraint set $C$.
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% \begin{description}
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% \item[If] $\Delta \vdash \overline{\type{T} <: \type{S}}$ and $\type{T} \doteq \type{S}$
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% \item[Then] $\Delta \vdash [\type{T}/\type{S}](\overline{\type{T} <: \type{S}})$
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% \end{description}
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% \end{lemma}
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\begin{lemma} \label{lemma:unifyCC}
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\begin{description}
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\item[If] $(\sigma, \Delta) = \unify{}( C \cup \set{ \overline{\wcNtype{\Delta}{N} \lessdot \type{T} } } )$
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\item[and] $\text{fv}(\ol{T}) \notin C$, $\text{fv}(\ol{\wcNtype{\Delta}{N}}) \notin C$
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\item[Then] $\Delta, \overline{\Delta} \vdash \overline{ \sigma(\type{N}) <: \sigma(\type{T}) }$
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\end{description}
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% Wildcard place holders only use free variables occuring in the same constraint as them.
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% \begin{description}
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% \item[If] $(\sigma, \Delta) = \unify{}( C )$
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% \item[Then] $\text{fv}(\sigma(\tv{a})) \subseteq \set{ \text{fv}(\type{T}) \cup \text{fv}(\type{S}) \mid \tv{a} \in \text{tph}(\type{T},\type{S}), (\type{T} \lessdot \type{S}) \in C }$
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% \end{description}
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\end{lemma}
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\textit{Proof:}
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%TODO: is it true even?
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\begin{lemma}
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The \unify{} algorithm only produces correct output for constraints not containing free variables.
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\begin{description}
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\item[If] $(\sigma, \Delta) = \unify{}( \overline{ \type{S} \lessdot \type{T} } \cup \overline{ \type{S'} \lessdot \type{T'} } )$ %\cup \overline{ \type{S} \doteq \type{S'} })$
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\item[and] $fv(\overline{ \type{S} }) = \emptyset$, $fv(\overline{ \type{T} }) = \emptyset$
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\item[Then] there exists a $\Delta'$ with:
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$\Delta \vdash \overline{\sigma(\type{S}) <: \sigma(\type{T})}$
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and $\Delta, \Delta' \vdash \overline{\CC{}(\sigma(\type{S'})) <: \sigma(\type{T'})}$
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% and $\sigma(\type{T'}) = \sigma(\type{T'})$.
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The function $\CC{}$ is given as $\CC{}(\wcNtype{\Delta}{N}) = \type{N}$
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% Unify cannot guarantee that only wildcards declared on the left side are used. Example input:
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% List<b?> <. List<b?>
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% b? =. X
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% but if the left side is a normal type (no fv's) then this can be guaranteed (maybe?)
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%String <. X_String % here X is not in the environment
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% What is the problem?
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% a? can only use wildcards, which are declared in \Delta or on the left side
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% Otherwise constraints of the form a <. T cannot ensure soundness!
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% We could say that a? only gets free variables hold by the constraints a? appears in
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% a constraint a <. T means that there are free variables
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\end{description}
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\end{lemma}
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% \begin{lemma} % a lemma where we distinguis between free variable on the left or the right side of a constraint (not needed anymore)
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% The \unify{} algorithm only produces correct output for constraints not containing free variables.
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% \begin{description}
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% \item[If] $(\sigma, \Delta) = \unify{}( \overline{ \type{S} \lessdot \type{T} } \cup \overline{ \wcNtype{\Delta'}{N} \lessdot \type{T'} } \cup \overline{ \type{S'} \lessdot \wcNtype{\Delta'}{N'} })$
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% \item[and] $fv(\overline{ \type{S} }) = \emptyset, fv(\overline{ \type{T} }) = \emptyset, fv(\overline{ \type{T'} }) = \emptyset, fv(\overline{ \type{S'} }) = \emptyset$
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% \item[Then] $\Delta \vdash \overline{\sigma(\type{S}) <: \sigma(\type{T})}$
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% and $\Delta, \Delta' \vdash \overline{\sigma(\type{N}) <: \sigma(\type{T'})}, \overline{\sigma(\type{S'}) <: \sigma(\type{N'})}$
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% %TODO: Rephrase (\Delta' is used three times!)
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% %The function $\CC{}$ is given as $\CC{}(\wcNtype{\Delta}{N}) = \type{N}$
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% \end{description}
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% \end{lemma}
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\textit{Proof:}
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%(we are going backwards over the algorithm)
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%first we have to determine the \Delta'' -> it's only the wildcards which are free in N
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% during this proof we can use Delta'' as we like
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For every step in the \unify{} algorithm:
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Assuming the unifier $\sigma$ is correct for a constraint set $C'$, the unifier is also correct for the
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constraint set $C$ before the transformation.
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% \begin{description}
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% \item[Assumption:] $\unify{}(C) = (\Delta, \sigma)$, with $\Delta \vdash \sigma(C)$
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% \item[Induction step:] For every case $C'$ which can be transformed to $C$ we have to show $\Delta \vdash \sigma(C')$
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% \end{description}
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\unify{} terminates with $C = \emptyset$ for which the preposition holds:
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$\Delta \vdash \sigma(\emptyset)$
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We now show that for every transformation of a constraint set $C$ to a constraint set $C'$
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the preposition holds for $C$ using the assumption that it holds for $C'$ :
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$\Delta \vdash \sigma(C') \implies \Delta \vdash \sigma(C)$
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\begin{description}
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\item[AddDelta] $C$ is not changed
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\item[GenDelta] by definition, S-Var-Left, and S-Trans %The generated type variable is unique due to the solved form property. %and the solved form property (no $\tv{a}$ in $C$)
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\item[GenSigma] by definition and S-Refl.
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% holds for $\set{\tv{a} \doteq \type{G}}$ by definition.
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% Holds for $C$ by assumption and because $\tv{a} \notin C$ by solved form definition ($[\type{G}/\tv{a}]C = C$).
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\item[Ground] Assumption and S-Bot
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\item[Sub-Elim] Assumption and S-Refl
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\item[Force] by assumption and $\rwildcard{X} = \type{U}$ %TODO: step 5 should remove all X^T_T with T (make wildcards with same upper and lower bounds to normal types)
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\item[Raise] Assumption, S-Trans
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\item[Settle] Assumption, S-Trans
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\item[Super] S-Extends ($\vdash \wctype{\Delta}{C}{\ol{T}} <: \wctype{\Delta}{D}{[\ol{T}/\ol{X}]\ol{N}}$), S-Trans
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\item[\generalizeRule{}] by Assumption, because $C \subset C'$
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\item[Adapt] Assumption, S-Extends, S-Trans
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\item[Adopt] Assumption, because $C \subset C'$
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%\item[Capture, Reduce] are always applied together. We have to destinct between two cases:
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\item[Capture]
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If $\text{fv}(\wctype{\Delta}{C}{\ol{T}}) = \emptyset$ the preposition holds by Assumption and S-Exists.
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If $\text{fv}(\wctype{\Delta}{C}{\ol{T}}) \neq \emptyset$ the preposition holds because
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$\Delta, \Delta' \vdash [\ol{\rwildcard{C}}/\ol{\rwildcard{B}}] \exptype{C}{\ol{S}} <: \wctype{\Delta}{C}{\ol{T}}$ $\implies$
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$\Delta' \Delta \vdash \text{CC}(\sigma(\wctype{\overline{\wildcard{B}{\type{U'}}{\type{L'}}}}{C}{\ol{S}})) <: \sigma(\wctype{\Delta}{C}{\ol{T}})$
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%we have to show $\Delta' \vdash \text{CC}(\sigma(\wctype{\overline{\wildcard{B}{\type{U'}}{\type{L'}}}}{C}{\ol{S}})) <: \sigma(\wctype{\Delta}{C}{\ol{T}})$,
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%which holds by assumption with $\Delta'$ chosen in a way that $\text{fv}(\exptype{C}{\ol{S}}) \subseteq \Delta'$. The variables $\ol{C}$ in $\ol{S}$ can be renamed to $\ol{B}$, because $\ol{C}$ are fresh.
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% If $\text{fv}(\wctype{\Delta}{C}{\ol{T}}) \neq \emptyset$
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% we have to show $\set{\overline{\wildcard{B}{\type{U'}}{\type{L'}}}}, \Delta \vdash \sigma(\exptype{C}{\ol{S}}) <: \sigma(\wctype{\Delta}{C}{\ol{T}})$,
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% which holds by assumption with $\Delta'$ chosen in a way that $\text{fv}(\exptype{C}{\ol{S}}) \subseteq \Delta'$.
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% The variables $\ol{C}$ in $\ol{S}$ can be renamed to $\ol{B}$, because $\ol{C}$ are fresh.
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%If $\text{fv}(\wctype{\Delta}{C}{\ol{T}}) \neq \emptyset$ we have to show $\Delta \vdash \sigma(\exptype{C}{\ol{S}}) <: \sigma(\wctype{\Delta}{C}{\ol{T}})$.
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%$\Delta \vdash \sigma([\ol{C}/\ol{B}]\exptype{C}{\ol{S}}) <: \sigma(\wctype{\Delta}{C}{\ol{T}})$ holds by assumption and
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%the variables $\ol{B}$ in $\ol{S}$ can be renamed to $\ol{C}$, because $\ol{C} \notin \ol{S}$ ($\ol{C}$ are fresh).
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%The assumption implies $\text{fv}(\ol{S}) \subseteq \text{dom}(\Delta \cup \set{\overline{\wildcard{C}{\type{U'}}{\type{L'}}}})$
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%, which implies $\text{fv}(\ol{S}) \subseteq \text{dom}(\Delta \cup \set{\overline{\wildcard{B}{\type{U'}}{\type{L'}}}})$.
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%We are doing a capture conversion. If $\type{T}$ does not contain free variables, this does not affect the subtype relation.
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\item[Reduce] %Assumption and S-Exists.
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% Three different cases of the constraint $\exptype{C}{\ol{S}} \lessdot \wctype{\overline{\wildcard{A}{\type{U}}{\type{L}}}}{C}{\ol{T}}$:
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% \begin{description}
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% \item[$\text{fv}(\exptype{C}{\ol{S}}) = \emptyset, \text{fv}(\wctype{\overline{\wildcard{A}{\type{U}}{\type{L}}}}{C}{\ol{T}})$:]
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% the preposition holds by Assumption and S-Exists.
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% \item[$\text{fv}(\wctype{\overline{\wildcard{A}{\type{U}}{\type{L}}}}{C}{\ol{T}}) = \emptyset$:]
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% then $\text{fv}(\exptype{C}{\ol{S}}) \subseteq \Delta'$ with $\Delta' = \overline{\wildcard{A}{\type{U}}{\type{L}}}$
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% \item[$\text{fv}(\exptype{C}{\ol{S}}) = \emptyset$] $\Delta' = \emptyset$
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% \end{description}
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% List<X> <. Y.List<Y>, free variables are either in
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If $\text{fv}(\exptype{C}{\ol{S}}) = \text{fv}(\wctype{\overline{\wildcard{A}{\type{U}}{\type{L}}}}{C}{\ol{T}}) = \emptyset$ the preposition holds by Assumption and S-Exists.
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Otherwise
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$\Delta' \vdash \CC{}(\sigma(\exptype{C}{\ol{S}})) <: \sigma(\wctype{\overline{\wildcard{A}{\type{U}}{\type{L}}}}{C}{\ol{T}})$
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holds with any $\Delta'$ so that $(\text{fv}(\exptype{C}{\ol{S}}) \cup \text{fv}(\wctype{\overline{\wildcard{A}{\type{U}}{\type{L}}}}{C}{\ol{T}}) ) \subseteq \text{dom}(\Delta') $.
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\item[Match] Assumption, S-Trans
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\item[Trim] Assumption and S-Exists
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\item[Remove] $C$ is not changed
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\item[Circle] S-Refl
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\item[Swap] by definition
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\item[Erase] S-Refl
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\item[Equals] by definition
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% \item[Reduce]
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% The renaming from $\rwildcard{C}$ to $\rwildcard{B}$ is not a problem. It's allowed to rename wildcards inside a type.
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% Removing $\rwildcard{C}$ from the environment does not change anything because it was freshly generated and is used nowhere.
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% The rest follows directly from S-Exists.
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% We can say: $\text{fv}(\wctype{\overline{\wildcard{A}{\type{U}}{\type{L}}}}{C}{\ol{T}}) = \emptyset$,
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% because the input to the \unify{} algorithm has no free type variables and we never substitute a type with free type variables
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% and none of the other steps of the algorithm generates a $\lessdot$ constraint containing free type variables on the right side. %TODO: proof
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% $\text{fv}(\ol{T}) \subseteq \text{dom}(\wildcardEnv \cup \overline{\wildcard{B}{\type{U'}}{\type{L'}}})$
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% TODO: The capture conversion has to be when substituting a $\wtv{a}$ variable. Then we have to rename!
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% %Lets first try it without the capture conversion. And involve the wtvs in the second step
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% % The algorithm works by never substituting wildcards
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% \unify{} cannot guarantee the premise $\text{dom}(\Delta, \Delta') \cap \text{fv}(\wcNtype{\overline{\wildcard{X}{\type{U}}{\type{L}}}}{N}) = \emptyset$.
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% We loosen the soundness requirements and allow a arbitrary environment $\Delta''$ to be added to the right side of the subtype relation.
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% This is still sufficient to proof soundness for the whole algorithm.
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% We show that the need for the additional environment $\Delta''$ can be satisfied by a let statement.
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% \begin{itemize}
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% \item $\Delta, \Delta' \vdash [\ol{T}/\ol{X}]\ol{L} <: \ol{T}$: S-Exists
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% \item $\Delta, \Delta' \vdash \ol{T} <: [\ol{T}/\ol{X}]\ol{U}$: S-Exists
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% \item $\textit{fv}(\ol{T}) \subseteq \text{dom}{\Delta, \Delta'}$: $\wildcardEnv$ holds all variables %TODO: Proof
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% \end{itemize}
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\item[Normalize] Assumption and lemma 5 \emph{substitution preserves subtyping}.%\ref{lemma:wildcardReplacement}. (Or Lemma 5 from the wildcard paper. \emph{substitution preserves subtyping})
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% The GenSigma step replaces both sides of $\rwildcard{A} \doteq \rwildcard{B}$ with the upper bound $\type{U}$.
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% This works for every constraint, whether it contains free variables or not.
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% It does not add to free variables of constraints because the upper bound does not contain any.
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The GenSigma and Gen Delta steps remove Wildcards which have the same upper and lower bound.
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$\rwildcard{A},\rwildcard{B} \notin \sigma(C)$
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% sigma(T) = sigma(U) we have to show that T = U means \Delta \vdash [T/U]C \implies \Delta \vdash [U/T]C
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% the constraints L <. U, U <. L lead to L =. U
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%If L is List<X> with X being free wildcard
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%then U <. L will fail if U is type variable
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% this is because bounds never contain free variables (is that true?)
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%This type contains free variables when A is replaced by an CC wildcard
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%This must fail:
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\begin{verbatim}
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<A> A m(List<? extends List<A>> l, A)
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m(List<List<? super String>> l, "hi")
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\end{verbatim}
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%This fails because of Equals rule (TODO: proof)
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\item[Tame] same reasoning as Normalize
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\item[Bot] S-Bot
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\item[Pit] S-Bot
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\item[Upper] S-Trans and S-VarLeft
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\item[Lower] S-Trans and S-VerRight
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\item[Subst-WC] by S-Refl
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\item[Subst]
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$\sigma(C \cup \set{\tv{a} \doteq \type{T}}) = \sigma([\type{T}/\tv{a}]C \cup \set{\tv{a} \doteq \type{T}})$
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and
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$\sigma(\wildcardEnv) = \sigma([\type{T}/\tv{a}]\wildcardEnv)$
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\item[Subst-WC]
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%Proof by Lemma 5 \emph{Type substitution preserves subtyping} from \cite{WildcardsNeedWitnessProtection}.
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Same as Subst
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\end{description}
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